Lecture Notes, single file, Tuesday, February 18th ================================================== - Intro to CS4400/CS5400 - Intro to Programming Languages - Intro to Racket - Side-note: "Goto Statement Considered Harmful" - Quick Intro to Racket - Lists & Recursion - Some Style - Tail calls - Note on Types - Side-note: Names are important - BNF, Grammars, the AE Language - Simple Parsing - The `match` Form - The `define-type` Form - The `cases` Form - Semantics (= Evaluation) - Side-note: Compositionality - Implementing an Evaluator - Implementing The AE Language - Intro to Typed Racket - Bindings & Substitution - WAE: Adding Bindings to AE - Evaluation of `with` - Formal Specs - Lazy vs Eager Evaluation - de Bruijn Indexes - Functions & Function Values - Implementing First Class Functions - The FLANG Language - Introducing Racket's `lambda` - Using Functions as Objects - Using `define-type` for new "type aliases" - Currying - Using Higher-Order & Anonymous Functions - Side-note: "Point-Free" combinators - This is not Runtime Code Generation - Substitution Caches - Implementation of Cache Functionality - Formal Rules for Cached Substitutions - Evaluating with Substitution Caches - Dynamic and Lexical Scopes - Dynamic versus Lexical Scope - Implementing Lexical Scope: Closures and Environments - Fixing an Overlooked Bug - Lexical Scope using Racket Closures - More Closures (on both levels) - Types of Evaluators - Feature Embedding - Recursion, Recursion, Recursion - Recursion without the Magic - The Core of `make-recursive` - Denotational Explanation of Recursion - The Y Combinator - The main property of Y - Yet another explanation for Y - Typing the Y Combinator - Lambda Calculus --- Schlac - Church Numerals - More Encodings - Alternative Church Encoding - Recursive Environments - Recursion: Racket's `letrec` - Implementing Recursion using `letrec` - Implementing `rec` Using Cyclic Structures - Boxes and Mutation - Types for Boxes - `Boxof`'s Lack of Subtyping - Implementing a Circular Environment - Variable Mutation - State and Environments - Implementing Objects with State ------------------------------------------------------------------------ # Intro to CS4400/CS5400 [Tuesday, January 7th] * General plan for how the course will go. * Administrative stuff. (Mostly going over the web pages.) ___`https://pl.barzilay.org/`___ ------------------------------------------------------------------------ # Intro to Programming Languages [Tuesday, January 7th] > [PLAI §1] * Why should we care about programming languages? (Any examples of big projects *without* a little language?) * What defines a language? - syntax - semantics - libraries (runtime) - idioms (community) * How important is each of these? - libraries give you the run-time support, not an important part of the language itself. (BTW, the line between "a library" and "part of the language" is less obvious than it seems.) - idioms originate from both language design and culture. They are often misleading. For example, JavaScript programmers will often write: function explorer_move() { doThis(); } function mozilla_move() { doThat(); } if (isExplorer) document.onmousemove = explorer_move; else document.onmousemove = mozilla_move; or if (isExplorer) document.onmousemove = function() { doThis(); }; else document.onmousemove = function() { doThat(); }; or document.onmousemove = isExplorer ? function() { ... } : function() { ... }; or document.onmousemove = isExplorer ? () => { doThis(); } : () => { doThat(); }; or document.onmousemove = isExplorer ? doThis : doThat; How many JavaScript programmers will know what this does: function foo(n) { return function(m) { return m+n; }; } or these: n => m => m+n; (x,y) => s => s(x,y); ([Real example](https://www.websocket.org/echo.html)) Yet another example: let x = ""; while (foo()) x += whatever(); How would you *expect* this code perform? How do you think it does in the reality of many uses of JS by people who are not really programmers? - Compare: - `a[25]+5` (Java: exception) - `(+ (vector-ref a 25) 5)` (Racket: exception) - `a[25]+5` (JavaScript: exception (or NaN)) - `a[25]+5` (Python: exception) - `$a[25]+5` (Perl: 5) - `a[25]+5` (C: ___BOOM___) -> syntax is mostly in the cosmetics department; semantics is the real thing. * How should we talk about semantics? - A few well-known formalisms for semantics. - We will use programs to explain semantics: the best explanation *is* a program. - Ignore possible philosophical issues with circularity (but be aware of them). (Actually, they are solved: Scheme has a formal explanation that can be taken as a translation from Scheme to logic, which means that things that we write can be translated to logic.) - We will use Racket for many reasons (syntax, functional, practical, simple, formal, statically typed, environment). ------------------------------------------------------------------------ # Intro to Racket [Tuesday, January 7th] * General layout of the parts of Racket: - The Racket language is (mostly) in the Scheme family, or more generally in the Lisp family; - Racket: the core language implementation (language and runtime), written in C; - The actual language(s) that are available in Racket have lots of additional parts that are implemented in Racket itself; - GRacket: a portable Racket GUI extension, written in Racket too; - DrRacket: a GRacket application (also written in Racket); - Our language(s)... * Documentation: the Racket documentation is your friend (But beware that some things are provided in different forms from different places). ------------------------------------------------------------------------ ## Side-note: "Goto Statement Considered Harmful" [Tuesday, January 7th] > A review of "Goto Statement Considered Harmful", by E.W. DIJKSTRA > > This paper tries to convince us that the well-known goto statement > should be eliminated from our programming languages or, at least > (since I don't think that it will ever be eliminated), that > programmers should not use it. It is not clear what should replace > it. The paper doesn't explain to us what would be the use of the `if` > statement without a `goto` to redirect the flow of execution: Should > all our postconditions consist of a single statement, or should we > only use the arithmetic `if`, which doesn't contain the offensive > `goto`? > > And how will one deal with the case in which, having reached the end > of an alternative, the program needs to continue the execution > somewhere else? > > The author is a proponent of the so-called "structured programming" > style, in which, if I get it right, gotos are replaced by indentation. > Structured programming is a nice academic exercise, which works well > for small examples, but I doubt that any real-world program will ever > be written in such a style. More than 10 years of industrial > experience with Fortran have proved conclusively to everybody > concerned that, in the real world, the goto is useful and necessary: > its presence might cause some inconveniences in debugging, but it is a > de facto standard and we must live with it. It will take more than > the academic elucubrations of a purist to remove it from our > languages. > > Publishing this would waste valuable paper: Should it be published, I > am as sure it will go uncited and unnoticed as I am confident that, 30 > years from now, the goto will still be alive and well and used as > widely as it is today. > > Confidential comments to the editor: The author should withdraw the > paper and submit it someplace where it will not be peer reviewed. A > letter to the editor would be a perfect choice: Nobody will notice it > there! ------------------------------------------------------------------------ # Quick Intro to Racket [Tuesday, January 7th] Racket syntax: similar to other Sexpr-based languages. Reminder: the parens can be compared to C/etc function call parens --- they always mean that some function is applied. This is the reason why `(+ (1) (2))` won't work: if you use C syntax that is `+(1(), 2())` but `1` isn't a function so `1()` is an error. An important difference between *syntax* and *semantics*. A good way to think about this is the difference between the *string* `42` stored in a file somewhere (two ASCII values), and the *number* `42` stored in memory (in some representation). You could also continue with the above example: there is nothing wrong with "*murder*" --- it's just a word, but *murder* is something you'll go to jail for. The evaluation function that Racket uses is actually a function that takes a piece of syntax and returns (or executes) its semantics. ------------------------------------------------------------------------ `define` expressions are used for creating new bindings, do not try to use them to change values. For example, you should not try to write something like `(define x (+ x 1))` in an attempt to mimic `x = x+1`. It will not work. ------------------------------------------------------------------------ There are two boolean values built in to Racket: `#t` (true) and `#f` (false). They can be used in `if` statements, for example: (if (< 2 3) 10 20) --> 10 because `(< 2 3)` evaluates to `#t`. As a matter of fact, *any* value except for `#f` is considered to be true, so: (if 0 1 2) --> 1 ; all of these are "truthy" (if "false" 1 2) --> 1 (if "" 1 2) --> 1 (if null 1 2) --> 1 (if #t 1 2) --> 1 ; including the true value (if #f 1 2) --> 2 ; the only false value (if #false 1 2) --> 2 ; another way to write it (if false 1 2) --> 2 ; also false since it's bound to #f ------------------------------------------------------------------------ Note: Racket is a *functional* language --- so *everything* has a value. This means that the expression (if test consequent) has no meaning when `test` evaluates to `#f`. This is unlike Pascal/C where statements *do* something (side effect) like printing or an assignment --- here an `if` statement with no alternate part will just *do nothing* if the test is false... Racket, however, must return some value --- it could decide on simply returning `#f` (or some unspecified value) as the value of (if #f something) as some implementations do, but Racket just declares it a syntax error. (As we will see in the future, Racket has a more convenient `when` with a clearer intention.) ------------------------------------------------------------------------ Well, *almost* everything is a value... There are certain things that are part of Racket's syntax --- for example `if` and `define` are special forms, they do not have a value! More about this shortly. (Bottom line: much more things do have a value, compared with other languages.) ------------------------------------------------------------------------ `cond` is used for a `if` … `else if` … `else if` … `else` … sequence. The problem is that nested `if`s are inconvenient. For example, (define (digit-num n) (if (<= n 9) 1 (if (<= n 99) 2 (if (<= n 999) 3 (if (<= n 9999) 4 "a lot"))))) In C/Java/Whatever, you'd write: function digit_num(n) { if (n <= 9) return 1; else if (n <= 99) return 2; else if (n <= 999) return 3; else if (n <= 9999) return 4; else return "a lot"; } (Side question: why isn't there a `return` statement in Racket?) But trying to force Racket code to look similar: (define (digit-num n) (if (<= n 9) 1 (if (<= n 99) 2 (if (<= n 999) 3 (if (<= n 9999) 4 "a lot"))))) is more than just bad taste --- the indentation rules are there for a reason, the main one is that you can see the structure of your program at a quick glance, and this is no longer true in the above code. (Such code will be penalized!) So, instead of this, we can use Racket's `cond` statement, like this: (define (digit-num n) (cond [(<= n 9) 1] [(<= n 99) 2] [(<= n 999) 3] [(<= n 9999) 4] [else "a lot"])) Note that `else` is a keyword that is used by the `cond` form --- you should always use an `else` clause (for similar reasons as an `if`, to avoid an extra expression evaluation there, and we will need it when we use a typed language). Also note that square brackets are read by DrRacket like round parens, it will only make sure that the paren pairs match. We use this to make code more readable --- specifically, there is a major difference between the above use of `[]` from the conventional use of `()`. Can you see what it is? The general structure of a `cond`: (cond [test-1 expr-1] [test-2 expr-2] ... [test-n expr-n] [else else-expr]) ------------------------------------------------------------------------ Example for using an `if` expression, and a recursive function: (define (fact n) (if (zero? n) 1 (* n (fact (- n 1))))) Use this to show the different tools, especially: * special objects that *cannot* be used * syntax-checker * stepper * submission tool (installing, registering and submitting) An example of converting it to tail recursive form: (define (helper n acc) (if (zero? n) acc (helper (- n 1) (* acc n)))) (define (fact n) (helper n 1)) ------------------------------------------------------------------------ Additional notes about homework submissions: * Begin every function with clear documentation: a type followed by a purpose statement. * Document the function when needed, and according to the guidelines above and in the style guide. * After the function, always have a few test cases --- they should cover your complete code (make sure to include possible corner cases). Later on, we will switch to testing the whole file through it's "public interface", instead of testing each function. ------------------------------------------------------------------------ # Lists & Recursion [Tuesday, January 7th] Lists are a fundamental Racket data type. A list is defined as either: 1. the empty list (`null`, `empty`, or `'()`), 2. a pair (`cons` cell) of anything and a list. As simple as this may seem, it gives us precise *formal* rules to prove that something is a list. * Why is there a "the" in the first rule? Examples: null (cons 1 null) (cons 1 (cons 2 (cons 3 null))) (list 1 2 3) ; a more convenient function to get the above List operations --- predicates: null? ; true only for the empty list pair? ; true for any cons cell list? ; this can be defined using the above We can derive `list?` from the above rules: (define (list? x) (if (null? x) #t (and (pair? x) (list? (rest x))))) or better: (define (list? x) (or (null? x) (and (pair? x) (list? (rest x))))) But why can't we define `list?` more simply as (define (list? x) (or (null? x) (pair? x))) The difference between the above definition and the proper one can be observed in the full Racket language, not in the student languages (where there are no pairs with non-list values in their tails). List operations --- destructors for pairs (`cons` cells): first rest Traditionally called `car`, `cdr`. Also, any `cr` combination for `` that is made of up to four `a`s and/or `d`s --- we will probably not use much more than `cadr`, `caddr` etc. ------------------------------------------------------------------------ Example for recursive function involving lists: (define (list-length list) (if (null? list) 0 (+ 1 (list-length (rest list))))) Use different tools, esp: * syntax-checker * stepper How come we could use `list` as an argument --- use the syntax checker (define (list-length-helper list len) (if (null? list) len (list-length-helper (rest list) (+ len 1)))) (define (list-length list) (list-length-helper list 0)) Main idea: lists are a recursive structure, so functions that operate on lists should be recursive functions that follow the recursive definition of lists. Another example for list function --- summing a list of numbers (define (sum-list l) (if (null? l) 0 (+ (first l) (sum-list (rest l))))) Also show how to implement `rcons`, using this guideline. ------------------------------------------------------------------------ More examples: Define `reverse` --- solve the problem using `rcons`. `rcons` can be generalized into something very useful: `append`. * How would we use `append` instead of `rcons`? * How much time will this take? Does it matter if we use `append` or `rcons`? Redefine `reverse` using tail recursion. * Is the result more complex? (Yes, but not too bad because it collects the elements in reverse.) ------------------------------------------------------------------------ # Some Style [Tuesday, January 7th] When you have some common value that you need to use in several places, it is bad to duplicate it. For example: (define (how-many a b c) (cond [(> (* b b) (* 4 a c)) 2] [(= (* b b) (* 4 a c)) 1] [(< (* b b) (* 4 a c)) 0])) What's bad about it? * It's longer than necessary, which will eventually make your code less readable. * It's slower --- by the time you reach the last case, you have evaluated the two sequences three times. * It's more prone to bugs --- the above code is short enough, but what if it was longer so you don't see the three occurrences on the same page? Will you remember to fix all places when you debug the code months after it was written? In general, the ability to use names is probably the most fundamental concept in computer science --- the fact that makes computer programs what they are. We already have a facility to name values: function arguments. We could split the above function into two like this: (define (how-many-helper b^2 4ac) ; note: valid names! (cond [(> b^2 4ac) 2] [(= b^2 4ac) 1] [else 0])) (define (how-many a b c) (how-many-helper (* b b) (* 4 a c))) But instead of the awkward solution of coming up with a new function just for its names, we have a facility to bind local names --- `let`. In general, the syntax for a `let` special form is (let ([id expr] ...) expr) For example, (let ([x 1] [y 2]) (+ x y)) But note that the bindings are done "in parallel", for example, try this: (let ([x 1] [y 2]) (let ([x y] [y x]) (list x y))) (Note that "in parallel" is quoted here because it's not really parallelism, but just a matter of scopes: the RHSs are all evaluated in the surrounding scope!) Using this for the above problem: (define (how-many a b c) (let ([b^2 (* b b)] [4ac (* 4 a c)]) (cond [(> b^2 4ac) 2] [(= b^2 4ac) 1] [else 0]))) ------------------------------------------------------------------------ Some notes on writing code (also see the style-guide in the handouts section) ***Code quality will be graded to in this course!*** * Use abstractions whenever possible, as said above. This is bad: (define (how-many a b c) (cond [(> (* b b) (* 4 a c)) 2] [(= (* b b) (* 4 a c)) 1] [(< (* b b) (* 4 a c)) 0])) (define (what-kind a b c) (cond [(= a 0) 'degenerate] [(> (* b b) (* 4 a c)) 'two] [(= (* b b) (* 4 a c)) 'one] [(< (* b b) (* 4 a c)) 'none])) * But don't over abstract: `(define one 1)` or `(define two "two")` * Always do test cases (show coverage tool), you might want to comment them, but you should always make sure your code works. * Do not under-document, but also don't over-document. * ***INDENTATION!*** (Let DrRacket decide; get used to its rules) --> This is part of the culture that was mentioned last time, but it's done this way for good reason: decades of programming experience have shown this to be the most readable format. It's also extremely important to keep good indentation since programmers in all Lisps don't count parens --- they look at the structure. * As a general rule, `if` should be either all on one line, or the condition on the first and each consequent on a separate line. Similarly for `define` --- either all on one line or a newline after the object that is being define (either an identifier or a an identifier with arguments). * Another general rule: you should never have white space after an open-paren, or before a close paren (white space includes newlines). Also, before an open paren there should be either another open paren or white space, and the same goes for after a closing paren. * Use the tools that are available to you: for example, use `cond` instead of nested `if`s (definitely do not force the indentation to make a nested `if` look like its C counterpart --- remember to let DrRacket indent for you). Another example --- do not use `(+ 1 (+ 2 3))` instead of `(+ 1 2 3)` (this might be needed in *extremely* rare situations, only when you know your calculus and have extensive knowledge about round-off errors). Another example --- do not use `(cons 1 (cons 2 (cons 3 null)))` instead of `(list 1 2 3)`. Also --- don't write things like: (if (< x 100) #t #f) since it's the same as just (< x 100) A few more of these: (if x #t y) --same-as--> (or x y) (if x y #f) --same-as--> (and x y) (if x #f #t) --same-as--> (not x) (Actually the first two are almost the same, for example, `(and 1 2)` will return `2`, not `#t`.) * Use these as examples for many of these issues: (define (interest x) (* x (cond [(and (> x 0) (<= x 1000)) 0.04] [(and (> x 1000) (<= x 5000)) 0.045] [else 0.05]))) (define (how-many a b c) (cond ((> (* b b) (* (* 4 a) c)) 2) ((< (* b b) (* (* 4 a) c)) 0) (else 1))) (define (what-kind a b c) (if (equal? a 0) 'degenerate (if (equal? (how-many a b c) 0) 'zero (if (equal? (how-many a b c) 1) 'one 'two) ) ) ) (define (interest deposit) (cond [(< deposit 0) "invalid deposit"] [(and (>= deposit 0) (<= deposit 1000)) (* deposit 1.04) ] [(and (> deposit 1000) (<= deposit 5000)) (* deposit 1.045)] [(> deposit 5000) (* deposit 1.05)])) (define (interest deposit) (if (< deposit 1001) (* 0.04 deposit) (if (< deposit 5001) (* 0.045 deposit) (* 0.05 deposit)))) (define (what-kind a b c) (cond ((= 0 a) 'degenerate) (else (cond ((> (* b b)(*(* 4 a) c)) 'two) (else (cond ((= (* b b)(*(* 4 a) c)) 'one) (else 'none))))))); ------------------------------------------------------------------------ The fact that in Racket we can use functions as values is very useful --- for example, `map`, `foldl` & `foldr`, many more. Example: ;; every? : (A -> Boolean) (Listof A) -> Boolean ;; Returns false if any element of lst fails ;; the given pred, true if all pass pred. (define (every? pred lst) (or (null? lst) (and (pred (first lst)) (every? pred (rest lst))))) ------------------------------------------------------------------------ # Tail calls [Tuesday, January 14th] You should generally know what tail calls are, but here's a quick review of the subject. A function call is said to be in tail position if there is no context to "remember" when you're calling it. Very roughly, this means that function calls that are not nested in argument expressions of another call are tail calls. This definition is something that depends on a context, for example, in an expression like (if (huh?) (foo (add1 (* x 3))) (foo (/ x 2))) both calls to `foo` are tail calls, but they're tail calls of *this* expression and therefore apply to *this* context. It might be that this code is inside another call, as in (blah (if (huh?) (foo (add1 (* x 3))) (foo (/ x 2))) something-else) and the `foo` calls are now not in tail position. The main feature of all Scheme implementations including Racket wrt tail calls is that calls that are in tail position of a function are said to be "eliminated". That means that if we're in an `f` function, and we're about to call `g` in tail position and therefore whatever `g` returns would be the result of `f` too, then when Racket does the call to `g` it doesn't bother keeping the `f` context --- it won't remember that it needs to "return" to `f` and will instead return straight to its caller. In other words, when you think about a conventional implementation of function calls as frames on a stack, Racket will get rid of a stack frame when it can. Another way to see this is to use DrRacket's stepper to step through a function call. The stepper is generally an alternative debugger, where instead of visualizing stack frames it assembles an expression that represents these frames. Now, in the case of tail calls, there is no room in such a representation to keep the call --- and the thing is that in Racket that's perfectly fine since these calls are not kept on the call stack. Note that there are several names for this feature: * "Tail recursion". This is a common way to refer to the more limited optimization of *only* tail-recursive functions into loops. In languages that have tail calls as a feature, this is too limited, since they also optimize cases of mutual recursion, or any case of a tail call. * "Tail call optimization". In some languages, or more specifically in some compilers, you'll hear this term. This is fine when tail calls are considered only an "optimization" --- but in Racket's case (as well as Scheme), it's more than just an optimization: it's a *language feature* that you can rely on. For example, a tail-recursive function like `(define (loop) (loop))` *must* run as an infinite loop, not just optimized to one when the compiler feels like it. * "Tail call elimination". This is the so far the most common proper name for the feature: it's not just recursion, and it's not an optimization. ### When should you use tail calls? #################################### Often, people who are aware of tail calls will try to use them *always*. That's not always a good idea. You should generally be aware of the tradeoffs when you consider what style to use. The main thing to remember is that tail-call elimination is a property that helps reducing *space* use (stack space) --- often reducing it from linear space to constant space. This can obviously make things faster, but usually the speedup is just a constant factor since you need to do the same number of iterations anyway, so you just reduce the time spent on space allocation. Here is one such example that we've seen: (define (list-length-1 list) (if (null? list) 0 (+ 1 (list-length-1 (rest list))))) ;; versus (define (list-length-helper list len) (if (null? list) len (list-length-helper (rest list) (+ len 1)))) (define (list-length-2 list) (list-length-helper list 0)) In this case the first (recursive) version version consumes space linear to the length of the list, whereas the second version needs only constant space. But if you consider only the asymptotic runtime, they are both O(length(l)). A second example is a simple implementation of `map`: (define (map-1 f l) (if (null? l) l (cons (f (first l)) (map-1 f (rest l))))) ;; versus (define (map-helper f l acc) (if (null? l) (reverse acc) (map-helper f (rest l) (cons (f (first l)) acc)))) (define (map-2 f l) (map-helper f l '())) In this case, both the asymptotic space and the runtime consumption are the same. In the recursive case we have a constant factor for the stack space, and in the iterative one (the tail-call version) we also have a similar factor for accumulating the reversed list. In this case, it is probably better to keep the first version since the code is simpler. In fact, Racket's stack space management can make the first version run faster than the second --- so optimizing it into the second version is useless. ------------------------------------------------------------------------ # Note on Types [Tuesday, January 14th] Types can become interesting when dealing with higher-order functions. For example, `map` receives a function and a list of some type, and applies the function over this list to accumulate its output, so its type is: ;; map : (A -> B) (Listof A) -> (Listof B) Actually, `map` can use more than a single list, it will apply the function on the first element in all lists, then the second and so on. So the type of `map` with two lists can be described as: ;; map : (A B -> C) (Listof A) (Listof B) -> (Listof C) Here's a hairy example --- what is the type of this function: (define (foo x y) (map map x y)) Begin by what we know --- both `map`s, call them `map1` and `map2`, have the double- and single-list types of `map` respectively, here they are, with different names for types: ;; the first `map', consumes a function and two lists map1 : (A B -> C) (Listof A) (Listof B) -> (Listof C) ;; the second `map', consumes a function and one list map2 : (X -> Y) (Listof X) -> (Listof Y) Now, we know that `map2` is the first argument to `map1`, so the type of `map1`s first argument should be the type of `map2`: (A B -> C) = (X -> Y) (Listof X) -> (Listof Y) From here we can conclude that A = (X -> Y) B = (Listof X) C = (Listof Y) If we use these equations in `map1`'s type, we get: map1 : ((X -> Y) (Listof X) -> (Listof Y)) (Listof (X -> Y)) (Listof (Listof X)) -> (Listof (Listof Y)) Now, `foo`'s two arguments are the 2nd and 3rd arguments of `map1`, and its result is `map1`s result, so we can now write the type of `foo`: ;; foo : (Listof (X -> Y)) ;; (Listof (Listof X)) ;; -> (Listof (Listof Y)) (define (foo x y) (map map x y)) This should help you understand why, for example, this will cause a type error: (foo (list add1 sub1 add1) (list 1 2 3)) and why this is valid: (foo (list add1 sub1 add1) (map list (list 1 2 3))) ------------------------------------------------------------------------ ## Side-note: Names are important [Tuesday, January 14th] An important "discovery" in computer science is that we *don't* need names for every intermediate sub-expression --- for example, in almost any language we can write the equivalent of: s = (-b + sqrt(b^2 - 4*a*c)) / (2*a) instead of x = b * b y = 4 * a y = y * c x = x - y x = sqrt(x) y = -b x = y + x y = 2 * a s = x / y Such languages are put in contrast to assembly languages, and were all put under the generic label of "high level languages". (Here's an interesting idea --- why not do the same for function values?) ------------------------------------------------------------------------ # BNF, Grammars, the AE Language [Tuesday, January 14th] Getting back to the theme of the course: we want to investigate programming languages, and we want to do that *using* a programming language. The first thing when we design a language is to specify the language. For this we use BNF (Backus-Naur Form). For example, here is the definition of a simple arithmetic language: ::= | + | - Explain the different parts. Specifically, this is a mixture of low-level (concrete) syntax definition with parsing. We use this to derive expressions in some language. We start with ``, which should be one of these: * a number `` * an ``, the text "`+`", and another `` * the same but with "`-`" `` is a terminal: when we reach it in the derivation, we're done. `` is a non-terminal: when we reach it, we have to continue with one of the options. It should be clear that the `+` and the `-` are things we expect to find in the input --- because they are not wrapped in `<>`s. We could specify what `` is (turning it into a `` non-terminal): ::= | + | - ::= 0 | 1 | 2 | 3 | 4 | 5 | 6 | 7 | 8 | 9 | But we don't --- why? Because in Racket we have numbers as primitives and we want to use Racket to implement our languages. This makes life a lot easier, and we get free stuff like floats, rationals etc. To use a BNF formally, for example, to prove that `1-2+3` is a valid `` expression, we first label the rules: ::= (1) | + (2) | - (3) and then we can use them as formal justifications for each derivation step: + ; (2) + ; (1) - + ; (3) - + 3 ; (num) - + 3 ; (1) - + 3 ; (1) 1 - + 3 ; (num) 1 - 2 + 3 ; (num) This would be one way of doing this. Alternatively, we can can visualize the derivation using a tree, with the rules used at the nodes. These specifications suffer from being ambiguous: an expression can be derived in multiple ways. Even the little syntax for a number is ambiguous --- a number like `123` can be derived in two ways that result in trees that look different. This ambiguity is not a "real" problem now, but it will become one very soon. We want to get rid of this ambiguity, so that there is a single (= deterministic) way to derive all expressions. There is a standard way to resolve that --- we add another non-terminal to the definition, and make it so that each rule can continue to exactly one of its alternatives. For example, this is what we can do with numbers: ::= | ::= 0 | 1 | 2 | 3 | 4 | 5 | 6 | 7 | 8 | 9 Similar solutions can be applied to the `` BNF --- we either restrict the way derivations can happen or we come up with new non-terminals to force a deterministic derivation trees. As an example of restricting derivations, we look at the current grammar: ::= | + | - and instead of allowing an `` on both sides of the operation, we force one to be a number: ::= | + | - Now there is a single way to derive any expression, and it is always associating operations to the right: an expression like `1+2+3` can only be derived as `1+(2+3)`. To change this to left-association, we would use this: ::= | + | - But what if we want to force precedence? Say that our AE syntax has addition and multiplication: ::= | + | * We can do that same thing as above and add new non-terminals --- say one for "products": ::= | + | ::= | * Now we must parse any AE expression as additions of multiplications (or numbers). First, note that if `` goes to `` and that goes to ``, then there is no need for an `` to go to a ``, so this is the same syntax: ::= + | ::= | * Now, if we want to still be able to multiply additions, we can force them to appear in parentheses: ::= + | ::= | * | ( ) Next, note that `` is still ambiguous about additions, which can be fixed by forcing the left hand side of an addition to be a factor: ::= + | ::= | * | ( ) We still have an ambiguity for multiplications, so we do the same thing and add another non-terminal for "atoms": ::= + | ::= * | ::= | ( ) And you can try to derive several expressions to be convinced that derivation is always deterministic now. But as you can see, this is exactly the cosmetics that we want to avoid --- it will lead us to things that might be interesting, but unrelated to the principles behind programming languages. It will also become much much worse when we have a real language rather such a tiny one. Is there a good solution? --- It is right in our face: do what Racket does --- always use fully parenthesized expressions: ::= | ( + ) | ( - ) To prevent confusing Racket code with code in our language(s), we also change the parentheses to curly ones: ::= | { + } | { - } But in Racket *everything* has a value --- including those `+`s and `-`s, which makes this extremely convenient with future operations that might have either more or less arguments than 2 as well as treating these arithmetic operators as plain functions. In our toy language we will not do this initially (that is, `+` and `-` are second order operators: they cannot be used as values). But since we will get to it later, we'll adopt the Racket solution and use a fully-parenthesized prefix notation: ::= | { + } | { - } (Remember that in a sense, Racket code is written in a form of already-parsed syntax...) ------------------------------------------------------------------------ # Simple Parsing [Tuesday, January 14th] On to an implementation of a "parser": Unrelated to what the syntax actually looks like, we want to parse it as soon as possible --- converting the concrete syntax to an abstract syntax tree. No matter how we write our syntax: - `3+4` (infix), - `3 4 +` (postfix), - `+(3,4)` (prefix with args in parens), - `(+ 3 4)` (parenthesized prefix), we always mean the same abstract thing --- adding the number `3` and the number `4`. The essence of this is basically a tree structure with an addition operation as the root and two leaves holding the two numerals. With the right data definition, we can describe this in Racket as the expression `(Add (Num 3) (Num 4))` where `Add` and `Num` are constructors of a tree type for syntax, or in a C-like language, it could be something like `Add(Num(3),Num(4))`. Similarly, the expression `(3-4)+7` will be described in Racket as the expression: (Add (Sub (Num 3) (Num 4)) (Num 7)) Important note: "expression" was used in two *different* ways in the above --- each way corresponds to a different language, and the result of evaluating the second "expression" is a Racket value that *represents* the first expression. To define the data type and the necessary constructors we will use this: (define-type AE [Num Number] [Add AE AE] [Sub AE AE]) * Note --- Racket follows the tradition of Lisp which makes syntax issues almost negligible --- the language we use is almost as if we are using the parse tree directly. Actually, it is a very simple syntax for parse trees, one that makes parsing extremely easy. [This has an interesting historical reason... Some Lisp history --- *M-expressions* vs. *S-expressions*, and the fact that we write code that is isomorphic to an AST. Later we will see some of the advantages that we get by doing this. See also "*The Evolution of Lisp*", section 3.5.1 (especially the last sentence).] To make things very simple, we will use the above fact through a double-level approach: * we first "parse" our language into an intermediate representation --- a Racket list --- this is mostly done by a modified version of Racket's `read` function that uses curly `{}` braces instead of round `()` parens, * then we write our own `parse` function that will parse the resulting list into an instance of the `AE` type --- an abstract syntax tree (AST). This is achieved by the following simple recursive function: (: parse-sexpr : Sexpr -> AE) ;; parses s-expressions into AEs (define (parse-sexpr sexpr) (cond [(number? sexpr) (Num sexpr)] [(and (list? sexpr) (= 3 (length sexpr))) (let ([make-node (match (first sexpr) ['+ Add] ['- Sub] [else (error 'parse-sexpr "unknown op: ~s" (first sexpr))]) #| the above is the same as: (cond [(equal? '+ (first sexpr)) Add] [(equal? '- (first sexpr)) Sub] [else (error 'parse-sexpr "unknown op: ~s" (first sexpr))]) |#]) (make-node (parse-sexpr (second sexpr)) (parse-sexpr (third sexpr))))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) This function is pretty simple, but as our languages grow, they will become more verbose and more difficult to write. So, instead, we use a new special form: `match`, which is matching a value and binds new identifiers to different parts (try it with "Check Syntax"). Re-writing the above code using `match`: (: parse-sexpr : Sexpr -> AE) ;; parses s-expressions into AEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(list '+ left right) (Add (parse-sexpr left) (parse-sexpr right))] [(list '- left right) (Sub (parse-sexpr left) (parse-sexpr right))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) And finally, to make it more uniform, we will combine this with the function that parses a string into a sexpr so we can use strings to represent our programs: (: parse : String -> AE) ;; parses a string containing an AE expression to an AE (define (parse str) (parse-sexpr (string->sexpr str))) ------------------------------------------------------------------------ # The `match` Form [Tuesday, January 14th] The syntax for `match` is (match value [pattern result-expr] ...) The value is matched against each pattern, possibly binding names in the process, and if a pattern matches it evaluates the result expression. The simplest form of a pattern is simply an identifier --- it always matches and binds that identifier to the value: (match (list 1 2 3) [x x]) ; evaluates to the list Another simple pattern is a quoted symbol, which matches that symbol. For example: (match foo ['x "yes"] [else "no"]) will evaluate to `"yes"` if `foo` is the symbol `x`, and to `"no"` otherwise. Note that `else` is not a keyword here --- it happens to be a pattern that always succeeds, so it behaves like an else clause except that it binds `else` to the unmatched-so-far value. Many patterns look like function application --- but don't confuse them with applications. A `(list x y z)` pattern matches a list of exactly three items and binds the three identifiers; or if the "arguments" are themselves patterns, `match` will descend into the values and match them too. More specifically, this means that patterns can be nested: (match (list 1 2 3) [(list x y z) (+ x y z)]) ; evaluates to 6 (match '((1) (2) 3) [(list (list x) (list y) z) (+ x y z)]) ; also 6 There is also a `cons` pattern that matches a non-empty list and then matches the first part against the head for the list and the second part against the tail of the list. In a `list` pattern, you can use `...` to specify that the previous pattern is repeated zero or more times, and bound names get bound to the list of respective matching. One simple consequent is that the `(list hd tl ...)` pattern is exactly the same as `(cons hd tl)`, but being able to repeat an arbitrary pattern is very useful: > (match '((1 2) (3 4) (5 6) (7 8)) [(list (list x y) ...) (list x y)]) '((1 3 5 7) (2 4 6 8)) A few more useful patterns: id -- matches anything, binds `id' to it _ -- matches anything, but does not bind (number: n) -- matches any number and binds it to `n' (symbol: s) -- same for symbols (string: s) -- strings (sexpr: s) -- S-expressions (needed sometimes for Typed Racket) (and pat1 pat2) -- matches both patterns (or pat1 pat2) -- matches either pattern (careful with bindings) Note that the `foo:` patterns are all specific to our `#lang pl`, they are not part of `#lang racket` or `#lang typed/racket` The patterns are tried one by one *in-order*, and if no pattern matches the value, an error is raised. Note that `...` in a `list` pattern can follow *any* pattern, including all of the above, and including nested list patterns. Here are a few examples --- you can try them out with `#lang pl untyped` at the top of the definitions window. This: (match x [(list (symbol: syms) ...) syms]) matches `x` against a pattern that accepts only a list of symbols, and binds `syms` to those symbols. And here's an example that matches a list of any number of lists, where each of the sub-lists begins with a symbol and then has any number of numbers. Note how the `n` and `s` bindings get values for a list of all symbols and a list of lists of the numbers: > (define (foo x) (match x [(list (list (symbol: s) (number: n) ...) ...) (list 'symbols: s 'numbers: n)])) > (foo (list (list 'x 1 2 3) (list 'y 4 5))) '(symbols: (x y) numbers: ((1 2 3) (4 5))) Here is a quick example for how `or` is used with two literal alternatives, how `and` is used to name a specific piece of data, and how `or` is used with a binding: > (define (foo x) (match x [(list (or 1 2 3)) 'single] [(list (and x (list 1 _)) 2) x] [(or (list 1 x) (list 2 x)) x])) > (foo (list 3)) 'single > (foo (list (list 1 99) 2)) '(1 99) > (foo (list 1 10)) 10 > (foo (list 2 10)) 10 ------------------------------------------------------------------------ # The `define-type` Form [Tuesday, January 14th] The class language that we're using, `#lang pl`, is based on *Typed Racket*: a statically-typed dialect of Racket. It is not exactly the same as Typed Racket --- it is restricted in many ways, and extended in a few ways. (You should therefore try to avoid looking at the Typed Racket documentation and expect things to be the same in `#lang pl`.) The most important extension is `define-type`, which is the construct we will be using to create new user-defined types. In general, such definitions looks like what we just used: (define-type AE [Num Number] [Add AE AE] [Sub AE AE]) This defines a *new type* called `AE`, an `AE?` predicate for this type, and a few *variants* for this type: `Num`, `Add`, and `Sub` in this case. Each of these variant names is a constructor, taking in arguments with the listed types, where these types can include the newly defined type itself in (the very common) case we're defining a recursive type. The return type is always the newly defined type, `AE` here. To summarize, this definition gives us a new `AE` type, and three constructors, as if we wrote the following type declarations: * `(: Num : Number -> AE)` * `(: Add : AE AE -> AE)` * `(: Sub : AE AE -> AE)` The newly defined types are known as *"disjoint unions"*, since values in these types are disjoint --- there is no overlap between the different variants. As we will see, this is what makes this such a useful construct for our needs: the compiler knows about the variants of each newly defined type, which will make it possible for it to complain if we extend a type with more variants but not update all uses of the type. Furthermore, since the return types of these constructors are all the new type itself, there is *no way* for us to write code that expects just *one* of these variants. We will use a second form, `cases`, to handle these values. ------------------------------------------------------------------------ # The `cases` Form [Tuesday, January 14th] A `define-type` declaration defines *only* what was described above: one new type name and a matching predicate, and a few variants as constructor functions. Unlike HtDP, we don't get predicates for each of the variants, and we don't get accessor functions for the fields of the variants. The way that we handle the new kind of values is with `cases`: this is a form that is very similar to `match`, but is specific to instances of the user-defined type. > Many students find it confusing to distinguish `match` and `cases` > since they are so similar. Try to remember that `match` is for > primitive Racket values (we'll mainly use them for S-expression > values), while `cases` is for user-defined values. The distinction > between the two forms is unfortunate, and doesn't serve any purpose. > It is just technically difficult to unify the two. For example, code that handles `AE` values (as defined above) can look as follows: (cases some-ae-value [(Num n) "a number"] [(Add l r) "an addition"] [(Sub l r) "a subtraction"]) As you can see, we need to have patterns for each of the listed variants (and the compiler will throw an error if some are missing), and each of these patterns specifies bindings that will get the field values contained in a given variant object. We can also use nested patterns: (cases some-ae-value [(Num n) "a number"] [(Add (Num m) (Num n)) "a simple addition"] [(Add l r) "an addition"] [(Sub (Num m) (Num n)) "a simple subtraction"] [(Sub l r) "a subtraction"]) but this is a feature that we will not use too often. The final clause in a `cases` form can be an `else` clause, which serves as a fallback in case none of the previous clauses matched the input value. However, using an `else` like this is ***strongly discouraged!*** The problem with using it is that it effectively eliminates the advantage in getting the type-checker to complain when a type definition is extended with new variants. Using these `else` clauses, we can actually mimic all of the functionality that you expect in HtDP-style code, which demonstrates that this is equivalent to HtDP-style definitions. For example: (: Add? : AE -> Boolean) ;; identifies instances of the `Add` variant (define (Add? ae) (cases ae [(Add l r) #t] [else #f])) (: Add-left : AE -> AE) ;; get the left-hand subexpression of an addition (define (Add-left ae) (cases ae [(Add l r) l] [else (error 'Add-left "expecting an Add value, got ~s" ae)])) ... ***Important reminder:*** this is code that ***you should not write!*** Doing so will lead to code that is more fragile than just using `cases`, since you'd be losing the protection the compiler gives you in the form of type errors on occurrences of `cases` that need to be updated when a type is extended with new variants. You would therefore end up writing a bunch of boiler-plate code only to end up with lower-quality code. The core of the problem is in the prevalent use of `else` which gives up that protection. In these examples the `else` clause is justified because even if `AE` is extended with new variants, functions like `Add?` and `Add-left` should not be affected and treat the new variants as they treat all other non-`Add` instances. (And since `else` is inherent to these functions, using them in our code is inherently a bad idea.) We will, however, have a few (*very few!*) places where we'll need to use `else` --- but this will always be done only on some specific functionality rather than a wholesale approach of defining a different interface for user-defined types. ------------------------------------------------------------------------ # Semantics (= Evaluation) [Tuesday, January 14th] > [PLAI §2] Back to BNF --- now, meaning. An important feature of these BNF specifications: we can use the derivations to specify *meaning* (and meaning in our context is "running" a program (or "interpreting", "compiling", but we will use "evaluating")). For example: ::= ; evaluates to the number | + ; evaluates to the sum of evaluating ; and | - ; ... the subtraction of from (... roughly!) To do this a little more formally: a. eval() = ;*** special rule: translate syntax to value b. eval( + ) = eval() + eval() c. eval( - ) = eval() - eval() Note the completely different roles of the two `+`s and `-`s. In fact, it might have been more correct to write: a. eval("") = b. eval(" + ") = eval("") + eval("") c. eval(" - ") = eval("") - eval("") or even using a marker to denote meta-holes in these strings: a. eval("$") = b. eval("$ + $") = eval("$") + eval("$") c. eval("$ - $") = eval("$") - eval("$") but we will avoid pretending that we're doing that kind of string manipulation. (For example, it will require specifying what does it mean to return `` for `$` (involves `string->number`), and the fragments on the right side mean that we need to specify these as substring operations.) Note that there's a similar kind of informality in our BNF specifications, where we assume that `` refers to some terminal or non-terminal. In texts that require more formal specifications (for example, in RFC specifications), each literal part of the BNF is usually double-quoted, so we'd get ::= | "+" | "-" An alternative popular notation for `eval(X)` is `[[X]]`: a. [[]] = b. [[ + ]] = [[]] + [[]] c. [[ - ]] = [[]] - [[]] Is there a problem with this definition? Ambiguity: eval(1 - 2 + 3) = ? Depending on the way the expression is parsed, we can get either a result of `2` or `-4`: eval(1 - 2 + 3) = eval(1 - 2) + eval(3) [b] = eval(1) - eval(2) + eval(3) [c] = 1 - 2 + 3 [a,a,a] = 2 eval(1 - 2 + 3) = eval(1) - eval(2 + 3) [c] = eval(1) - (eval(2) + eval(3)) [a] = 1 - (2 + 3) [a,a,a] = -4 Again, be very aware of confusing subtleties which are extremely important: We need parens around a sub-expression only in one side, why? --- When we write: eval(1 - 2 + 3) = ... = 1 - 2 + 3 we have two expressions, but one stands for an *input syntax*, and one stands for a real mathematical expression. In a case of a computer implementation, the syntax on the left is (as always) an AE syntax, and the real expression on the right is an expression in whatever language we use to implement our AE language. Like we said earlier, ambiguity is not a real problem until the actual parse tree matters. With `eval` it definitely matters, so we must not make it possible to derive any syntax in multiple ways or our evaluation will be non-deterministic. ------------------------------------------------------------------------ Quick exercise: We can define a meaning for ``s and then ``s in a similar way: ::= | eval(0) = 0 eval(1) = 1 eval(2) = 2 ... eval(9) = 9 eval() = eval( ) = 10*eval() + eval() Is this exactly what we want? --- Depends on what we actually want... * First, there's a bug in this code --- having a BNF derivation like ::= | is unambiguous, but makes it hard to parse a number. We get: eval(123) = 10*eval(1) + eval(23) = 10*1 + 10*eval(2) + eval(3) = 10*1 + 10*2 + 3 = 33 Changing the order of the last rule works much better: ::= | and then: eval( ) = 10*eval() + eval() * As a concrete example see how you would make it work with `107`, which demonstrates why compositionality is important. * Example for free stuff that looks trivial: if we were to define the meaning of numbers this way, would it always work? Think an average language that does not give you bignums, making the above rules fail when the numbers are too big. In Racket, we happen to be using an integer representation for the syntax of integers, and both are unlimited. But what if we wanted to write a Racket compiler in C or a C compiler in Racket? What about a C compiler in C, where the compiler runs on a 64 bit machine, and the result needs to run on a 32 bit machine? ------------------------------------------------------------------------ ## Side-note: Compositionality [Tuesday, January 14th] The example of ::= | being a language that is easier to write an evaluator for leads us to an important concept --- compositionality. This definition is easier to write an evaluator for, since the resulting language is compositional: the meaning of an expression --- for example `123` --- is composed out of the meaning of its two parts, which in this BNF are `12` and `3`. Specifically, the evaluation of ` ` is `10 *` the evaluation of the first, plus the evaluation of the second. In the ` ` case this is more difficult --- the meaning of such a number depends not only on the *meaning* of the two parts, but also on the `` *syntax*: eval( ) = eval() * 10^length() + eval() This this case this can be tolerable, since the meaning of the expression is still made out of its parts --- but imperative programming (when you use side effects) is much more problematic since it is not compositional (at least not in the obvious sense). This is compared to functional programming, where the meaning of an expression is a combination of the meanings of its subexpressions. For example, every sub-expression in a functional program has some known meaning, and these all make up the meaning of the expression that contains them --- but in an imperative program we can have a part of the code be `x++` --- and that doesn't have a meaning by itself, at least not one that contributes to the meaning of the whole program in a direct way. (Actually, we can have a well-defined meaning for such an expression: the meaning is going from a world where `x` is a container of some value N, to a world where the same container has a different value N+1. You can probably see now how this can make things more complicated. On an intuitive level --- if we look at a random part of a functional program we can tell its meaning, so building up the meaning of the whole code is easy, but in an imperative program, the meaning of a random part is pretty much useless.) ------------------------------------------------------------------------ # Implementing an Evaluator [Tuesday, January 14th] Now continue to implement the semantics of our syntax --- we express that through an `eval` function that evaluates an expression. We use a basic programming principle --- splitting the code into two layers, one for parsing the input, and one for doing the evaluation. Doing this avoids the mess we'd get into otherwise, for example: (define (eval sexpr) (match sexpr [(number: n) n] [(list '+ left right) (+ (eval left) (eval right))] [(list '- left right) (- (eval left) (eval right))] [else (error 'eval "bad syntax in ~s" sexpr)])) This is messy because it combines two very different things --- syntax and semantics --- into a single lump of code. For this particular kind of evaluator it looks simple enough, but this is only because it's simple enough that all we do is replace constructors by arithmetic operations. Later on things will get more complex, and bundling the evaluator with the parser will be more problematic. (Note: the fact that we can replace constructors with the run-time operators mean that we have a very simple, calculator-like language, and that we can, in fact, "compile" all programs down to a number.) If we split the code, we can easily include decisions like making {+ 1 {- 3 "a"}} syntactically invalid. (Which is not, BTW, what Racket does...) (Also, this is like the distinction between XML syntax and well-formed XML syntax.) An additional advantage is that by using two separate components, it is simple to replace each one, making it possible to change the input syntax, and the semantics independently --- we only need to keep the same interface data (the AST) and things will work fine. Our `parse` function converts an input syntax to an abstract syntax tree (AST). It is abstract exactly because it is independent of any actual concrete syntax that you type in, print out etc. ------------------------------------------------------------------------ # Implementing The AE Language [Tuesday, January 14th] Back to our `eval` --- this will be its (obvious) type: (: eval : AE -> Number) ;; consumes an AE and computes ;; the corresponding number which leads to some obvious test cases: (equal? 3 (eval (parse "3"))) (equal? 7 (eval (parse "{+ 3 4}"))) (equal? 6 (eval (parse "{+ {- 3 4} 7}"))) which from now on we will write using the new `test` form that the `#lang pl` language provides: (test (eval (parse "3")) => 3) (test (eval (parse "{+ 3 4}")) => 7) (test (eval (parse "{+ {- 3 4} 7}")) => 6) Note that we're testing *only* at the interface level --- only running whole functions. For example, you could think about a test like: (test (parse "{+ {- 3 4} 7}") => (Add (Sub (Num 3) (Num 4)) (Num 7))) but the details of parsing and of the constructor names are things that nobody outside of our evaluator cares about --- so we're not testing them. In fact, we shouldn't even mention `parse` in these tests, since it is not part of the public interface of our users; they only care about using it as a compiler-like black box. (This is sometimes called "integration tests".) We'll address this shortly. Like everything else, the structure of the recursive `eval` code follows the recursive structure of its input. In HtDP terms, our template is: (: eval : AE -> Number) (define (eval expr) (cases expr [(Num n) ... n ...] [(Add l r) ... (eval l) ... (eval r) ...] [(Sub l r) ... (eval l) ... (eval r) ...])) In this case, filling in the gaps is very simple (: eval : AE -> Number) (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))])) We now further combine `eval` and `parse` into a single `run` function that evaluates an AE string. (: run : String -> Number) ;; evaluate an AE program contained in a string (define (run str) (eval (parse str))) This function becomes the single public entry point into our code, and the only thing that should be used in tests that verify our interface: (test (run "3") => 3) (test (run "{+ 3 4}") => 7) (test (run "{+ {- 3 4} 7}") => 6) The resulting *full* code is: ;;; ---<<>>----------------------------------------------------- #lang pl #| BNF for the AE language: ::= | { + } | { - } | { * } | { / } |# ;; AE abstract syntax trees (define-type AE [Num Number] [Add AE AE] [Sub AE AE] [Mul AE AE] [Div AE AE]) (: parse-sexpr : Sexpr -> AE) ;; parses s-expressions into AEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> AE) ;; parses a string containing an AE expression to an AE AST (define (parse str) (parse-sexpr (string->sexpr str))) (: eval : AE -> Number) ;; consumes an AE and computes the corresponding number (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))] [(Div l r) (/ (eval l) (eval r))])) (: run : String -> Number) ;; evaluate an AE program contained in a string (define (run str) (eval (parse str))) ;; tests (test (run "3") => 3) (test (run "{+ 3 4}") => 7) (test (run "{+ {- 3 4} 7}") => 6) (Note that the tests are done with a `test` form, which we mentioned above.) For anyone who thinks that Racket is a bad choice, this is a good point to think how much code would be needed in some other language to do the same as above. ------------------------------------------------------------------------ # Intro to Typed Racket [Tuesday, January 14th] The plan: * Why Types? * Why Typed Racket? * What's Different about Typed Racket? * Some Examples of Typed Racket for Course Programs ### Types ############################################################## - Who has used a (statically) typed language? - Who has used a typed language that's not Java? Typed Racket will be both similar to and very different from anything you've seen before. ### Why types? ######################################################### - Types help structure programs. - Types provide enforced and mandatory documentation. - Types help catch errors. --> They *will* help you. A *lot*. ### Structuring programs ############################################### - Data definitions ;; An AE is one of: ; \ ;; (make-Num Number) ; > HtDP ;; (make-Add AE AE) ; / (define-type AE ; \ [Num number?] ; > Predicates =~= contracts (PLAI) [Add AE? AE?]) ; / (has names of defined types too) (define-type AE ; \ [Num Number] ; > Typed Racket (our PL) [Add AE AE]) ; / - Data-first The structure of your program is derived from the structure of your data. You have seen this in Fundamentals with the design recipe and with templates. In this class, we will see it extensively with type definitions and the (cases ...) form. Types make this pervasive --- we have to think about our data before our code. - A language for describing data Instead of having an informal language for describing types in contract lines, and a more formal description of predicates in a `define-type` form, we will have a single, unified language for both of these. Having such a language means that we get to be more precise and more expressive (since the typed language covers cases that you would otherwise dismiss with some hand waving, like "a function"). ### Why Typed Racket? ################################################## Racket is the language we all know, and it has the benefits that we discussed earlier. Mainly, it is an excellent language for experimenting with programming languages. - Typed Racket allows us to take our Racket programs and typecheck them, so we get the benefits of a statically typed language. - Types are an important programming language feature; Typed Racket will help us understand them. [Also: the development of Typed Racket is happening here in Northeastern, and will benefit from your feedback.] ### How is Typed Racket different from Racket ########################## - Typed Racket will reject your program if there are type errors! This means that it does that at compile-time, *before* any code gets to run. - Typed Racket files start like this: #lang typed/racket ;; Program goes here. but we will use a variant of the Typed Racket language, which has a few additional constructs: #lang pl ;; Program goes here. - Typed Racket requires you to write the contracts on your functions. Racket: ;; f : Number -> Number (define (f x) (* x (+ x 1))) Typed Racket: #lang pl (: f : Number -> Number) (define (f x) (* x (+ x 1))) [In the "real" Typed Racket the preferred style is with prefix arrows: #lang typed/racket (: f (-> Number Number)) (define (f x) : Number (* x (+ x 1))) and you can also have the type annotations appear inside the definition: #lang typed/racket (define (f [x : Number]) : Number (* x (+ x 1))) but we will not use these form.] - As we've seen, Typed Racket uses types, not predicates, in `define-type`. (define-type AE [Num Number] [Add AE AE]) versus (define-type AE [Num number?] [Add AE? AE?]) - There are other differences, but these will suffice for now. ### Examples ########################################################### (: digit-num : Number -> (U Number String)) (define (digit-num n) (cond [(<= n 9) 1] [(<= n 99) 2] [(<= n 999) 3] [(<= n 9999) 4] [else "a lot"])) (: fact : Number -> Number) (define (fact n) (if (zero? n) 1 (* n (fact (- n 1))))) (: helper : Number Number -> Number) (define (helper n acc) (if (zero? n) acc (helper (- n 1) (* acc n)))) (: fact : Number -> Number) (define (fact n) (helper n 1)) (: fact : Number -> Number) (define (fact n) (: helper : Number Number -> Number) (define (helper n acc) (if (zero? n) acc (helper (- n 1) (* acc n)))) (helper n 1)) (: every? : (All (A) (A -> Boolean) (Listof A) -> Boolean)) ;; Returns false if any element of lst fails the given pred, ;; true if all pass pred. (define (every? pred lst) (or (null? lst) (and (pred (first lst)) (every? pred (rest lst))))) (define-type AE [Num Number] [Add AE AE] [Sub AE AE]) ;; the only difference in the following definition is ;; using (: : ) instead of ";; : " (: parse-sexpr : Sexpr -> AE) ;; parses s-expressions into AEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(list '+ left right) (Add (parse-sexpr left) (parse-sexpr right))] [(list '- left right) (Sub (parse-sexpr left) (parse-sexpr right))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) ### More interesting examples ########################################## * Typed Racket is designed to be a language that is friendly to the kind of programs that people write in Racket. For example, it has unions: (: foo : (U String Number) -> Number) (define (foo x) (if (string? x) (string-length x) ;; at this point it knows that `x' is not a ;; string, therefore it must be a number (+ 1 x))) This is not common in statically typed languages, which are usually limited to only *disjoint* unions. For example, in OCaml you'd write this definition: type string_or_number = Str of string | Int of int ;; let foo x = match x with Str s -> String.length s | Int i -> i+1 ;; And use it with an explicit constructor: foo (Str "bar") ;; foo (Int 3) ;; * Note that in the Typed Racket case, the language keeps track of information that is gathered via predicates --- which is why it knows that one `x` is a String, and the other is a Number. * Typed Racket has a concept of subtypes --- which is also something that most statically typed languages lack. In fact, the fact that it has (arbitrary) unions means that it must have subtypes too, since a type is always a subtype of a union that contains this type. * Another result of this feature is that there is an `Any` type that is the union of all other types. Note that you can always use this type since everything is in it --- but it gives you the *least* information about a value. In other words, Typed Racket gives you a choice: *you* decide which type to use, one that is very restricted but has a lot of information about its values to a type that is very permissive but has almost no useful information. This is in contrast to other type system (HM systems) where there is always exactly one correct type. To demonstrate, consider the identity function: (define (id x) x) You could use a type of `(: id : Integer -> Integer)` which is very restricted, but you know that the function always returns an integer value. Or you can make it very permissive with a `(: id : Any -> Any)`, but then you know nothing about the result --- in fact, `(+ 1 (id 2))` will throw a type error. It *does* return `2`, as expected, but the type checker doesn't know the type of that `2`. If you wanted to use this type, you'd need to check that the result is a number, eg: (let ([x (id 123)]) (if (number? x) (+ x 10) 999)) This means that for this particular function there is no good *specific* type that we can choose --- but there are *polymorphic* types. These types allow propagating their input type(s) to their output type. In this case, it's a simple "my output type is the same as my input type": (: id : (All (A) A -> A)) This makes the output preserve the same level of information that you had on its input. * Another interesting thing to look at is the type of `error`: it's a function that returns a type of `Nothing` --- a type that is the same as an *empty* union: `(U)`. It's a type that has no values in it --- it fits `error` because it *is* a function that doesn't return any value, in fact, it doesn't return at all. In addition, it means that an `error` expression can be used anywhere you want because it is a subtype of anything at all. * An `else` clause in a `cond` expression is almost always needed, for example: (: digit-num : Number -> (U Number String)) (define (digit-num n) (cond [(<= n 9) 1] [(<= n 99) 2] [(<= n 999) 3] [(<= n 9999) 4] [(> n 9999) "a lot"])) (and if you think that the type checker should know what this is doing, then how about (> (* n 10) (/ (* (- 10000 1) 20) 2)) or (>= n 10000) for the last test?) * In some rare cases you will run into one limitation of Typed Racket: it is difficult (that is: a generic solution is not known at the moment) to do the right inference when polymorphic functions are passed around to higher-order functions. For example: (: call : (All (A B) (A -> B) A -> B)) (define (call f x) (f x)) (call rest (list 4)) In such cases, we can use `inst` to *instantiate* a function with a polymorphic type to a given type --- in this case, we can use it to make it treat `rest` as a function that is specific for numeric lists: (call (inst rest Number) (list 4)) In other rare cases, Typed Racket will infer a type that is not suitable for us --- there is another form, `ann`, that allows us to specify a certain type. Using this in the `call` example is more verbose: (call (ann rest : ((Listof Number) -> (Listof Number))) (list 4)) However, these are going to be rare and will be mentioned explicitly whenever they're needed. ------------------------------------------------------------------------ # Bindings & Substitution [Tuesday, January 14th] We now get to an important concept: substitution. Even in our simple language, we encounter repeated expressions. For example, if we want to compute the square of some expression: {* {+ 4 2} {+ 4 2}} Why would we want to get rid of the repeated sub-expression? * It introduces a redundant computation. In this example, we want to avoid computing the same sub-expression a second time. * It makes the computation more complicated than it could be without the repetition. Compare the above with: with x = {+ 4 2}, {* x x} * This is related to a basic fact in programming that we have already discussed: duplicating information is always a bad thing. Among other bad consequences, it can even lead to bugs that could not happen if we wouldn't duplicate code. A toy example is "fixing" one of the numbers in one expression and forgetting to fix the corresponding one: {* {+ 4 2} {+ 4 1}} Real world examples involve much more code, which make such bugs very difficult to find, but they still follow the same principle. * This gives us more expressive power --- we don't just say that we want to multiply two expressions that both happen to be `{+ 4 2}`, we say that we multiply the `{+ 4 2}` expression by *itself*. It allows us to express identity of two values as well as using two values that happen to be the same. So, the normal way to avoid redundancy is to introduce an identifier. Even when we speak, we might say: "let x be 4 plus 2, multiply x by x". (These are often called "variables", but we will try to avoid this name: what if the identifier does not change (vary)?) To get this, we introduce a new form into our language: {with {x {+ 4 2}} {* x x}} We expect to be able to reduce this to: {* 6 6} by substituting 6 for `x` in the body sub-expression of `with`. A little more complicated example: {with {x {+ 4 2}} {with {y {* x x}} {+ y y}}} [add] = {with {x 6} {with {y {* x x}} {+ y y}}} [subst]= {with {y {* 6 6}} {+ y y}} [mul] = {with {y 36} {+ y y}} [subst]= {+ 36 36} [add] = 72 ------------------------------------------------------------------------ # WAE: Adding Bindings to AE [Tuesday, January 21st] > [PLAI §3] To add this to our language, we start with the BNF. We now call our language "WAE" (With+AE): ::= | { + } | { - } | { * } | { / } | { with { } } | Note that we had to introduce *two* new rules: one for introducing an identifier, and one for using it. This is common in many language specifications, for example `define-type` introduces a new type, and it comes with `cases` that allows us to destruct its instances. For `` we need to use some form of identifiers, the natural choice in Racket is to use symbols. We can therefore write the corresponding type definition: (define-type WAE [Num Number] [Add WAE WAE] [Sub WAE WAE] [Mul WAE WAE] [Div WAE WAE] [Id Symbol] [With Symbol WAE WAE]) The parser is easily extended to produce these syntax objects: (: parse-sexpr : Sexpr -> WAE) ;; parses s-expressions into WAEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) But note that this parser is inconvenient --- if any of these expressions: {* 1 2 3} {foo 5 6} {with x 5 {* x 8}} {with {5 x} {* x 8}} would result in a "bad syntax" error, which is not very helpful. To make things better, we can add another case for `with` expressions that are malformed, and give a more specific message in that case: (: parse-sexpr : Sexpr -> WAE) ;; parses s-expressions into WAEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [(cons 'with more) (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) and finally, to group all of the parsing code that deals with `with` expressions (both valid and invalid ones), we can use a single case for both of them: (: parse-sexpr : Sexpr -> WAE) ;; parses s-expressions into WAEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) ;; go in here for all sexpr that begin with a 'with (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) And now we're done with the syntactic part of the `with` extension. > Quick note --- why would we indent `With` like a normal function in > code like this > > (With 'x > (Num 2) > (Add (Id 'x) (Num 4))) > > instead of an indentation that looks like a `let` > > (With 'x (Num 2) > (Add (Id 'x) (Num 4))) > > ? > > The reason for this is that the second indentation looks like a > binding construct (eg, the indentation used in a `let` expression), > but `With` is *not* a binding form --- it's a *plain function* because > it's at the Racket level. You should therefore keep in mind the huge > difference between that `With` and the `with` that appears in WAE > programs: > > {with {x 2} > {+ x 4}} > > Another way to look at it: imagine that we intend for the language to > be used by Spanish/Chinese/German/French speakers. In this case we > would translate "`with`": > > {con {x 2} {+ x 4}} > {he {x 2} {+ x 4}} > {mit {x 2} {+ x 4}} > {avec {x 2} {+ x 4}} > {c {x 2} {+ x 4}} > > but we will *not* do the same for `With` if we (the language > implementors) are English speakers. ------------------------------------------------------------------------ # Evaluation of `with` [Tuesday, January 21st] Now, to make this work, we will need to do some substitutions. We basically want to say that to evaluate: {with {id WAE1} WAE2} we need to evaluate `WAE2` with id substituted by `WAE1`. Formally: eval( {with {id WAE1} WAE2} ) = eval( subst(WAE2,id,WAE1) ) There is a more common syntax for substitution (quick: what do I mean by this use of "syntax"?): eval( {with {id WAE1} WAE2} ) = eval( WAE2[WAE1/id] ) > Side-note: this syntax originates with logicians who used `[x/v]e`, > and later there was a convention that mimicked the more natural order > of arguments to a function with `e[x->v]`, and eventually both of > these got combined into `e[v/x]` which is a little confusing in that > the left-to-right order of the arguments is not the same as for the > `subst` function. Now all we need is an exact definition of substitution. > Note that substitution is not the same as evaluation, it's only a part > of the evaluation process. In the previous examples, when we > evaluated the expression we did substitutions as well as the usual > arithmetic operations that were already part of the AE evaluator. In > this last definition there is still a missing evaluation step, see if > you can find it. So let us try to define substitution now: > Substitution (take 1): `e[v/i]` \ > To substitute an identifier `i` in an expression `e` with an > expression `v`, replace all identifiers in `e` that have the same > name `i` by the expression `v`. This seems to work with simple expressions, for example: {with {x 5} {+ x x}} --> {+ 5 5} {with {x 5} {+ 10 4}} --> {+ 10 4} however, we crash with an invalid syntax if we try: {with {x 5} {+ x {with {x 3} 10}}} --> {+ 5 {with {5 3} 10}} ??? --- we got to an invalid expression. To fix this, we need to distinguish normal occurrences of identifiers, and ones that are used as new bindings. We need a few new terms for this: 1. Binding Instance: a binding instance of an identifier is one that is used to name it in a new binding. In our `` syntax, binding instances are only the `` position of the `with` form. 2. Scope: the scope of a binding instance is the region of program text in which instances of the identifier refer to the value bound in the binding instance. (Note that this definition actually relies on a definition of substitution, because that is what is used to specify how identifiers refer to values.) 3. Bound Instance (or Bound Occurrence): an instance of an identifier is bound if it is contained within the scope of a binding instance of its name. 4. Free Instance (or Free Occurrence): An identifier that is not contained in any binding instance of its name is said to be free. Using this we can say that the problem with the previous definition of substitution is that it failed to distinguish between bound instances (which should be substituted) and binding instances (which should not). So we try to fix this: > Substitution (take 2): `e[v/i]` \ > To substitute an identifier `i` in an expression `e` with an > expression `v`, replace all instances of `i` that are not > themselves binding instances with the expression `v`. First of all, check the previous examples: {with {x 5} {+ x x}} --> {+ 5 5} {with {x 5} {+ 10 4}} --> {+ 10 4} still work, and {with {x 5} {+ x {with {x 3} 10}}} --> {+ 5 {with {x 3} 10}} --> {+ 5 10} also works. However, if we try this: {with {x 5} {+ x {with {x 3} x}}} we get: --> {+ 5 {with {x 3} 5}} --> {+ 5 5} --> 10 but we want that to be `8`: the inner `x` should be bound by the closest `with` that binds it. The problem is that the new definition of substitution that we have respects binding instances, but it fails to deal with their scope. In the above example, we want the inner `with` to *shadow* the outer `with`'s binding for `x`. > Substitution (take 3): `e[v/i]` \ > To substitute an identifier `i` in an expression `e` with an > expression `v`, replace all instances of `i` that are not > themselves binding instances, and that are not in any nested scope, > with the expression `v`. This avoids bad substitution above, but it is now doing things too carefully: {with {x 5} {+ x {with {y 3} x}}} becomes --> {+ 5 {with {y 3} x}} --> {+ 5 x} which is an error because `x` is unbound (and there is reasonable no rule that we can specify to evaluate it). The problem is that our substitution halts at every new scope, in this case, it stopped at the new `y` scope, but it shouldn't have because it uses a different name. In fact, that last definition of substitution cannot handle any nested scope. Revise again: > Substitution (take 4): `e[v/i]` \ > To substitute an identifier `i` in an expression `e` with an > expression `v`, replace all instances of `i` that are not > themselves binding instances, and that are not in any nested scope > of `i`, with the expression `v`. which, finally, is a good definition. This is just a little too mechanical. Notice that we actually refer to all instances of `i` that are not in a scope of a binding instance of `i`, which simply means all *free occurrences* of `i` --- free in `e` (why? --- remember the definition of "free"?): > Substitution (take 4b): `e[v/i]` \ > To substitute an identifier `i` in an expression `e` with an > expression `v`, replace all instances of `i` that are free in `e` > with the expression `v`. Based on this we can finally write the code for it: (: subst : WAE Symbol WAE -> WAE) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) ; returns expr[to/from] (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (if (eq? bound-id from) expr ;*** don't go in! (With bound-id named-expr (subst bound-body from to)))])) ... and this is just the same as writing a formal "paper version" of the substitution rule. We still have bugs: but we'll need some more work to get to them. ------------------------------------------------------------------------ Before we find the bugs, we need to see when and how substitution is used in the evaluation process. To modify our evaluator, we will need rules to deal with the new syntax pieces --- `with` expressions and identifiers. When we see an expression that looks like: {with {x E1} E2} we continue by *evaluating* `E1` to get a value `V1`, we then substitute the identifier `x` with the expression `V1` in `E2`, and continue by evaluating this new expression. In other words, we have the following evaluation rule: eval( {with {x E1} E2} ) = eval( E2[eval(E1)/x] ) So we know what to do with `with` expressions. How about identifiers? The main feature of `subst`, as said in the purpose statement, is that it leaves no free instances of the substituted variable around. This means that if the initial expression is valid (did not contain any free variables), then when we go from {with {x E1} E2} to E2[E1/x] the result is an expression that has *no* free instances of `x`. So we don't need to handle identifiers in the evaluator --- substitutions make them all go away. We can now extend the formal definition of AE to that of WAE: eval(...) = ... same as the AE rules ... eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) eval(id) = error! If you're paying close attention, you might catch a potential problem in this definition: we're substituting `eval(E1)` for `x` in `E2` --- an operation that requires a WAE expression, but `eval(E1)` is a number. (Look at the type of the `eval` definition we had for AE, then look at the above definition of `subst`.) This seems like being overly pedantic, but we it will require some resolution when we get to the code. The above rules are easily coded as follows: (: eval : WAE -> Number) ;; evaluates WAE expressions by reducing them to numbers (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))] [(Div l r) (/ (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (Num (eval named-expr))))] ;*** [(Id name) (error 'eval "free identifier: ~s" name)])) Note the `Num` expression in the marked line: evaluating the named expression gives us back a number --- we need to convert this number into a syntax to be able to use it with `subst`. The solution is to use `Num` to convert the resulting number into a numeral (the syntax of a number). It's not an elegant solution, but it will do for now. Finally, here are a few test cases. We use a new `test` special form which is part of the course plugin. The way to use `test` is with two expressions and an `=>` arrow --- DrRacket evaluates both, and nothing will happen if the results are equal. If the results are different, you will get a warning line, but evaluation will continue so you can try additional tests. You can also use an `=error>` arrow to test an error message --- use it with some text from the expected error, `?` stands for any single character, and `*` is a sequence of zero or more characters. (When you use `test` in your homework, the handin server will abort when tests fail.) We expect these tests to succeed (make sure that you understand *why* they should succeed). ;; tests (test (run "5") => 5) (test (run "{+ 5 5}") => 10) (test (run "{with {x {+ 5 5}} {+ x x}}") => 20) (test (run "{with {x 5} {+ x x}}") => 10) (test (run "{with {x {+ 5 5}} {with {y {- x 3}} {+ y y}}}") => 14) (test (run "{with {x 5} {with {y {- x 3}} {+ y y}}}") => 4) (test (run "{with {x 5} {+ x {with {x 3} 10}}}") => 15) (test (run "{with {x 5} {+ x {with {x 3} x}}}") => 8) (test (run "{with {x 5} {+ x {with {y 3} x}}}") => 10) (test (run "{with {x 5} {with {y x} y}}") => 5) (test (run "{with {x 5} {with {x x} x}}") => 5) (test (run "{with {x 1} y}") =error> "free identifier") Putting this all together, we get the following code; trying to run this code will raise an unexpected error... #lang pl #| BNF for the WAE language: ::= | { + } | { - } | { * } | { / } | { with { } } | |# ;; WAE abstract syntax trees (define-type WAE [Num Number] [Add WAE WAE] [Sub WAE WAE] [Mul WAE WAE] [Div WAE WAE] [Id Symbol] [With Symbol WAE WAE]) (: parse-sexpr : Sexpr -> WAE) ;; parses s-expressions into WAEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> WAE) ;; parses a string containing a WAE expression to a WAE AST (define (parse str) (parse-sexpr (string->sexpr str))) (: subst : WAE Symbol WAE -> WAE) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (if (eq? bound-id from) expr (With bound-id named-expr (subst bound-body from to)))])) (: eval : WAE -> Number) ;; evaluates WAE expressions by reducing them to numbers (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))] [(Div l r) (/ (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (Num (eval named-expr))))] [(Id name) (error 'eval "free identifier: ~s" name)])) (: run : String -> Number) ;; evaluate a WAE program contained in a string (define (run str) (eval (parse str))) ;; tests (test (run "5") => 5) (test (run "{+ 5 5}") => 10) (test (run "{with {x {+ 5 5}} {+ x x}}") => 20) (test (run "{with {x 5} {+ x x}}") => 10) (test (run "{with {x {+ 5 5}} {with {y {- x 3}} {+ y y}}}") => 14) (test (run "{with {x 5} {with {y {- x 3}} {+ y y}}}") => 4) (test (run "{with {x 5} {+ x {with {x 3} 10}}}") => 15) (test (run "{with {x 5} {+ x {with {x 3} x}}}") => 8) (test (run "{with {x 5} {+ x {with {y 3} x}}}") => 10) (test (run "{with {x 5} {with {y x} y}}") => 5) (test (run "{with {x 5} {with {x x} x}}") => 5) (test (run "{with {x 1} y}") =error> "free identifier") ------------------------------------------------------------------------ ------------------------------------------------------------------------ # Evaluation of `with` (contd.) [Tuesday, January 21st] Oops, this program still has problems that were caught by the tests --- we encounter unexpected free identifier errors. What's the problem now? In expressions like: {with {x 5} {with {y x} y}} we forgot to substitute `x` in the expression that `y` is bound to. We need to the recursive substitute in both the with's body expression as well as its named expression: (: subst : WAE Symbol WAE -> WAE) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (if (eq? bound-id from) expr (With bound-id (subst named-expr from to) ;*** new (subst bound-body from to)))])) And *still* we have a problem... Now it's {with {x 5} {with {x x} x}} that halts with an error, but we want it to evaluate to `5`! Carefully trying out our substitution code reveals the problem: when we substitute `5` for the outer `x`, we don't go inside the inner `with` because it has the same name --- but we *do* need to go into its named expression. We need to substitute in the named expression even if the identifier is the *same* one we're substituting: (: subst : WAE Symbol WAE -> WAE) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (With bound-id (subst named-expr from to) (if (eq? bound-id from) bound-body (subst bound-body from to)))])) The complete (and, finally, correct) version of the code is now: ;;; ---<<>>---------------------------------------------------- #lang pl #| BNF for the WAE language: ::= | { + } | { - } | { * } | { / } | { with { } } | |# ;; WAE abstract syntax trees (define-type WAE [Num Number] [Add WAE WAE] [Sub WAE WAE] [Mul WAE WAE] [Div WAE WAE] [Id Symbol] [With Symbol WAE WAE]) (: parse-sexpr : Sexpr -> WAE) ;; parses s-expressions into WAEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> WAE) ;; parses a string containing a WAE expression to a WAE AST (define (parse str) (parse-sexpr (string->sexpr str))) #| Formal specs for `subst': (`N' is a , `E1', `E2' are s, `x' is some , `y' is a *different* ) N[v/x] = N {+ E1 E2}[v/x] = {+ E1[v/x] E2[v/x]} {- E1 E2}[v/x] = {- E1[v/x] E2[v/x]} {* E1 E2}[v/x] = {* E1[v/x] E2[v/x]} {/ E1 E2}[v/x] = {/ E1[v/x] E2[v/x]} y[v/x] = y x[v/x] = v {with {y E1} E2}[v/x] = {with {y E1[v/x]} E2[v/x]} {with {x E1} E2}[v/x] = {with {x E1[v/x]} E2} |# (: subst : WAE Symbol WAE -> WAE) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (With bound-id (subst named-expr from to) (if (eq? bound-id from) bound-body (subst bound-body from to)))])) #| Formal specs for `eval': eval(N) = N eval({+ E1 E2}) = eval(E1) + eval(E2) eval({- E1 E2}) = eval(E1) - eval(E2) eval({* E1 E2}) = eval(E1) * eval(E2) eval({/ E1 E2}) = eval(E1) / eval(E2) eval(id) = error! eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) |# (: eval : WAE -> Number) ;; evaluates WAE expressions by reducing them to numbers (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))] [(Div l r) (/ (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (Num (eval named-expr))))] [(Id name) (error 'eval "free identifier: ~s" name)])) (: run : String -> Number) ;; evaluate a WAE program contained in a string (define (run str) (eval (parse str))) ;; tests (test (run "5") => 5) (test (run "{+ 5 5}") => 10) (test (run "{with {x 5} {+ x x}}") => 10) (test (run "{with {x {+ 5 5}} {+ x x}}") => 20) (test (run "{with {x 5} {with {y {- x 3}} {+ y y}}}") => 4) (test (run "{with {x {+ 5 5}} {with {y {- x 3}} {+ y y}}}") => 14) (test (run "{with {x 5} {+ x {with {x 3} 10}}}") => 15) (test (run "{with {x 5} {+ x {with {x 3} x}}}") => 8) (test (run "{with {x 5} {+ x {with {y 3} x}}}") => 10) (test (run "{with {x 5} {with {y x} y}}") => 5) (test (run "{with {x 5} {with {x x} x}}") => 5) (test (run "{with {x 1} y}") =error> "free identifier") ------------------------------------------------------------------------ Reminder: * We started doing substitution, with a `let`-like form: `with`. * Reasons for using bindings: - Avoid writing expressions twice. * More expressive language (can express identity). * Duplicating is bad! ("DRY": *Don't Repeat Yourself*.) * Avoids *static* redundancy. - Avoid redundant computations. * More than *just* an optimization when it avoids exponential resources. * Avoids *dynamic* redundancy. * BNF: ::= | { + } | { - } | { * } | { / } | { with { } } | Note that we had to introduce two new rules: one for introducing an identifier, and one for using it. * Type definition: (define-type WAE [Num Number] [Add WAE WAE] [Sub WAE WAE] [Mul WAE WAE] [Div WAE WAE] [Id Symbol] [With Symbol WAE WAE]) * Parser: (: parse-sexpr : Sexpr -> WAE) ;; parses s-expressions into WAEs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) * We need to define substitution. Terms: 1. Binding Instance. 2. Scope. 3. Bound Instance. 4. Free Instance. * After lots of attempts: > e[v/i] --- To substitute an identifier `i` in an expression `e` > with an expression `v`, replace all instances of `i` that are free > in `e` with the expression `v`. * Implemented the code, and again, needed to fix a few bugs: (: subst : WAE Symbol WAE -> WAE) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (With bound-id (subst named-expr from to) (if (eq? bound-id from) bound-body (subst bound-body from to)))])) (Note that the bugs that we fixed clarify the exact way that our scopes work: in `{with {x 2} {with {x {+ x 2}} x}}`, the scope of the first `x` is the `{+ x 2}` expression.) * We then extended the AE evaluation rules: eval(...) = ... same as the AE rules ... eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) eval(id) = error! and noted the possible type problem. * The above translated into a Racket definition for an `eval` function (with a hack to avoid the type issue): (: eval : WAE -> Number) ;; evaluates WAE expressions by reducing them to numbers (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))] [(Div l r) (/ (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (Num (eval named-expr))))] [(Id name) (error 'eval "free identifier: ~s" name)])) ------------------------------------------------------------------------ # Formal Specs [Tuesday, January 21st] Note the formal definitions that were included in the WAE code. They are ways of describing pieces of our language that are more formal than plain English, but still not as formal (and as verbose) as the actual code. A formal definition of `subst`: (`N` is a ``, `E1`, `E2` are ``s, `x` is some ``, `y` is a *different* ``) N[v/x] = N {+ E1 E2}[v/x] = {+ E1[v/x] E2[v/x]} {- E1 E2}[v/x] = {- E1[v/x] E2[v/x]} {* E1 E2}[v/x] = {* E1[v/x] E2[v/x]} {/ E1 E2}[v/x] = {/ E1[v/x] E2[v/x]} y[v/x] = y x[v/x] = v {with {y E1} E2}[v/x] = {with {y E1[v/x]} E2[v/x]} {with {x E1} E2}[v/x] = {with {x E1[v/x]} E2} And a formal definition of `eval`: eval(N) = N eval({+ E1 E2}) = eval(E1) + eval(E2) eval({- E1 E2}) = eval(E1) - eval(E2) eval({* E1 E2}) = eval(E1) * eval(E2) eval({/ E1 E2}) = eval(E1) / eval(E2) eval(id) = error! eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) ------------------------------------------------------------------------ # Lazy vs Eager Evaluation [Tuesday, January 21st] As we have previously seen, there are two basic approaches for evaluation: either eager or lazy. In lazy evaluation, bindings are used for sort of textual references --- it is only for avoiding writing an expression twice, but the associated computation is done twice anyway. In eager evaluation, we eliminate not only the textual redundancy, but also the computation. Which evaluation method did our evaluator use? The relevant piece of formalism is the treatment of `with`: eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) And the matching piece of code is: [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (Num (eval named-expr))))] How do we make this lazy? In the formal equation: eval({with {x E1} E2}) = eval(E2[E1/x]) and in the code: (: eval : WAE -> Number) ;; evaluates WAE expressions by reducing them to numbers (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id named-expr))] ;*** no eval and no Num wrapping [(Id name) (error 'eval "free identifier: ~s" name)])) We can verify the way this works by tracing `eval` (compare the trace you get for the two versions): > (trace eval) ; (put this in the definitions window) > (run "{with {x {+ 1 2}} {* x x}}") Ignoring the traces for now, the modified WAE interpreter works as before, specifically, all tests pass. So the question is whether the language we get is actually different than the one we had before. One difference is in execution speed, but we can't really notice a difference, and we care more about meaning. Is there any program that will run differently in the two languages? The main feature of the lazy evaluator is that it is not evaluating the named expression until it is actually needed. As we have seen, this leads to duplicating computations if the bound identifier is used more than once --- meaning that it does not eliminate the dynamic redundancy. But what if the bound identifier is not used at all? In that case the named expression simply evaporates. This is a good hint at an expression that behaves differently in the two languages --- if we add division to both languages, we get a different result when we try running: {with {x {/ 8 0}} 7} The eager evaluator stops with an error when it tries evaluating the division --- and the lazy evaluator simply ignores it. Even without division, we get a similar behavior for {with {x y} 7} but it is questionable whether the fact that this evaluates to 7 is correct behavior --- we really want to forbid program that use free variable. Furthermore, there is an issue with name capturing --- we don't want to substitute an expression into a context that captures some of its free variables. But our substitution allows just that, which is usually not a problem because by the time we do the substitution, the named expression should not have free variables that need to be replaced. However, consider evaluating this program: {with {y x} {with {x 2} {+ x y}}} under the two evaluation regimens: the eager version stops with an error, and the lazy version succeed. This points at a bug in our substitution, or rather not dealing with an issue that we do not encounter. So the summary is: as long as the initial program is correct, both evaluation regimens produce the same results. If a program contains free variables, they might get captured in a naive lazy evaluator implementation (but this is a bug that should be fixed). Also, there are some cases where eager evaluation runs into a run-time problem which does not happen in a lazy evaluator because the expression is not used. It is possible to prove that when you evaluate an expression, if there is an error that can be avoided, lazy evaluation will always avoid it, whereas an eager evaluator will always run into it. On the other hand, lazy evaluators are usually slower than eager evaluator, so it's a speed vs. robustness trade-off. Note that with lazy evaluation we say that an identifier is bound to an expression rather than a value. (Again, this is why the eager version needed to wrap `eval`'s result in a `Num` and this one doesn't.) (It is possible to change things and get a more well behaved substitution, we basically will need to find if a capture might happen, and rename things to avoid it. For example, {with {y E1} E2}[v/x] if `x' and `y' are equal = {with {y E1[v/x]} E2} = {with {x E1[v/x]} E2} if `y' has a free occurrence in `v' = {with {y1 E1[v/x]} E2[y1/y][v/x]} ; `y1' is "fresh" otherwise = {with {x E1[v/x]} E2[v/x]} With this, we might have gone through this path in evaluating the above: {with {y x} {with {x 2} {+ x y}}} {with {x₁ 2} {+ x₁ x}} ; note that x₁ is a fresh name, not x {+ 2 x} error: free `x` But you can see that this is much more complicated (more code: requires a `free-in` predicate, being able to invent new *fresh* names, etc). And it's not even the end of that story...) ------------------------------------------------------------------------ # de Bruijn Indexes [Tuesday, January 21st] This whole story revolves around names, specifically, name capture is a problem that should always be avoided (it is one major source of PL headaches). But are names the only way we can use bindings? There is a least one alternative way: note that the only thing we used names for are for references. We don't really care what the name is, which is pretty obvious when we consider the two WAE expressions: {with {x 5} {+ x x}} {with {y 5} {+ y y}} or the two Racket function definitions: (define (foo x) (list x x)) (define (foo y) (list y y)) Both of these show a pair of expressions that we should consider as equal in some sense (this is called "alpha-equality"). The only thing we care about is what variable points where: the binding structure is the only thing that matters. In other words, as long as DrRacket produces the same arrows when we use Check Syntax, we consider the program to be the same, regardless of name choices (for argument names and local names, not for global names like `foo` in the above). The alternative idea uses this principle: if all we care about is where the arrows go, then simply get rid of the names... Instead of referencing a binding through its name, just specify which of the surrounding scopes we want to refer to. For example, instead of: {with {x 5} {with {y 6} {+ x y}}} we can use a new "reference" syntax --- `[N]` --- and use this instead of the above: {with 5 {with 6 {+ [1] [0]}}} So the rules for `[N]` are --- `[0]` is the value bound in the current scope, `[1]` is the value from the next one up etc. Of course, to do this translation, we have to know the precise scope rules. Two more complicated examples: {with {x 5} {+ x {with {y 6} {+ x y}}}} is translated to: {with 5 {+ [0] {with 6 {+ [1] [0]}}}} (note how `x` appears as a different reference based on where it appeared in the original code.) Even more subtle: {with {x 5} {with {y {+ x 1}} {+ x y}}} is translated to: {with 5 {with {+ [0] 1} {+ [1] [0]}}} because the inner `with` does not have its own named expression in its scope, so the named expression is immediately in the scope of the outer `with`. This is called "de Bruijn Indexes": instead of referencing identifiers by their name, we use an index into the surrounding binding context. The major disadvantage, as can be seen in the above examples, is that it is not convenient for humans to work with. Specifically, the same identifier is referenced using different numbers, which makes it hard to understand what some code is doing. After all, *abstractions* are the main thing we deal with when we write programs, and having labels make the bindings structure much easier to understand than scope counts. However, practically all compilers use this for compiled code (think about stack pointers). For example, GCC compiles this code: { int x = 5; { int y = x + 1; return x + y; } } to: subl $8, %esp movl $5, -4(%ebp) ; int x = 5 movl -4(%ebp), %eax incl %eax movl %eax, -8(%ebp) ; int y = %eax movl -8(%ebp), %eax addl -4(%ebp), %eax ------------------------------------------------------------------------ # Functions & Function Values [Tuesday, January 21st] > [PLAI §4] Now that we have a form for local bindings, which forced us to deal with proper substitutions and everything that is related, we can get to functions. The concept of a function is itself very close to substitution, and to our `with` form. For example, when we write: {with {x 5} {* x x}} then the `{* x x}` body is itself parametrized over some value for `x`. If we take this expression and take out the `5`, we're left with something that has all of the necessary ingredients of a function --- a bunch of code that is parameterized over some input identifier: {with {x} {* x x}} We only need to replace `with` and use a proper name that indicates that it's a function: {fun {x} {* x x}} Now we have a new form in our language, one that should have a function as its meaning. As we have seen in the case of `with` expressions, we also need a new form to *use* these functions. We will use `call` for this, so that {call {fun {x} {* x x}} 5} will be the same as the original `with` expression that we started with --- the `fun` expression is like the `with` expression with no value, and applying it on `5` is providing that value back: {with {x 5} {* x x}} Of course, this does not help much --- all we get is a way to use local bindings that is more verbose from what we started with. What we're really missing is a way to *name* these functions. If we get the right evaluation rules, we can evaluate a `fun` expression to some value --- which will allow us to bind it to a variable using `with`. Something like this: {with {sqr {fun {x} {* x x}}} {+ {call sqr 5} {call sqr 6}}} In this expression, we say that `x` is the formal parameter (or argument), and the `5` and `6` are actual parameters (sometimes abbreviated as formals and actuals). Note that naming functions often helps, but many times there are small functions that are fine to specify without a name --- for example, consider a two-stage addition function, where there is no apparent good name for the returned function: {with {add {fun {x} {fun {y} {+ x y}}}} {call {call add 8} 9}} ------------------------------------------------------------------------ # Implementing First Class Functions [Tuesday, January 28th] > [PLAI §6] (uses some stuff from [PLAI §5], which we do later) This is a simple plan, but it is directly related to how functions are going to be used in our language. We know that `{call {fun {x} E1} E2}` is equivalent to a `with` expression, but the new thing here is that we do allow writing just the `{fun ...}` expression by itself, and therefore we need to have some meaning for it. The meaning, or the value of this expression, should roughly be "an expression that needs a value to be plugged in for `x`". In other words, our language will have these new kinds of values that contain an expression to be evaluated later on. There are three basic approaches that classify programming languages in relation to how the deal with functions: 1. First order: functions are not real values. They cannot be used or returned as values by other functions. This means that they cannot be stored in data structures. This is what most "conventional" languages used to have in the past. (You will be implementing such a language in homework 4.) An example of such a language is the Beginner Student language that is used in HtDP, where the language is intentionally first-order to help students write correct code (at the early stages where using a function as a value is usually an error). It's hard to find practical modern languages that fall in this category. 2. Higher order: functions can receive and return other functions as values. This is what you get with C and modern Fortran. 3. First class: functions are values with all the rights of other values. In particular, they can be supplied to other functions, returned from functions, stored in data structures, and new functions can be created at run-time. (And most modern languages have first class functions.) The last category is the most interesting one. Back in the old days, complex expressions were not first-class in that they could not be freely composed. This is still the case in machine-code: as we've seen earlier, to compute an expression such as (-b + sqrt(b^2 - 4*a*c)) / 2a you have to do something like this: x = b * b y = 4 * a y = y * c x = x - y x = sqrt(x) y = -b x = y + x y = 2 * a s = x / y In other words, every intermediate value needs to have its own name. But with proper ("high-level") programming languages (at least most of them...) you can just write the original expression, with no names for these values. With first-class functions something similar happens --- it is possible to have complex expressions that consume and return functions, and they do not need to be named. What we get with our `fun` expression (if we can make it work) is exactly this: it generates a function, and you can choose to either bind it to a name, or not. The important thing is that the value exists independently of a name. This has a major effect on the "personality" of a programming language as we will see. In fact, just adding this feature will make our language much more advanced than languages with just higher-order or first-order functions. ------------------------------------------------------------------------ Quick Example: the following is working JavaScript code, that uses first class functions. function foo(x) { function bar(y) { return x + y; } return bar; } function main() { var f = foo(1); var g = foo(10); return [f(2), g(2)]; } Note that the above definition of `foo` does *not* use an anonymous "lambda expression" --- in Racket terms, it's translated to (define (foo x) (define (bar y) (+ x y)) bar) The returned function is not anonymous, but it's not really named either: the `bar` name is bound only inside the body of `foo`, and outside of it that name no longer exists since it's not its scope. It gets used in the printed form if the function value is displayed, but this is merely a debugging aid. The anonymous `lambda` version that is common in Racket can be used in JavaScript too: function foo(x) { return function(y) { return x + y; } } > Side-note: GCC includes extensions that allow internal function > definitions, but it still does not have first class functions --- > trying to do the above is broken: > > #include > typedef int(*int2int)(int); > int2int foo(int x) { > int bar(int y) { return x + y; } > return bar; > } > int main() { > int2int f = foo(1); > int2int g = foo(10); > printf(">> %d, %d\n", f(2), g(2)); > } ------------------------------------------------------------------------ # The FLANG Language [Tuesday, January 28th] Now for the implementation --- we call this new language FLANG. First, the BNF: ::= | { + } | { - } | { * } | { / } | { with { } } | | { fun { } } | { call } And the matching type definition: (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] ; No named-expression [Call FLANG FLANG]) The parser for this grammar is, as usual, straightforward: (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) We also need to patch up the substitution function to deal with these things. The scoping rule for the new function form is, unsurprisingly, similar to the rule of `with`, except that there is no extra expression now, and the scoping rule for `call` is the same as for the arithmetic operators: N[v/x] = N {+ E1 E2}[v/x] = {+ E1[v/x] E2[v/x]} {- E1 E2}[v/x] = {- E1[v/x] E2[v/x]} {* E1 E2}[v/x] = {* E1[v/x] E2[v/x]} {/ E1 E2}[v/x] = {/ E1[v/x] E2[v/x]} y[v/x] = y x[v/x] = v {with {y E1} E2}[v/x] = {with {y E1[v/x]} E2[v/x]} {with {x E1} E2}[v/x] = {with {x E1[v/x]} E2} {call E1 E2}[v/x] = {call E1[v/x] E2[v/x]} {fun {y} E}[v/x] = {fun {y} E[v/x]} {fun {x} E}[v/x] = {fun {x} E} And the matching code: (: subst : FLANG Symbol FLANG -> FLANG) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (With bound-id (subst named-expr from to) (if (eq? bound-id from) bound-body (subst bound-body from to)))] [(Call l r) (Call (subst l from to) (subst r from to))] [(Fun bound-id bound-body) (if (eq? bound-id from) expr (Fun bound-id (subst bound-body from to)))])) ------------------------------------------------------------------------ Now, before we start working on an evaluator, we need to decide on what exactly do we use to represent values of this language. Before we had functions, we had only number values and we used Racket numbers to represent them. Now we have two kinds of values --- numbers and functions. It seems easy enough to continue using Racket numbers to represent numbers, but what about functions? What should be the result of evaluating {fun {x} {+ x 1}} ? Well, this is the new toy we have: it should be a function value, which is something that can be used just like numbers, but instead of arithmetic operations, we can `call` these things. What we need is a way to avoid evaluating the body expression of the function --- *delay* it --- and instead use some value that will contain this delayed expression in a way that can be used later. To accommodate this, we will change our implementation strategy a little: we will use our syntax objects for numbers (`(Num n)` instead of just `n`), which will be a little inconvenient when we do the arithmetic operations, but it will simplify life by making it possible to evaluate functions in a similar way: simply return their own syntax object as their values. The syntax object has what we need: the body expression that needs to be evaluated later when the function is called, and it also has the identifier name that should be replaced with the actual input to the function call. This means that evaluating: (Add (Num 1) (Num 2)) now yields (Num 3) and a number `(Num 5)` evaluates to `(Num 5)`. In a similar way, `(Fun 'x (Num 2))` evaluates to `(Fun 'x (Num 2))`. Why would this work? Well, because `call` will be very similar to `with` --- the only difference is that its arguments are ordered a little differently, being retrieved from the function that is applied and the argument. The formal evaluation rules are therefore treating functions like numbers, and use the syntax object to represent both values: eval(N) = N eval({+ E1 E2}) = eval(E1) + eval(E2) eval({- E1 E2}) = eval(E1) - eval(E2) eval({* E1 E2}) = eval(E1) * eval(E2) eval({/ E1 E2}) = eval(E1) / eval(E2) eval(id) = error! eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) eval(FUN) = FUN ; assuming FUN is a function expression eval({call E1 E2}) = eval(Ef[eval(E2)/x]) if eval(E1) = {fun {x} Ef} = error! otherwise Note that the last rule could be written using a translation to a `with` expression: eval({call E1 E2}) = eval({with {x E2} Ef}) if eval(E1) = {fun {x} Ef} = error! otherwise And alternatively, we could specify `with` using `call` and `fun`: eval({with {x E1} E2}) = eval({call {fun {x} E2} E1}) There is a small problem in these rules which is intuitively seen by the fact that the evaluation rule for a `call` is expected to be very similar to the one for arithmetic operations. We now have two kinds of values, so we need to check the arithmetic operation's arguments too: eval({+ E1 E2}) = N1 + N2 if eval(E1), eval(E2) evaluate to numbers N1, N2 otherwise error! ... The corresponding code is: (: eval : FLANG -> FLANG) ;*** note return type ;; evaluates FLANG expressions by reducing them to *expressions* but ;; only expressions that stand for values: only `Fun`s and `Num`s (define (eval expr) (cases expr [(Num n) expr] ;*** change here [(Add l r) (arith-op + (eval l) (eval r))] [(Sub l r) (arith-op - (eval l) (eval r))] [(Mul l r) (arith-op * (eval l) (eval r))] [(Div l r) (arith-op / (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (eval named-expr)))] ;*** no `(Num ...)' [(Id name) (error 'eval "free identifier: ~s" name)] [(Fun bound-id bound-body) expr] ;*** similar to `Num' [(Call (Fun bound-id bound-body) arg-expr) ;*** nested pattern (eval (subst bound-body ;*** just like `with' bound-id (eval arg-expr)))] [(Call something arg-expr) (error 'eval "`call' expects a function, got: ~s" something)])) Note that the `Call` case is doing the same thing we do in the `With` case. In fact, we could have just *generated* a `With` expression and evaluate that instead: ... [(Call (Fun bound-id bound-body) arg-expr) (eval (With bound-id arg-expr bound-body))] ... The `arith-op` function is in charge of checking that the input values are numbers (represented as FLANG numbers), translating them to plain numbers, performing the Racket operation, then re-wrapping the result in a `Num`. Note how its type indicates that it is a higher-order function. (: arith-op : (Number Number -> Number) FLANG FLANG -> FLANG) ;; gets a Racket numeric binary operator, and uses it within a FLANG ;; `Num' wrapper (note the H.O. type, and note the hack of the `val` ;; name which is actually an AST that represents a runtime value) (define (arith-op op val1 val2) (Num (op (Num->number val1) (Num->number val2)))) It uses the following function to convert FLANG numbers to Racket numbers. (Note that `else` is almost always a bad idea since it can prevent the compiler from showing you places to edit code --- but this case is an exception since we never want to deal with anything other than `Num`s.) The reason that this function is relatively trivial is that we chose the easy way and represented numbers using Racket numbers, but we could have used strings or anything else. (: Num->number : FLANG -> Number) ;; convert a FLANG number to a Racket one (define (Num->number e) (cases e [(Num n) n] [else (error 'arith-op "expected a number, got: ~s" e)])) We can also make things a little easier to use if we make `run` convert the result to a number: (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str))]) (cases result [(Num n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) Adding few simple tests we get: ;; The Flang interpreter #lang pl #| The grammar: ::= | { + } | { - } | { * } | { / } | { with { } } | | { fun { } } | { call } Evaluation rules: subst: N[v/x] = N {+ E1 E2}[v/x] = {+ E1[v/x] E2[v/x]} {- E1 E2}[v/x] = {- E1[v/x] E2[v/x]} {* E1 E2}[v/x] = {* E1[v/x] E2[v/x]} {/ E1 E2}[v/x] = {/ E1[v/x] E2[v/x]} y[v/x] = y x[v/x] = v {with {y E1} E2}[v/x] = {with {y E1[v/x]} E2[v/x]} ; if y =/= x {with {x E1} E2}[v/x] = {with {x E1[v/x]} E2} {call E1 E2}[v/x] = {call E1[v/x] E2[v/x]} {fun {y} E}[v/x] = {fun {y} E[v/x]} ; if y =/= x {fun {x} E}[v/x] = {fun {x} E} eval: eval(N) = N eval({+ E1 E2}) = eval(E1) + eval(E2) \ if both E1 and E2 eval({- E1 E2}) = eval(E1) - eval(E2) \ evaluate to numbers eval({* E1 E2}) = eval(E1) * eval(E2) / otherwise error! eval({/ E1 E2}) = eval(E1) / eval(E2) / eval(id) = error! eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) eval(FUN) = FUN ; assuming FUN is a function expression eval({call E1 E2}) = eval(Ef[eval(E2)/x]) if eval(E1)={fun {x} Ef}, otherwise error! |# (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) (: subst : FLANG Symbol FLANG -> FLANG) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (With bound-id (subst named-expr from to) (if (eq? bound-id from) bound-body (subst bound-body from to)))] [(Call l r) (Call (subst l from to) (subst r from to))] [(Fun bound-id bound-body) (if (eq? bound-id from) expr (Fun bound-id (subst bound-body from to)))])) (: Num->number : FLANG -> Number) ;; convert a FLANG number to a Racket one (define (Num->number e) (cases e [(Num n) n] [else (error 'arith-op "expected a number, got: ~s" e)])) (: arith-op : (Number Number -> Number) FLANG FLANG -> FLANG) ;; gets a Racket numeric binary operator, and uses it within a FLANG ;; `Num' wrapper (define (arith-op op val1 val2) (Num (op (Num->number val1) (Num->number val2)))) (: eval : FLANG -> FLANG) ;; evaluates FLANG expressions by reducing them to *expressions* (define (eval expr) (cases expr [(Num n) expr] [(Add l r) (arith-op + (eval l) (eval r))] [(Sub l r) (arith-op - (eval l) (eval r))] [(Mul l r) (arith-op * (eval l) (eval r))] [(Div l r) (arith-op / (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (eval named-expr)))] [(Id name) (error 'eval "free identifier: ~s" name)] [(Fun bound-id bound-body) expr] [(Call (Fun bound-id bound-body) arg-expr) (eval (subst bound-body bound-id (eval arg-expr)))] [(Call something arg-expr) (error 'eval "`call' expects a function, got: ~s" something)])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str))]) (cases result [(Num n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) ------------------------------------------------------------------------ ------------------------------------------------------------------------ # The FLANG Language (contd.) [Tuesday, January 28th] There is still a problem with this version. First a question --- if `call` is similar to arithmetic operations (and to `with` in what it actually does), then how come the code is different enough that it doesn't even need an auxiliary function? Second question: what *should* happen if we evaluate these code snippets: (run "{with {add {fun {x} {fun {y} {+ x y}}}} {call {call add 8} 9}}") (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") Third question, what *will* happen if we do the above? What we're missing is an evaluation of the function expression, in case it's not a literal `fun` form. The following fixes this: (: eval : FLANG -> FLANG) ;; evaluates FLANG expressions by reducing them to *expressions* (define (eval expr) (cases expr [(Num n) expr] [(Add l r) (arith-op + (eval l) (eval r))] [(Sub l r) (arith-op - (eval l) (eval r))] [(Mul l r) (arith-op * (eval l) (eval r))] [(Div l r) (arith-op / (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (eval named-expr)))] [(Id name) (error 'eval "free identifier: ~s" name)] [(Fun bound-id bound-body) expr] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr)]) ;*** need to evaluate this! (cases fval [(Fun bound-id bound-body) (eval (subst bound-body bound-id (eval arg-expr)))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) The complete code is: ;;; ---<<>>-------------------------------------------------- ;; The Flang interpreter #lang pl #| The grammar: ::= | { + } | { - } | { * } | { / } | { with { } } | | { fun { } } | { call } Evaluation rules: subst: N[v/x] = N {+ E1 E2}[v/x] = {+ E1[v/x] E2[v/x]} {- E1 E2}[v/x] = {- E1[v/x] E2[v/x]} {* E1 E2}[v/x] = {* E1[v/x] E2[v/x]} {/ E1 E2}[v/x] = {/ E1[v/x] E2[v/x]} y[v/x] = y x[v/x] = v {with {y E1} E2}[v/x] = {with {y E1[v/x]} E2[v/x]} ; if y =/= x {with {x E1} E2}[v/x] = {with {x E1[v/x]} E2} {call E1 E2}[v/x] = {call E1[v/x] E2[v/x]} {fun {y} E}[v/x] = {fun {y} E[v/x]} ; if y =/= x {fun {x} E}[v/x] = {fun {x} E} eval: eval(N) = N eval({+ E1 E2}) = eval(E1) + eval(E2) \ if both E1 and E2 eval({- E1 E2}) = eval(E1) - eval(E2) \ evaluate to numbers eval({* E1 E2}) = eval(E1) * eval(E2) / otherwise error! eval({/ E1 E2}) = eval(E1) / eval(E2) / eval(id) = error! eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) eval(FUN) = FUN ; assuming FUN is a function expression eval({call E1 E2}) = eval(Ef[eval(E2)/x]) if eval(E1)={fun {x} Ef}, otherwise error! |# (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) (: subst : FLANG Symbol FLANG -> FLANG) ;; substitutes the second argument with the third argument in the ;; first argument, as per the rules of substitution; the resulting ;; expression contains no free instances of the second argument (define (subst expr from to) (cases expr [(Num n) expr] [(Add l r) (Add (subst l from to) (subst r from to))] [(Sub l r) (Sub (subst l from to) (subst r from to))] [(Mul l r) (Mul (subst l from to) (subst r from to))] [(Div l r) (Div (subst l from to) (subst r from to))] [(Id name) (if (eq? name from) to expr)] [(With bound-id named-expr bound-body) (With bound-id (subst named-expr from to) (if (eq? bound-id from) bound-body (subst bound-body from to)))] [(Call l r) (Call (subst l from to) (subst r from to))] [(Fun bound-id bound-body) (if (eq? bound-id from) expr (Fun bound-id (subst bound-body from to)))])) (: Num->number : FLANG -> Number) ;; convert a FLANG number to a Racket one (define (Num->number e) (cases e [(Num n) n] [else (error 'arith-op "expected a number, got: ~s" e)])) (: arith-op : (Number Number -> Number) FLANG FLANG -> FLANG) ;; gets a Racket numeric binary operator, and uses it within a FLANG ;; `Num' wrapper (define (arith-op op val1 val2) (Num (op (Num->number val1) (Num->number val2)))) (: eval : FLANG -> FLANG) ;; evaluates FLANG expressions by reducing them to *expressions* (define (eval expr) (cases expr [(Num n) expr] [(Add l r) (arith-op + (eval l) (eval r))] [(Sub l r) (arith-op - (eval l) (eval r))] [(Mul l r) (arith-op * (eval l) (eval r))] [(Div l r) (arith-op / (eval l) (eval r))] [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (eval named-expr)))] [(Id name) (error 'eval "free identifier: ~s" name)] [(Fun bound-id bound-body) expr] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr)]) (cases fval [(Fun bound-id bound-body) (eval (subst bound-body bound-id (eval arg-expr)))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str))]) (cases result [(Num n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {add {fun {x} {fun {y} {+ x y}}}} {call {call add 8} 9}}") => 17) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Introducing Racket's `lambda` [Tuesday, January 28th] `fun` & `lambda` difference between lambda and simple values not being able to do recursive functions with `let` let* as a derived form let with lambda in Racket --> can be a derived form how `if` can be used to implement `and` and `or` as derived forms Newtonian syntax vs. a lambda expression. Don't be fooled into making a bogus connection between Racket's syntax, and its `unique` powers... The fact is that it is not the only language that has this capability. For example, this: (define (f g) (g 2 3)) (f +) ==> 5 (f *) ==> 6 (f (lambda (x y) (+ (square x) (square y)))) ==> 13 Can be written in JavaScript like this: function f(g) { return g(2,3); } function square(x) { return x*x; } console.log(f(function (x,y) { return square(x) + square(y); })); or in ES6 JavaScript: let f = (g) => g(2,3); let square = (x) => x*x; console.log(f((x,y) => square(x) + square(y))); In Perl: sub f { my ($g) = @_; return $g->(2,3); } sub square { my ($x) = @_; return $x * $x; } print f(sub { my ($x, $y) = @_; return square($x) + square($y); }); In Ruby: def f(g) g.call(2,3) end def square(x) x*x end puts f(lambda{|x,y| square(x) + square(y)}) etc. Even [Java has lambda expressions], and recently [C++ added them too]. [Java has lambda expressions]: http://www.drdobbs.com/jvm/lambda-expressions-in-java-8/240166764 [C++ added them too]: http://www.cprogramming.com/c++11/c++11-lambda-closures.html ------------------------------------------------------------------------ # Using Functions as Objects [Tuesday, January 28th] A very important aspect of Racket --- using "higher order" functions --- functions that get and return functions. Here is a very simple example: (define (f x) (lambda () x)) (define a (f 2)) (a) --> 2 (define b (f 3)) (b) --> 3 Note: what we get is actually an object that remembers (by the substitution we're doing) a number. How about: (define aa (f a)) (aa) --> # (this is a) ((aa)) --> 2 Take this idea to the next level: (define (kons x y) (lambda (b) (if b x y))) (define (kar p) (p #t)) (define (kdr p) (p #f)) (define a (kons 1 2)) (define b (kons 3 4)) (list (kar a) (kdr a)) (list (kar b) (kdr b)) Or, with types: (: kons : (All (A B) A B -> (Boolean -> (U A B)))) (define (kons x y) (lambda (b) (if b x y))) (: kar : (All (T) (Boolean -> T) -> T)) (define (kar p) (p #t)) (: kdr : (All (T) (Boolean -> T) -> T)) (define (kdr p) (p #f)) (define a (kons 1 2)) (define b (kons 3 4)) (list (kar a) (kdr a)) (list (kar b) (kdr b)) Even more --- why should the internal function expect a boolean and choose what to return? We can simply expect a function that will take the two values and return one: (define (kons x y) (lambda (s) (s x y))) (define (kar p) (p (lambda (x y) x))) (define (kdr p) (p (lambda (x y) y))) (define a (kons 1 2)) (define b (kons 3 4)) (list (kar a) (kdr a)) (list (kar b) (kdr b)) And a typed version, using our own constructor to make it a little less painful: (define-type (Kons A B) = ((A B -> (U A B)) -> (U A B))) (: kons : (All (A B) A B -> (Kons A B))) (define (kons x y) (lambda (s) (s x y))) (: kar : (All (A B) (Kons A B) -> (U A B))) (define (kar p) (p (lambda (x y) x))) (: kdr : (All (A B) (Kons A B) -> (U A B))) (define (kdr p) (p (lambda (x y) y))) (define a (kons 1 2)) (define b (kons 3 4)) (list (kar a) (kdr a)) (list (kar b) (kdr b)) Note that the `Kons` type definition is the same as: (define-type Kons = (All (A B) (A B -> (U A B)) -> (U A B))) so `All` is to polymorphic type definitions what `lambda` is for function definitions. Finally in JavaScript: function kons(x,y) { return function(s) { return s(x, y); } } function kar(p) { return p(function(x,y){ return x; }); } function kdr(p) { return p(function(x,y){ return y; }); } a = kons(1,2); b = kons(3,4); console.log('a = <' + kar(a) + ',' + kdr(a) + '>' ); console.log('b = <' + kar(b) + ',' + kdr(b) + '>' ); Or with ES6 *arrow functions*, the function definitionss become: const kons = (x,y) => s => s(x,y); const kar = p => p((x,y) => x); const kdr = p => p((x,y) => y); ------------------------------------------------------------------------ # Using `define-type` for new "type aliases" [Tuesday, January 28th] As seen in these examples, there is another way to use `define-type`, using a `=` to create a new type name "alias" for an *existing* type. For example: (define-type Strings = (Listof String)) These uses of `define-type` do not define any new kind of type, they are essentially a convenience tool for making code shorter and more readable. (define-type NumericFunction = Number -> Number) (: square : NumericFunction) (define (square n) (* n n)) Note in particular that this can also be used to define "alias type constructors" too: somewhat similar to creating new "type functions". For example: (define-type (BinaryFun In Out) = In In -> Out) (: diagonal : (BinaryFun Natural Number)) (define (diagonal width height) (sqrt (+ (* width width) (* height height)))) This is something that we will only need in a few rare cases. ------------------------------------------------------------------------ # Currying [Tuesday, January 28th] A *curried* function is a function that, instead of accepting two (or more) arguments, accepts only one and returns a function that accepts the rest. For example: (: plus : Number -> (Number -> Number)) (define (plus x) (lambda (y) (+ x y))) It's easy to write functions for translating between normal and curried versions. (define (currify f) (lambda (x) (lambda (y) (f x y)))) Typed version of that, with examples: (: currify : (All (A B C) (A B -> C) -> (A -> (B -> C)))) ;; convert a double-argument function to a curried one (define (currify f) (lambda (x) (lambda (y) (f x y)))) (: add : Number Number -> Number) (define (add x y) (+ x y)) (: plus : Number -> (Number -> Number)) (define plus (currify add)) (test ((plus 1) 2) => 3) (test (((currify add) 1) 2) => 3) (test (map (plus 1) '(1 2 3)) => '(2 3 4)) (test (map ((currify add) 1) '(1 2 3)) => '(2 3 4)) (test (map ((currify +) 1) '(1 2 3)) => '(2 3 4)) Usages --- common with H.O. functions like map, where we want to *fix* one argument. When dealing with such higher-order code, the types are very helpful, since every arrow corresponds to a function: (: currify : (All (A B C) (A B -> C) -> (A -> (B -> C)))) It is common to make the `->` function type associate to the right, so you can find this type written as: currify : (A B -> C) -> (A -> B -> C) or even as currify : (A B -> C) -> A -> B -> C but that can be a little confusing... ------------------------------------------------------------------------ # Using Higher-Order & Anonymous Functions [Tuesday, January 28th] Say that we have a function for estimating derivatives of a function at a specific point: (define dx 0.01) (: deriv : (Number -> Number) Number -> Number) ;; compute the derivative of `f' at the given point `x' (define (deriv f x) (/ (- (f (+ x dx)) (f x)) dx)) (: integrate : (Number -> Number) Number -> Number) ;; compute an integral of `f' at the given point `x' (define (integrate f x) (: loop : Number Number -> Number) (define (loop y acc) (if (> y x) (* acc dx) (loop (+ y dx) (+ acc (f y))))) (loop 0 0)) And say that we want to try out various functions given some `plot` function that draws graphs of numeric functions, for example: (plot sin) The problem is that `plot` expects a single `(Number -> Number)` function --- if we want to try it with a derivative, we can do this: (: sin-deriv : Number -> Number) ;; the derivative of sin (define sin-deriv (lambda (x) (deriv sin x))) (plot sin-deriv) But this will get very tedious very fast --- it is much simpler to use an anonymous function: (plot (lambda (x) (deriv sin x))) we can even verify that our derivative is correct by comparing a known function to its derivative (plot (lambda (x) (- (deriv sin x) (cos x)))) But it's still not completely natural to do these things --- you need to explicitly combine functions, which is not too convenient. Instead of doing this, we can write H.O. functions that will work with functional inputs and outputs. For example, we can write a function to subtract functions: (: fsub : (Number -> Number) (Number -> Number) -> (Number -> Number)) ;; subtracts two numeric 1-argument functions (define (fsub f g) (lambda (x) (- (f x) (g x)))) and the same for the derivative: (: fderiv : (Number -> Number) -> (Number -> Number)) ;; compute the derivative function of `f' (define (fderiv f) (lambda (x) (deriv f x))) Now we can try the same in a much easier way: (plot (fsub (fderiv sin) cos)) More than that --- our `fderiv` could be created from `deriv` automatically: (: currify : (All (A B C) (A B -> C) -> (A -> B -> C))) ;; convert a double-argument function to a curried one (define (currify f) (lambda (x) (lambda (y) (f x y)))) (: fderiv : (Number -> Number) -> (Number -> Number)) ;; compute the derivative function of `f' (define fderiv (currify deriv)) Same principle with `fsub`: we can write a function that converts a binary arithmetical function into a function that operates on unary numeric function. But to make things more readable we can define new types for unary and binary numeric functions: (define-type UnaryFun = (Number -> Number)) (define-type BinaryFun = (Number Number -> Number)) (: binop->fbinop : BinaryFun -> (UnaryFun UnaryFun -> UnaryFun)) ;; turns an arithmetic binary operator to a function operator (define (binop->fbinop op) (lambda (f g) (lambda (x) (op (f x) (g x))))) (: fsub : UnaryFun UnaryFun -> UnaryFun) ;; functional pointwise subtraction (define fsub (binop->fbinop -)) We can do this with anything --- developing a rich library of functions and functionals (functions over functions) is extremely easy... Here's a pretty extensive yet very short library of functions: #lang pl untyped (define (currify f) (lambda (x) (lambda (y) (f x y)))) (define (binop->fbinop op) (lambda (f g) (lambda (x) (op (f x) (g x))))) (define (compose f g) (lambda (x) (f (g x)))) (define dx 0.01) (define (deriv f x) (/ (- (f (+ x dx)) (f x)) dx)) (define (integrate f x) (define over? (if (< x 0) < >)) (define step (if (< x 0) - +)) (define add (if (< x 0) - +)) (define (loop y acc) (if (over? y x) (* acc dx) (loop (step y dx) (add acc (f y))))) (loop 0 0)) (define fadd (binop->fbinop +)) (define fsub (binop->fbinop -)) (define fmul (binop->fbinop *)) (define fdiv (binop->fbinop /)) (define fderiv (currify deriv)) (define fintegrate (currify integrate)) ;; ... This is written in the "untyped dialect" of the class language, but it should be easy now to add the types. Examples: ;; want to verify that `integrate' is the opposite of `deriv': ;; take a function, subtract it from its derivative's integral (plot (fsub sin (fintegrate (fderiv sin)))) ;; want to magnify the errors? -- here's how you magnify: (plot (compose ((currify *) 5) sin)) ;; so: (plot (compose ((currify *) 20) (fsub sin (fintegrate (fderiv sin))))) ------------------------------------------------------------------------ ## Side-note: "Point-Free" combinators [Tuesday, January 28th] > Forming functions without using `lambda` (or an implicit `lambda` > using a `define` syntactic sugar) is called *point-free style*. It's > especially popular in Haskell, where it is easier to form functions > this way because of implicit currying and a large number of higher > level function combinators. If used too much, it can easily lead to > obfuscated code. ------------------------------------------------------------------------ ------------------------------------------------------------------------ ## This is not Runtime Code Generation [Tuesday, January 28th] All of this is similar to run-time code generation, but not really. The only thing that `fderiv` does is take a function and store it somewhere in the returned function, then when that function receives a number, it uses the stored function and send it to deriv with the number. We could simply write deriv as what `fderiv` is --- which is the *real* derivative function: (define (deriv f) (lambda (x) (/ (- (f (+ x dx)) (f x)) dx))) but again, this is not faster or slower than the plain `deriv`. However, there are some situations where we can do some of the computation on the first-stage argument, saving work from the second stage. Here is a cooked-to-exaggeration example --- we want a function that receives two inputs `x`, `y` and returns `fib(x)*y`, but we must use a stupid `fib`: (define (fib n) (if (<= n 1) n (+ (fib (- n 1)) (fib (- n 2))))) The function we want is: (define (bogus x y) (* (fib x) y)) If we currify it as usual (or just use `currify`), we get: (define (bogus x) (lambda (y) (* (fib x) y))) And try this several times: (define bogus24 (bogus 24)) (map bogus24 '(1 2 3 4 5)) But in the definition of `bogus`, notice that `(fib x)` does not depend on `y` --- so we can rewrite it a little differently: (define (bogus x) (let ([fibx (fib x)]) (lambda (y) (* fibx y)))) and trying the above again is much faster now: (define bogus24 (bogus 24)) (map bogus24 '(1 2 3 4 5)) This is therefore not doing any kind of runtime code generation, but it *enables* doing similar optimizations in our code. A proper RTCG facility would recompile the curried function for a given first input, and (hopefully) automatically achieve the optimization that we did in a manual way. ------------------------------------------------------------------------ # Substitution Caches [Tuesday, January 28th] > [PLAI §5] (called "deferred substitutions" there) Evaluating using substitutions is very inefficient --- at each scope, we copy a piece of the program AST. This includes all function calls which implies an impractical cost (function calls should be *cheap*!). To get over this, we want to use a cache of substitutions. Basic idea: we begin evaluating with no cached substitutions, then collect them as we encounter bindings. [Implies another change for our evaluator: we don't really substitute identifiers until we get to them; when we reach an identifier, it is no longer an error -- we must consult the substitution cache.] ------------------------------------------------------------------------ # Implementation of Cache Functionality [Tuesday, February 4th] First, we need a type for a substitution cache. For this we will use a list of lists of two elements each --- a name and its value FLANG: ;; a type for substitution caches: (define-type SubstCache = (Listof (List Symbol FLANG))) We need to have an empty substitution cache, a way to extend it, and a way to look things up: (: empty-subst : SubstCache) (define empty-subst null) (: extend : Symbol FLANG SubstCache -> SubstCache) ;; extend a given substitution cache with a new mapping (define (extend id expr sc) (cons (list id expr) sc)) (: lookup : Symbol SubstCache -> FLANG) ;; lookup a symbol in a substitution cache, return the value it is ;; bound to (or throw an error if it isn't bound) (define (lookup name sc) (cond [(null? sc) (error 'lookup "no binding for ~s" name)] [(eq? name (first (first sc))) (second (first sc))] [else (lookup name (rest sc))])) Actually, the reason to use such list of lists is that Racket has a built-in function called `assq` that will do this kind of search (`assq` is a search in an association list using `eq?` for the key comparison). This is a version of `lookup` that uses `assq`: (define (lookup name sc) (let ([cell (assq name sc)]) (if cell (second cell) (error 'lookup "no binding for ~s" name)))) ------------------------------------------------------------------------ # Formal Rules for Cached Substitutions [Tuesday, February 4th] The formal evaluation rules are now different. Evaluation carries along a *substitution cache* that begins its life as empty: so `eval` needs an extra argument. We begin by writing the rules that deal with the cache, and use the above function names for simplicity --- the behavior of the three definitions can be summed up in a single rule for `lookup`: lookup(x,empty-subst) = error! lookup(x,extend(x,E,sc)) = E lookup(x,extend(y,E,sc)) = lookup(x,sc) if `x' is not `y' And now we can write the new rules for `eval` eval(N,sc) = N eval({+ E1 E2},sc) = eval(E1,sc) + eval(E2,sc) eval({- E1 E2},sc) = eval(E1,sc) - eval(E2,sc) eval({* E1 E2},sc) = eval(E1,sc) * eval(E2,sc) eval({/ E1 E2},sc) = eval(E1,sc) / eval(E2,sc) eval(x,sc) = lookup(x,sc) eval({with {x E1} E2},sc) = eval(E2,extend(x,eval(E1,sc),sc)) eval({fun {x} E},sc) = {fun {x} E} eval({call E1 E2},sc) = eval(Ef,extend(x,eval(E2,sc),sc)) if eval(E1,sc) = {fun {x} Ef} = error! otherwise Note that there is no mention of `subst` --- the whole point is that we don't really do substitution, but use the cache instead. The `lookup` rules, and the places where `extend` is used replaces `subst`, and therefore specifies our scoping rules. Also note that the rule for `call` is still very similar to the rule for `with`, but it looks like we have lost something --- the interesting bit with substituting into `fun` expressions. ------------------------------------------------------------------------ # Evaluating with Substitution Caches [Tuesday, February 4th] Implementing the new `eval` is easy now --- it is extended in the same way that the formal `eval` rule is extended: (: eval : FLANG SubstCache -> FLANG) ;; evaluates FLANG expressions by reducing them to expressions (define (eval expr sc) (cases expr [(Num n) expr] [(Add l r) (arith-op + (eval l sc) (eval r sc))] [(Sub l r) (arith-op - (eval l sc) (eval r sc))] [(Mul l r) (arith-op * (eval l sc) (eval r sc))] [(Div l r) (arith-op / (eval l sc) (eval r sc))] [(With bound-id named-expr bound-body) (eval bound-body (extend bound-id (eval named-expr sc) sc))] [(Id name) (lookup name sc)] [(Fun bound-id bound-body) expr] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr sc)]) (cases fval [(Fun bound-id bound-body) (eval bound-body (extend bound-id (eval arg-expr sc) sc))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) Again, note that we don't need `subst` anymore, but the rest of the code (the data type definition, parsing, and `arith-op`) is exactly the same. Finally, we need to make sure that `eval` is initially called with an empty cache. This is easy to change in our main `run` entry point: (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) empty-subst)]) (cases result [(Num n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) The full code (including the same tests, but not including formal rules for now) follows. Note that one test does not pass. #lang pl (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; a type for substitution caches: (define-type SubstCache = (Listof (List Symbol FLANG))) (: empty-subst : SubstCache) (define empty-subst null) (: extend : Symbol FLANG SubstCache -> SubstCache) ;; extend a given substitution cache with a new mapping (define (extend name val sc) (cons (list name val) sc)) (: lookup : Symbol SubstCache -> FLANG) ;; lookup a symbol in a substitution cache, return the value it is ;; bound to (or throw an error if it isn't bound) (define (lookup name sc) (let ([cell (assq name sc)]) (if cell (second cell) (error 'lookup "no binding for ~s" name)))) (: Num->number : FLANG -> Number) ;; convert a FLANG number to a Racket one (define (Num->number e) (cases e [(Num n) n] [else (error 'arith-op "expected a number, got: ~s" e)])) (: arith-op : (Number Number -> Number) FLANG FLANG -> FLANG) ;; gets a Racket numeric binary operator, and uses it within a FLANG ;; `Num' wrapper (define (arith-op op val1 val2) (Num (op (Num->number val1) (Num->number val2)))) (: eval : FLANG SubstCache -> FLANG) ;; evaluates FLANG expressions by reducing them to expressions (define (eval expr sc) (cases expr [(Num n) expr] [(Add l r) (arith-op + (eval l sc) (eval r sc))] [(Sub l r) (arith-op - (eval l sc) (eval r sc))] [(Mul l r) (arith-op * (eval l sc) (eval r sc))] [(Div l r) (arith-op / (eval l sc) (eval r sc))] [(With bound-id named-expr bound-body) (eval bound-body (extend bound-id (eval named-expr sc) sc))] [(Id name) (lookup name sc)] [(Fun bound-id bound-body) expr] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr sc)]) (cases fval [(Fun bound-id bound-body) (eval bound-body (extend bound-id (eval arg-expr sc) sc))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) empty-subst)]) (cases result [(Num n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => "???") (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Dynamic and Lexical Scopes [Tuesday, February 4th] This seems like it should work, and it even worked on a few examples, except for one which was hard to follow. Seems like we have a bug... Now we get to a tricky issue that managed to be a problem for *lots* of language implementors, including the first version of Lisp. Lets try to run the following expression --- try to figure out what it will evaluate to: (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") We expect it to return `7` (at least I do!), but we get `9` instead... The question is --- *should* it return `9`? What we have arrived to is called *dynamic scope*. Scope is determined by the dynamic run-time environment (which is represented by our substitution cache). This is *almost always* undesirable, as I hope to convince you. Before we start, we define two options for a programming language: * Static Scope (also called Lexical Scope): In a language with static scope, each identifier gets its value from the scope of its definition, not its use. * Dynamic Scope: In a language with dynamic scope, each identifier gets its value from the scope of its use, not its definition. Racket uses lexical scope, our new evaluator uses dynamic, the old substitution-based evaluator was static etc. As a side-remark, Lisp began its life as a dynamically-scoped language. The artifacts of this were (sort-of) dismissed as an implementation bug. When Scheme was introduced, it was the first Lisp dialect that used strictly lexical scoping, and Racket is obviously doing the same. (Some Lisp implementations used dynamic scope for interpreted code and lexical scope for compiled code!) In fact, Emacs Lisp is the only *live* dialects of Lisp that is still dynamically scoped by default. To see this, compare a version of the above code in Racket: (let ((x 3)) (let ((f (lambda (y) (+ x y)))) (let ((x 5)) (f 4)))) and the Emacs Lisp version (which looks almost the same): (let ((x 3)) (let ((f (lambda (y) (+ x y)))) (let ((x 5)) (funcall f 4)))) which also happens when we use another function on the way: (defun blah (func val) (funcall func val)) (let ((x 3)) (let ((f (lambda (y) (+ x y)))) (let ((x 5)) (blah f 4)))) and note that renaming identifiers can lead to different code --- change that `val` to `x`: (defun blah (func x) (funcall func x)) (let ((x 3)) (let ((f (lambda (y) (+ x y)))) (let ((x 5)) (blah f 4)))) and you get `8` because the argument name changed the `x` that the internal function sees! Consider also this Emacs Lisp function: (defun return-x () x) which has no meaning by itself (`x` is unbound), (return-x) but can be given a dynamic meaning using a `let`: (let ((x 5)) (return-x)) or a function application: (defun foo (x) (return-x)) (foo 5) There is also a dynamically-scoped language in the course languages: #lang pl dynamic (define x 123) (define (getx) x) (define (bar1 x) (getx)) (define (bar2 y) (getx)) (test (getx) => 123) (test (let ([x 456]) (getx)) => 456) (test (getx) => 123) (test (bar1 999) => 999) (test (bar2 999) => 123) (define (foo x) (define (helper) (+ x 1)) helper) (test ((foo 0)) => 124) ;; and *much* worse: (define (add x y) (+ x y)) (test (let ([+ *]) (add 6 7)) => 42) Note how bad the last example gets: you basically cannot call any function and know in advance what it will do. There are some cases where dynamic scope can be useful in that it allows you to "remotely" customize any piece of code. A good example of where this is taken to an extreme is Emacs: originally, it was based on an ancient Lisp dialect that was still dynamically scoped, but it retained this feature even when practically all Lisp dialects moved on to having lexical scope by default. The reason for this is that the danger of dynamic scope is also a way to make a very open system where almost anything can be customized by changing it "remotely". Here's a concrete example for a similar kind of dynamic scope usage that makes a very hackable and open system: #lang pl dynamic (define tax% 6.25) (define (with-tax n) (+ n (* n (/ tax% 100)))) (with-tax 10) ; how much do we pay? (let ([tax% 17.0]) (with-tax 10)) ; how much would we pay in Israel? ;; make that into a function (define il-tax% 18.0) (define (us-over-il-saving n) (- (let ([tax% il-tax%]) (with-tax n)) (with-tax n))) (us-over-il-saving 10) ;; can even control that: how much would we save if ;; the tax in israel went down one percent? (let ([il-tax% (- il-tax% 1)]) (us-over-il-saving 10)) ;; or change both: how much savings in NH instead of MA? (let ((tax% 0.0) (il-tax% tax%)) (us-over-il-saving 1000)) Obviously, this power to customize everything is also the main source of problems with getting no guarantees for code. A common way to get the best of both worlds is to have *controllable* dynamic scope. For example, Common Lisp also has lexical scope everywhere by default, but some variables can be declared as *special*, which means that they are dynamically scoped. The main problem with that is that you can't tell when a variable is special by just looking at the code that uses it, so a more popular approach is the one that is used in Racket: all bindings are always lexically scoped, but there are *parameters* which are a kind of dynamically scoped value containers --- but they are bound to plain (lexically scoped) identifiers. Here's the same code as above, translated to Racket with parameters: #lang racket (define tax% (make-parameter 6.5)) ; create the dynamic container (define (with-tax n) (+ n (* n (/ (tax%) 100)))) ; note how its value is accessed (with-tax 10) ; how much do we pay? (parameterize ([tax% 18.0]) (with-tax 10)) ; not a `let' ;; make that into a function (define il-tax% (make-parameter 18.0)) (define (us-over-il-saving n) (- (parameterize ([tax% (il-tax%)]) (with-tax n)) (with-tax n))) (us-over-il-saving 10) (parameterize ([il-tax% (- (il-tax%) 1)]) (us-over-il-saving 10)) The main point here is that the points where a dynamically scoped value is used are under the programmer's control --- you cannot "customize" what `-` is doing, for example. This gives us back the guarantees that we like to have (= that code works), but of course these points are pre-determined, unlike an environment where everything can be customized including things that are unexpectedly useful. > As a side-note, after many decades of debating this, Emacs has finally > added lexical scope in its core language, but this is still determined > by a flag --- a global `lexical-binding` variable. ------------------------------------------------------------------------ # Dynamic versus Lexical Scope [Tuesday, February 4th] And back to the discussion of whether we should use dynamic or lexical scope: * The most important fact is that we want to view programs as executed by the normal substituting evaluator. Our original motivation was to optimize evaluation only --- not to *change* the semantics! It follows that we want the result of this optimization to behave in the same way. All we need is to evaluate: (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") in the original evaluator to get convinced that `7` should be the correct result (note also that the same code, when translated into Racket, evaluates to `7`). (Yet, this is a very important optimization, which without it lots of programs become too slow to be feasible, so you might claim that you're fine with the modified semantics...) * It does not allow using functions as objects, for example, we have seen that we have a functional representation for pairs: (define (kons x y) (lambda (n) (match n ['first x] ['second y] [else (error ...)]))) (define my-pair (kons 1 2)) If this is evaluated in a dynamically-scoped language, we do get a function as a result, but the values bound to `x` and `y` are now gone! Using the substitution model we substituted these values in, but now they were only held in a cache which no has no entries for them... In the same way, currying would not work, our nice `deriv` function would not work etc etc etc. * Makes reasoning impossible, because any piece of code behaves in a way that *cannot* be predicted until run-time. For example, if dynamic scoping was used in Racket, then you wouldn't be able to know what this function is doing: (define (foo) x) As it is, it will cause a run-time error, but if you call it like this: (let ([x 1]) (foo)) then it will return `1`, and if you later do this: (define (bar x) (foo)) (let ([x 1]) (bar 2)) then you would get `2`! These problems can be demonstrated in Emacs Lisp too, but Racket goes one step further --- it uses the same rule for evaluating a function as well as its values (Lisp uses a different name-space for functions). Because of this, you cannot even rely on the following function: (define (add x y) (+ x y)) to always add `x` and `y`! --- A similar example to the above: (let ([+ -]) (add 1 2)) would return `-1`! * Many so-called "scripting" languages begin their lives with dynamic scoping. The main reason, as we've seen, is that implementing it is extremely simple (no, *nobody* does substitution in the real world! (Well, *almost* nobody...)). Another reason is that these problems make life impossible if you want to use functions as object like you do in Racket, so you notice them very fast --- but in a `normal` language without first-class functions, problems are not as obvious. * For example, bash has `local` variables, but they have dynamic scope: x="the global x" print_x() { echo "The current value of x is \"$x\""; } foo() { local x="x from foo"; print_x; } print_x; foo; print_x Perl began its life with dynamic scope for variables that are declared `local`: $x="the global x"; sub print_x { print "The current value of x is \"$x\"\n"; } sub foo { local($x); $x="x from foo"; print_x; } print_x; foo; print_x; When faced with this problem, "the Perl way" was, obviously, not to remove or fix features, but to pile them up --- so `local` *still* behaves in this way, and now there is a `my` declaration which achieves proper lexical scope (and every serious Perl programmer knows that you should always use `my`)... There are other examples of languages that changed, and languages that want to change (e.g, nobody likes dynamic scope in Emacs Lisp, but there's just too much code now). * This is still a tricky issue, like any other issue with bindings. For example, googling got me quickly to [a Python blog post] which is confused about what "dynamic scoping" is... It claims that Python uses dynamic scope (Search for "Python uses dynamic as opposed to lexical scoping"), yet python always used lexical scope rules, as can be seen by translating their code to Racket (ignore side-effects in this computation): (define (orange-juice) (* x 2)) (define x 3) (define y (orange-juice)) ; y is now 6 (define x 1) (define y (orange-juice)) ; y is now 2 or by trying this in Python: def orange_juice(): return x*2 def foo(x): return orange_juice() foo(2) The real problem of python (pre 2.1, and pre 2.2 without the funny from __future__ import nested_scope line) is that it didn't create closures, which we will talk about shortly. [a Python blog post]: https://folk.idi.ntnu.no/mlh/hetland_org/writing/instant-python.html * Another example, which is an indicator of how easy it is to mess up your scope is the following Ruby bug --- running in `irb`: % irb irb(main):001:0> x = 0 => 0 irb(main):002:0> lambda{|x| x}.call(5) => 5 irb(main):003:0> x => 5 (This is a bug due to weird scoping rules for variables, which was fixed in newer versions of Ruby. See [this Ruby rant] for details, or read about [Ruby and the principle of unwelcome surprise] for additional gems.) [this Ruby rant]: http://innig.net/software/ruby/closures-in-ruby [Ruby and the principle of unwelcome surprise]: http://ceaude.twoticketsplease.de/articles/ruby-and-the-principle-of-unwelcome-surprise.html * Another thing to consider is the fact that compilation is something that you do based only on the lexical structure of programs, since compilers never actually run code. This means that dynamic scope makes compilation close to impossible. * There are some advantages for dynamic scope too. Two notable ones are: - Dynamic scope makes it easy to have a "configuration variable" easily change for the extent of a calling piece of code (this is used extensively in Emacs, for example). The thing is that usually we want to control which variables are "configurable" in this way, statically scoped languages like Racket often choose a separate facility for these. To rephrase the problem of dynamic scoping, it's that *all* variables are modifiable. The same can be said about functions: it is sometimes desirable to change a function dynamically (for example, see "Aspect Oriented Programming"), but if there is no control and all functions can change, we get a world where no code can every be reliable. - It makes recursion immediately available --- for example, {with {f {fun {x} {call f x}}} {call f 0}} is an infinite loop with a dynamically scoped language. But in a lexically scoped language we will need to do some more work to get recursion going. ------------------------------------------------------------------------ # Implementing Lexical Scope: Closures and Environments [Tuesday, February 4th] So how do we fix this? Lets go back to the root of the problem: the new evaluator does not behave in the same way as the substituting evaluator. In the old evaluator, it was easy to see how functions can behave as objects that remember values. For example, when we do this: {with {x 1} {fun {y} {+ x y}}} the result was a function value, which actually was the syntax object for this: {fun {y} {+ 1 y}} Now if we call this function from someplace else like: {with {f {with {x 1} {fun {y} {+ x y}}}} {with {x 2} {call f 3}}} it is clear what the result will be: f is bound to a function that adds 1 to its input, so in the above the later binding for `x` has no effect at all. But with the caching evaluator, the value of {with {x 1} {fun {y} {+ x y}}} is simply: {fun {y} {+ x y}} and there is no place where we save the 1 --- *that's* the root of our problem. (That's also what makes people suspect that using `lambda` in Racket and any other functional language involves some inefficient code-recompiling magic.) In fact, we can verify that by inspecting the returned value, and see that it does contain a free identifier. Clearly, we need to create an object that contains the body and the argument list, like the function syntax object --- but we don't do any substitution, so in addition to the body an argument name(s) we need to remember that we still need to substitute `x` by `1` . This means that the pieces of information we need to know are: - formal argument(s): {y} - body: {+ x y} - pending substitutions: [1/x] and that last bit has the missing `1`. The resulting object is called a `closure` because it closes the function body over the substitutions that are still pending (its environment). So, the first change is in the value of functions which now need all these pieces, unlike the `Fun` case for the syntax object. A second place that needs changing is the when functions are called. When we're done evaluating the `call` arguments (the function value and the argument value) but before we apply the function we have two *values* --- there is no more use for the current substitution cache at this point: we have finished dealing with all substitutions that were necessary over the current expression --- we now continue with evaluating the body of the function, with the new substitutions for the formal arguments and actual values given. But the body itself is the same one we had before --- which is the previous body with its suspended substitutions that we *still* did not do. Rewrite the evaluation rules --- all are the same except for evaluating a `fun` form and a `call` form: eval(N,sc) = N eval({+ E1 E2},sc) = eval(E1,sc) + eval(E2,sc) eval({- E1 E2},sc) = eval(E1,sc) - eval(E2,sc) eval({* E1 E2},sc) = eval(E1,sc) * eval(E2,sc) eval({/ E1 E2},sc) = eval(E1,sc) / eval(E2,sc) eval(x,sc) = lookup(x,sc) eval({with {x E1} E2},sc) = eval(E2,extend(x,eval(E1,sc),sc)) eval({fun {x} E},sc) = <{fun {x} E}, sc> eval({call E1 E2},sc1) = eval(Ef,extend(x,eval(E2,sc1),sc2)) if eval(E1,sc1) = <{fun {x} Ef}, sc2> = error! otherwise As a side note, these substitution caches are a little more than "just a cache" now --- they actually hold an *environment* of substitutions in which expression should be evaluated. So we will switch to the common *environment* name now.: eval(N,env) = N eval({+ E1 E2},env) = eval(E1,env) + eval(E2,env) eval({- E1 E2},env) = eval(E1,env) - eval(E2,env) eval({* E1 E2},env) = eval(E1,env) * eval(E2,env) eval({/ E1 E2},env) = eval(E1,env) / eval(E2,env) eval(x,env) = lookup(x,env) eval({with {x E1} E2},env) = eval(E2,extend(x,eval(E1,env),env)) eval({fun {x} E},env) = <{fun {x} E}, env> eval({call E1 E2},env1) = eval(Ef,extend(x,eval(E2, env1),env2)) if eval(E1,env1) = <{fun {x} Ef}, env2> = error! otherwise In case you find this easier to follow, the "flat algorithm" for evaluating a `call` is: 1. f := evaluate E1 in env1 2. if f is not a <{fun ...},...> closure then error! 3. x := evaluate E2 in env1 4. new_env := extend env_of(f) by mapping arg_of(f) to x 5. evaluate (and return) body_of(f) in new_env Note how the scoping rules that are implied by this definition match the scoping rules that were implied by the substitution-based rules. (It should be possible to prove that they are the same.) The changes to the code are almost trivial, except that we need a way to represent `<{fun {x} Ef}, env>` pairs. ------------------------------------------------------------------------ The implication of this change is that we now cannot use the same type for function syntax and function values since function values have more than just syntax. There is a simple solution to this --- we never do any substitutions now, so we don't need to translate values into expressions --- we can come up with a new type for values, separate from the type of abstract syntax trees. When we do this, we will also fix our hack of using FLANG as the type of values: this was merely a convenience since the AST type had cases for all kinds of values that we needed. (In fact, you should have noticed that Racket does this too: numbers, strings, booleans, etc are all used by both programs and syntax representation (s-expressions) --- but note that function values are *not* used in syntax.) We will now implement a separate `VAL` type for runtime values. First, we need now a type for such environments --- we can use `Listof` for this: ;; a type for environments: (define-type ENV = (Listof (List Symbol VAL))) but we can just as well define a new type for environment values: (define-type ENV [EmptyEnv] [Extend Symbol VAL ENV]) Reimplementing `lookup` is now simple: (: lookup : Symbol ENV -> VAL) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (cases env [(EmptyEnv) (error 'lookup "no binding for ~s" name)] [(Extend id val rest-env) (if (eq? id name) val (lookup name rest-env))])) ... we don't need `extend` because we get `Extend` from the type definition, and we also get `(EmptyEnv)` instead of `empty-subst`. We now use this with the new type for values --- two variants of these: (define-type VAL [NumV Number] [FunV Symbol FLANG ENV]) ; arg-name, body, scope And now the new implementation of `eval` which uses the new type and implements lexical scope: (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) (NumV n)] [(Add l r) (arith-op + (eval l env) (eval r env))] [(Sub l r) (arith-op - (eval l env) (eval r env))] [(Mul l r) (arith-op * (eval l env) (eval r env))] [(Div l r) (arith-op / (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (FunV bound-id bound-body env)] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV bound-id bound-body f-env) (eval bound-body (Extend bound-id (eval arg-expr env) f-env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) We also need to update `arith-op` to use `VAL` objects. The full code follows --- it now passes all tests, including the example that we used to find the problem. ;;; ---<<>>---------------------------------------------- ;; The Flang interpreter, using environments #lang pl #| The grammar: ::= | { + } | { - } | { * } | { / } | { with { } } | | { fun { } } | { call } Evaluation rules: eval(N,env) = N eval({+ E1 E2},env) = eval(E1,env) + eval(E2,env) eval({- E1 E2},env) = eval(E1,env) - eval(E2,env) eval({* E1 E2},env) = eval(E1,env) * eval(E2,env) eval({/ E1 E2},env) = eval(E1,env) / eval(E2,env) eval(x,env) = lookup(x,env) eval({with {x E1} E2},env) = eval(E2,extend(x,eval(E1,env),env)) eval({fun {x} E},env) = <{fun {x} E}, env> eval({call E1 E2},env1) = eval(Ef,extend(x,eval(E2,env1),env2)) if eval(E1,env1) = <{fun {x} Ef}, env2> = error! otherwise |# (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; Types for environments, values, and a lookup function (define-type ENV [EmptyEnv] [Extend Symbol VAL ENV]) (define-type VAL [NumV Number] [FunV Symbol FLANG ENV]) (: lookup : Symbol ENV -> VAL) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (cases env [(EmptyEnv) (error 'lookup "no binding for ~s" name)] [(Extend id val rest-env) (if (eq? id name) val (lookup name rest-env))])) (: NumV->number : VAL -> Number) ;; convert a FLANG runtime numeric value to a Racket one (define (NumV->number val) (cases val [(NumV n) n] [else (error 'arith-op "expected a number, got: ~s" val)])) (: arith-op : (Number Number -> Number) VAL VAL -> VAL) ;; gets a Racket numeric binary operator, and uses it within a NumV ;; wrapper (define (arith-op op val1 val2) (NumV (op (NumV->number val1) (NumV->number val2)))) (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) (NumV n)] [(Add l r) (arith-op + (eval l env) (eval r env))] [(Sub l r) (arith-op - (eval l env) (eval r env))] [(Mul l r) (arith-op * (eval l env) (eval r env))] [(Div l r) (arith-op / (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (FunV bound-id bound-body env)] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV bound-id bound-body f-env) (eval bound-body (Extend bound-id (eval arg-expr env) f-env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) (EmptyEnv))]) (cases result [(NumV n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => 7) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Fixing an Overlooked Bug [Tuesday, February 4th] Incidentally, this version fixes a bug we had previously in the substitution version of FLANG: (run "{with {f {fun {y} {+ x y}}} {with {x 7} {call f 1}}}") This bug was due to our naive `subst`, which doesn't avoid capturing renames. But note that since that version of the evaluator makes its way from the outside in, there is no difference in semantics for *valid* programs --- ones that don't have free identifiers. (Reminder: This was *not* a dynamically scoped language, just a bug that happened when `x` wasn't substituted away before `f` was replaced with something that refers to `x`.) ------------------------------------------------------------------------ # Lexical Scope using Racket Closures [Tuesday, February 4th] > [PLAI §11] (without the last part about recursion) An alternative representation for an environment. We've already seen how first-class functions can be used to implement "objects" that contain some information. We can use the same idea to represent an environment. The basic intuition is --- an environment is a *mapping* (a function) between an identifier and some value. For example, we can represent the environment that maps `'a` to `1` and `'b` to `2` (using just numbers for simplicity) using this function: (: my-map : Symbol -> Number) (define (my-map id) (cond [(eq? 'a id) 1] [(eq? 'b id) 2] [else (error ...)])) An empty mapping that is implemented in this way has the same type: (: empty-mapping : Symbol -> Number) (define (empty-mapping id) (error ...)) We can use this idea to implement our environments: we only need to define three things --- `EmptyEnv`, `Extend`, and `lookup`. If we manage to keep the contract to these functions intact, we will be able to simply plug it into the same evaluator code with no other changes. It will also be more convenient to define `ENV` as the appropriate function type for use in the `VAL` type definition instead of using the actual type: ;; Define a type for functional environments (define-type ENV = Symbol -> VAL) Now we get to `EmptyEnv` --- this is expected to be a function that expects no arguments and creates an empty environment, one that behaves like the `empty-mapping` function defined above. We could define it like this (changing the `empty-mapping` type to return a `VAL`): (define (EmptyEnv) empty-mapping) but we can skip the need for an extra definition and simply return an empty mapping function: (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error ...))) (The un-Rackety name is to avoid replacing previous code that used the `EmptyEnv` name for the constructor that was created by the type definition.) The next thing we tackle is `lookup`. The previous definition that was used is: (: lookup : Symbol ENV -> VAL) (define (lookup name env) (cases env [(EmptyEnv) (error 'lookup "no binding for ~s" name)] [(Extend id val rest-env) (if (eq? id name) val (lookup name rest-env))])) How should it be modified now? Easy --- an environment is a mapping: a Racket function that will do the searching job itself. We don't need to modify the contract since we're still using `ENV`, except a different implementation for it. The new definition is: (: lookup : Symbol ENV -> VAL) (define (lookup name env) (env name)) Note that `lookup` does almost nothing --- it simply delegates the real work to the `env` argument. This is a good hint for the error message that empty mappings should throw --- (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error 'lookup "no binding for ~s" id))) Finally, `Extend` --- this was previously created by the variant case of the ENV type definition: [Extend Symbol VAL ENV] keeping the same type that is implied by this variant means that the new `Extend` should look like this: (: Extend : Symbol VAL ENV -> ENV) (define (Extend id val rest-env) ...) The question is --- how do we extend a given environment? Well, first, we know that the result should be mapping --- a `symbol -> VAL` function that expects an identifier to look for: (: Extend : Symbol VAL ENV -> ENV) (define (Extend id val rest-env) (lambda (name) ...)) Next, we know that in the generated mapping, if we look for `id` then the result should be `val`: (: Extend : Symbol VAL ENV -> ENV) (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val ...))) If the `name` that we're looking for is not the same as `id`, then we need to search through the previous environment, eg: (lookup name rest). But we know what `lookup` does --- it simply delegates back to the mapping function (which is our `rest` argument), so we can take a direct route: (: Extend : Symbol VAL ENV -> ENV) (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val (rest-env name)))) (Note that the last line is simply `(lookup name rest-env)`, but we know that we have a functional implementation.) To see how all this works, try out extending an empty environment a few times and examine the result. For example, the environment that we began with: (define (my-map id) (cond [(eq? 'a id) 1] [(eq? 'b id) 2] [else (error ...)])) behaves in the same way (if the type of values is numbers) as (Extend 'a 1 (Extend 'b 2 (EmptyEnv))) The new code is now the same, except for the environment code: #lang pl #| The grammar: ::= | { + } | { - } | { * } | { / } | { with { } } | | { fun { } } | { call } Evaluation rules: eval(N,env) = N eval({+ E1 E2},env) = eval(E1,env) + eval(E2,env) eval({- E1 E2},env) = eval(E1,env) - eval(E2,env) eval({* E1 E2},env) = eval(E1,env) * eval(E2,env) eval({/ E1 E2},env) = eval(E1,env) / eval(E2,env) eval(x,env) = lookup(x,env) eval({with {x E1} E2},env) = eval(E2,extend(x,eval(E1,env),env)) eval({fun {x} E},env) = <{fun {x} E}, env> eval({call E1 E2},env1) = eval(Ef,extend(x,eval(E2,env1),env2)) if eval(E1,env1) = <{fun {x} Ef}, env2> = error! otherwise |# (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; Types for environments, values, and a lookup function (define-type VAL [NumV Number] [FunV Symbol FLANG ENV]) ;; Define a type for functional environments (define-type ENV = Symbol -> VAL) (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error 'lookup "no binding for ~s" id))) (: Extend : Symbol VAL ENV -> ENV) ;; extend a given environment cache with a new binding (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val (rest-env name)))) (: lookup : Symbol ENV -> VAL) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (env name)) (: NumV->number : VAL -> Number) ;; convert a FLANG runtime numeric value to a Racket one (define (NumV->number val) (cases val [(NumV n) n] [else (error 'arith-op "expected a number, got: ~s" val)])) (: arith-op : (Number Number -> Number) VAL VAL -> VAL) ;; gets a Racket numeric binary operator, and uses it within a NumV ;; wrapper (define (arith-op op val1 val2) (NumV (op (NumV->number val1) (NumV->number val2)))) (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) (NumV n)] [(Add l r) (arith-op + (eval l env) (eval r env))] [(Sub l r) (arith-op - (eval l env) (eval r env))] [(Mul l r) (arith-op * (eval l env) (eval r env))] [(Div l r) (arith-op / (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (FunV bound-id bound-body env)] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV bound-id bound-body f-env) (eval bound-body (Extend bound-id (eval arg-expr env) f-env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) (EmptyEnv))]) (cases result [(NumV n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => 7) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # More Closures (on both levels) [Tuesday, February 4th] Racket closures (= functions) can be used in other places too, and as we have seen, they can do more than encapsulate various values --- they can also hold the behavior that is expected of these values. To demonstrate this we will deal with closures in our language. We currently use a variant that holds the three pieces of relevant information: [FunV Symbol FLANG ENV] We can replace this by a functional object, which will hold the three values. First, change the `VAL` type to hold functions for `FunV` values: (define-type VAL [NumV Number] [FunV (? -> ?)]) And note that the function should somehow encapsulate the same information that was there previously, the question is *how* this information is going to be done, and this will determine the actual type. This information plays a role in two places in our evaluator --- generating a closure in the `Fun` case, and using it in the `Call` case: [(Fun bound-id bound-body) (FunV bound-id bound-body env)] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV bound-id bound-body f-env) (eval bound-body ;*** (Extend bound-id ;*** (eval arg-expr env) ;*** f-env))] ;*** [else (error 'eval "`call' expects a function, got: ~s" fval)]))] we can simply fold the marked functionality bit of `Call` into a Racket function that will be stored in a `FunV` object --- this piece of functionality takes an argument value, extends the closure's environment with its value and the function's name, and continues to evaluate the function body. Folding all of this into a function gives us: (lambda (arg-val) (eval bound-body (Extend bound-id arg-val env))) where the values of `bound-body`, `bound-id`, and `val` are known at the time that the `FunV` is *constructed*. Doing this gives us the following code for the two cases: [(Fun bound-id bound-body) (FunV (lambda (arg-val) (eval bound-body (Extend bound-id arg-val env))))] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV proc) (proc (eval arg-expr env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))] And now the type of the function is clear: (define-type VAL [NumV Number] [FunV (VAL -> VAL)]) And again, the rest of the code is unmodified: #lang pl (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; Types for environments, values, and a lookup function (define-type VAL [NumV Number] [FunV (VAL -> VAL)]) ;; Define a type for functional environments (define-type ENV = Symbol -> VAL) (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error 'lookup "no binding for ~s" id))) (: Extend : Symbol VAL ENV -> ENV) ;; extend a given environment cache with a new binding (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val (rest-env name)))) (: lookup : Symbol ENV -> VAL) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (env name)) (: NumV->number : VAL -> Number) ;; convert a FLANG runtime numeric value to a Racket one (define (NumV->number val) (cases val [(NumV n) n] [else (error 'arith-op "expected a number, got: ~s" val)])) (: arith-op : (Number Number -> Number) VAL VAL -> VAL) ;; gets a Racket numeric binary operator, and uses it within a NumV ;; wrapper (define (arith-op op val1 val2) (NumV (op (NumV->number val1) (NumV->number val2)))) (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) (NumV n)] [(Add l r) (arith-op + (eval l env) (eval r env))] [(Sub l r) (arith-op - (eval l env) (eval r env))] [(Mul l r) (arith-op * (eval l env) (eval r env))] [(Div l r) (arith-op / (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (FunV (lambda (arg-val) (eval bound-body (Extend bound-id arg-val env))))] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV proc) (proc (eval arg-expr env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) (EmptyEnv))]) (cases result [(NumV n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => 7) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Types of Evaluators [Tuesday, February 4th] What we did just now is implement lexical environments and closures in the language we implement using lexical environments and closures in our own language (Racket)! This is another example of embedding a feature of the host language in the implemented language, an issue that we have already discussed. There are many examples of this, even when the two languages involved are different. For example, if we have this bit in the C implementation of Racket: // Disclaimer: not real Racket code Racket_Object *eval_and(int argc, Racket_Object *argv[]) { Racket_Object *tmp; if ( argc != 2 ) signal_racket_error("bad number of arguments"); else if ( racket_eval(argv[0]) != racket_false && (tmp = racket_eval(argv[1])) != racket_false ) return tmp; else return racket_false; } then the special semantics of evaluating a Racket `and` form is being inherited from C's special treatment of `&&`. You can see this by the fact that if there is a bug in the C compiler, then it will propagate to the resulting Racket implementation too. A different solution is to not use `&&` at all: // Disclaimer: not real Racket code Racket_Object *eval_and(int argc, Racket_Object *argv[]) { Racket_Object *tmp; if ( argc != 2 ) signal_racket_error("bad number of arguments"); else if ( racket_eval(argv[0]) != racket_false ) return racket_eval(argv[1]); else return racket_false; } and we can say that this is even better since it evaluates the second expression in tail position. But in this case we don't really get that benefit, since C itself is not doing tail-call optimization as a standard feature (though some compilers do so under some circumstances). We have seen a few different implementations of evaluators that are quite different in flavor. They suggest the following taxonomy. * A ___syntactic evaluator___ is one that uses its own language to represent expressions and semantic runtime values of the evaluated language, implementing all the corresponding behavior explicitly. * A ___meta evaluator___ is an evaluator that uses language features of its own language to directly implement behavior of the evaluated language. While our substitution-based FLANG evaluator was close to being a syntactic evaluator, we haven't written any purely syntactic evaluators so far: we still relied on things like Racket arithmetics etc. The most recent evaluator that we have studied, is even more of a *meta* evaluator than the preceding ones: it doesn't even implement closures and lexical scope, and instead, it uses the fact that Racket itself has them. With a good match between the evaluated language and the implementation language, writing a meta evaluator can be very easy. With a bad match, though, it can be very hard. With a syntactic evaluator, implementing each semantic feature will be somewhat hard, but in return you don't have to worry as much about how well the implementation and the evaluated languages match up. In particular, if there is a particularly strong mismatch between the implementation and the evaluated language, it may take less effort to write a syntactic evaluator than a meta evaluator. As an exercise, we can build upon our latest evaluator to remove the encapsulation of the evaluator's response in the VAL type. The resulting evaluator is shown below. This is a true meta evaluator: it uses Racket closures to implement FLANG closures, Racket function application for FLANG function application, Racket numbers for FLANG numbers, and Racket arithmetic for FLANG arithmetic. In fact, ignoring some small syntactic differences between Racket and FLANG, this latest evaluator can be classified as something more specific than a meta evaluator: * A ___meta-circular evaluator___ is a meta evaluator in which the implementation and the evaluated languages are the same. (Put differently, the trivial nature of the evaluator clues us in to the deep connection between the two languages, whatever their syntactic differences may be.) ------------------------------------------------------------------------ # Feature Embedding [Tuesday, February 4th] We saw that the difference between lazy evaluation and eager evaluation is in the evaluation rules for `with` forms, function applications, etc: eval({with {x E1} E2}) = eval(E2[eval(E1)/x]) is eager, and eval({with {x E1} E2}) = eval(E2[E1/x]) is lazy. But is the first rule *really* eager? The fact is that the only thing that makes it eager is the fact that our understanding of the mathematical notation is eager --- if we were to take math as lazy, then the description of the rule becomes a description of lazy evaluation. Another way to look at this is --- take the piece of code that implements this evaluation: (: eval : FLANG -> Number) ;; evaluates FLANG expressions by reducing them to numbers (define (eval expr) (cases expr ... [(With bound-id named-expr bound-body) (eval (subst bound-body bound-id (Num (eval named-expr))))] ...)) and the same question applies: is this really implementing eager evaluation? We know that this is indeed eager --- we can simply try it and check that it is, but it is only eager because we are using an eager language for the implementation! If our own language was lazy, then the evaluator's implementation would run lazily, which means that the above applications of the `eval` and the `subst` functions would also be lazy, making our evaluator lazy as well. This is a general phenomena where some of the semantic features of the language we use (math in the formal description, Racket in our code) gets *embedded* into the language we implement. Here's another example --- consider the code that implements arithmetics: (: eval : FLANG -> Number) ;; evaluates FLANG expressions by reducing them to numbers (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] ...)) what if it was written like this: FLANG eval(FLANG expr) { if (is_Num(expr)) return num_of_Num(expr); else if (is_Add(expr)) return eval(lhs_of_Add(expr)) + eval(rhs_of_Add(expr)); else if ... ... } Would it still implement unlimited integers and exact fractions? That depends on the language that was used to implement it: the above syntax suggests C, C++, Java, or some other relative, which usually come with limited integers and no exact fractions. But this really depends on the language --- even our own code has unlimited integers and exact rationals only because Racket has them. If we were using a language that didn't have such features (there are such Scheme implementations), then our implemented language would absorb these (lack of) features too, and its own numbers would be limited in just the same way. (And this includes the syntax for numbers, which we embedded intentionally, like the syntax for identifiers). The bottom line is that we should be aware of such issues, and be very careful when we talk about semantics. Even the language that we use to communicate (semi-formal logic) can mean different things. ------------------------------------------------------------------------ Aside: read "Reflections on Trusting Trust" by Ken Thompson (You can skip to the "Stage II" part to get to the interesting stuff.) (And when you're done, look for "XcodeGhost" to see a relevant example, and don't miss the leaked document on the wikipedia page...) ------------------------------------------------------------------------ Here is yet another variation of our evaluator that is even closer to a meta-circular evaluator. It uses Racket values directly to implement values, so arithmetic operations become straightforward. Note especially how the case for function application is similar to arithmetics: a FLANG function application translates to a Racket function application. In both cases (applications and arithmetics) we don't even check the objects since they are simple Racket objects --- if our language happens to have some meaning for arithmetics with functions, or for applying numbers, then we will inherit the same semantics in our language. This means that we now specify less behavior and fall back more often on what Racket does. We use Racket values with this type definition: (define-type VAL = (U Number (VAL -> VAL))) And the evaluation function can now be: (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) n] ;*** return the actual number [(Add l r) (+ (eval l env) (eval r env))] [(Sub l r) (- (eval l env) (eval r env))] [(Mul l r) (* (eval l env) (eval r env))] [(Div l r) (/ (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (lambda ([arg-val : VAL]) ;*** return the racket function ;; note that this requires input type specifications since ;; typed racket can't guess the right one (eval bound-body (Extend bound-id arg-val env)))] [(Call fun-expr arg-expr) ((eval fun-expr env) ;*** trivial like the arithmetics! (eval arg-expr env))])) Note how the arithmetics implementation is simple --- it's a direct translation of the FLANG syntax to Racket operations, and since we don't check the inputs to the Racket operations, we let Racket throw type errors for us. Note also how function application is just like the arithmetic operations: a FLANG application is directly translated to a Racket application. However, this does not work quite as simply in Typed Racket. The whole point of typechecking is that we never run into type errors --- so we cannot throw back on Racket errors since code that might produce them is forbidden! A way around this is to perform explicit checks that guarantee that Racket cannot run into type errors. We do this with the following two helpers that are defined inside `eval`: (: evalN : FLANG -> Number) (define (evalN e) (let ([n (eval e env)]) (if (number? n) n (error 'eval "got a non-number: ~s" n)))) (: evalF : FLANG -> (VAL -> VAL)) (define (evalF e) (let ([f (eval e env)]) (if (function? f) f (error 'eval "got a non-function: ~s" f)))) Note that Typed Racket is "smart enough" to figure out that in `evalF` the result of the recursive evaluation has to be either `Number` or `(VAL -> VAL)`; and since the `if` throws out on numbers, we're left with `(VAL -> VAL)` functions, not just any function. #lang pl (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; Types for environments, values, and a lookup function ;; Values are plain Racket values, no new VAL wrapper; ;; (but note that this is a recursive definition) (define-type VAL = (U Number (VAL -> VAL))) ;; Define a type for functional environments (define-type ENV = (Symbol -> VAL)) (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error 'lookup "no binding for ~s" id))) (: Extend : Symbol VAL ENV -> ENV) ;; extend a given environment cache with a new binding (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val (rest-env name)))) (: lookup : Symbol ENV -> VAL) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (env name)) (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (: evalN : FLANG -> Number) (define (evalN e) (let ([n (eval e env)]) (if (number? n) n (error 'eval "got a non-number: ~s" n)))) (: evalF : FLANG -> (VAL -> VAL)) (define (evalF e) (let ([f (eval e env)]) (if (function? f) f (error 'eval "got a non-function: ~s" f)))) (cases expr [(Num n) n] [(Add l r) (+ (evalN l) (evalN r))] [(Sub l r) (- (evalN l) (evalN r))] [(Mul l r) (* (evalN l) (evalN r))] [(Div l r) (/ (evalN l) (evalN r))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (lambda ([arg-val : VAL]) (eval bound-body (Extend bound-id arg-val env)))] [(Call fun-expr arg-expr) ((evalF fun-expr) (eval arg-expr env))])) (: run : String -> VAL) ; no need to convert VALs to numbers ;; evaluate a FLANG program contained in a string (define (run str) (eval (parse str) (EmptyEnv))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => 7) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Recursion, Recursion, Recursion [Tuesday, February 11th] > [PLAI §9] There is one major feature that is still missing from our language: we have no way to perform recursion (therefore no kind of loops). So far, we could only use recursion when we had *names*. In FLANG, the only way we can have names is through `with` which not good enough for recursion. To discuss the issue of recursion, we switch to a "broken" version of (untyped) Racket --- one where a `define` has a different scoping rules: the scope of the defined name does *not* cover the defined expression. Specifically, in this language, this doesn't work: #lang pl broken (define (fact n) (if (zero? n) 1 (* n (fact (- n 1))))) (fact 5) In our language, this translation would also not work (assuming we have `if` etc): {with {fact {fun {n} {if {= n 0} 1 {* n {call fact {- n 1}}}}}} {call fact 5}} And similarly, in plain Racket this won't work if `let` is the only tool you use to create bindings: (let ([fact (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))]) (fact 5)) In the broken-scope language, the `define` form is more similar to a mathematical definition. For example, when we write: (define (F x) x) (define (G y) (F y)) (G F) it is actually shorthand for (define F (lambda (x) x)) (define G (lambda (y) (F y))) (G F) we can then replace defined names with their definitions: (define F (lambda (x) x)) (define G (lambda (y) (F y))) ((lambda (y) (F y)) (lambda (x) x)) and this can go on, until we get to the actual code that we wrote: ((lambda (y) ((lambda (x) x) y)) (lambda (x) x)) This means that the above `fact` definition is similar to writing: fact := (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))) (fact 5) which is not a well-formed definition --- it is *meaningless* (this is a formal use of the word "meaningless"). What we'd really want, is to take the *equation* (using `=` instead of `:=`) fact = (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))) and find a solution which will be a value for `fact` that makes this true. If you look at the Racket evaluation rules handout on the web page, you will see that this problem is related to the way that we introduced the Racket `define`: there is a hand-wavy explanation that talks about *knowing* things. The big question is: can we define recursive functions without Racket's magical `define` form? > Note: This question is a little different than the question of > implementing recursion in our language --- in the Racket case we have > no control over the implementation of the language. As it will > eventually turn out, implementing recursion in our own language will > be quite easy when we use mutation in a specific way. So the question > that we're now facing can be phrased as either "can we get recursion > in Racket without Racket's magical definition forms?" or "can we get > recursion in our interpreter without mutation?". ------------------------------------------------------------------------ # Recursion without the Magic [Tuesday, February 11th] > [PLAI §22.4] (we go much deeper) > Note: This explanation is similar to the one you can find in "The Why > of Y", by Richard Gabriel. To implement recursion without the `define` magic, we first make an observation: this problem does *not* come up in a dynamically-scoped language. Consider the `let`-version of the problem: #lang pl dynamic (let ([fact (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))]) (fact 5)) This works fine --- because by the time we get to evaluate the body of the function, `fact` is already bound to itself in the current dynamic scope. (This is another reason why dynamic scope is perceived as a convenient approach in new languages.) Regardless, the problem that we have with lexical scope is still there, but the way things work in a dynamic scope suggest a solution that we can use now. Just like in the dynamic scope case, when `fact` is called, it does have a value --- the only problem is that this value is inaccessible in the lexical scope of its body. Instead of trying to get the value in via lexical scope, we can imitate what happens in the dynamically scoped language by passing the `fact` value to itself so it can call itself (going back to the original code in the broken-scope language): (define (fact self n) ;*** (if (zero? n) 1 (* n (self (- n 1))))) (fact fact 5) ;*** except that now the recursive call should still send itself along: (define (fact self n) (if (zero? n) 1 (* n (self self (- n 1))))) ;*** (fact fact 5) The problem is that this required rewriting calls to `fact` --- both outside and recursive calls inside. To make this an acceptable solution, calls from both places should not change. Eventually, we should be able to get a working `fact` definition that uses just (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))) The first step in resolving this problem is to curry the `fact` definition. (define (fact self) ;*** (lambda (n) ;*** (if (zero? n) 1 (* n ((self self) (- n 1)))))) ;*** ((fact fact) 5) ;*** Now `fact` is no longer our factorial function --- it's a function that constructs it. So call it `make-fact`, and bind `fact` to the actual factorial function. (define (make-fact self) ;*** (lambda (n) (if (zero? n) 1 (* n ((self self) (- n 1)))))) (define fact (make-fact make-fact)) ;*** (fact 5) ;*** We can try to do the same thing in the body of the factorial function: instead of calling `(self self)`, just bind `fact` to it: (define (make-fact self) (lambda (n) (let ([fact (self self)]) ;*** (if (zero? n) 1 (* n (fact (- n 1))))))) ;*** (define fact (make-fact make-fact)) (fact 5) This works fine, but if we consider our original goal, we need to get that local `fact` binding outside of the `(lambda (n) ...)` --- so we're left with a definition that uses the factorial expression as is. So, swap the two lines: (define (make-fact self) (let ([fact (self self)]) ;*** (lambda (n) ;*** (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact (make-fact make-fact)) (fact 5) But the problem is that this gets us into an infinite loop because we're trying to evaluate `(self self)` too early. In fact, if we ignore the body of the `let` and other details, we basically do this: (define (make-fact self) (self self)) (make-fact make-fact) --reduce-sugar--> (define make-fact (lambda (self) (self self))) (make-fact make-fact) --replace-definition--> ((lambda (self) (self self)) (lambda (self) (self self))) --rename-identifiers--> ((lambda (x) (x x)) (lambda (x) (x x))) And this expression has an interesting property: it reduces to itself, so evaluating it gets stuck in an infinite loop. So how do we solve this? Well, we know that `(self self)` *should* be the same value that is the factorial function itself --- so it must be a one-argument function. If it's such a function, we can use a value that is equivalent, except that it will not get evaluated until it is needed, when the function is called. The trick here is the observation that `(lambda (n) (add1 n))` is really the same function as `add1`, except that the `add1` part doesn't get evaluated until the function is called. Applying this trick to our code produces a version that does not get stuck in the same infinite loop: (define (make-fact self) (let ([fact (lambda (n) ((self self) n))]) ;*** (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact (make-fact make-fact)) (fact 5) Continuing from here --- we know that (let ([x v]) e) is the same as ((lambda (x) e) v) (remember how we derived `fun` from a `with`), so we can turn that `let` into the equivalent function application form: (define (make-fact self) ((lambda (fact) ;*** (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))) (lambda (n) ((self self) n)))) ;*** (define fact (make-fact make-fact)) (fact 5) And note now that the (lambda (fact) ...) expression is everything that we need for a recursive definition of `fact` --- it has the proper factorial body with a plain recursive call. It's almost like the usual value that we'd want to define `fact` as, except that we still have to abstract on the recursive value itself. So lets move this code into a separate definition for `fact-core`: (define fact-core ;*** (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define (make-fact self) (fact-core ;*** (lambda (n) ((self self) n)))) (define fact (make-fact make-fact)) (fact 5) We can now proceed by moving the `(make-fact make-fact)` self application into its own function which is what creates the real factorial: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define (make-fact self) (fact-core (lambda (n) ((self self) n)))) (define (make-real-fact) (make-fact make-fact)) ;*** (define fact (make-real-fact)) ;*** (fact 5) Rewrite the `make-fact` definition using an explicit `lambda`: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define make-fact ;*** (lambda (self) ;*** (fact-core (lambda (n) ((self self) n))))) (define (make-real-fact) (make-fact make-fact)) (define fact (make-real-fact)) (fact 5) and fold the functionality of `make-fact` and `make-real-fact` into a single `make-fact` function by just using the value of `make-fact` explicitly instead of through a definition: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define (make-real-fact) (let ([make (lambda (self) ;*** (fact-core ;*** (lambda (n) ((self self) n))))]) ;*** (make make))) (define fact (make-real-fact)) (fact 5) We can now observe that `make-real-fact` has nothing that is specific to factorial --- we can make it take a "core function" as an argument: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define (make-real-fact core) ;*** (let ([make (lambda (self) (core ;*** (lambda (n) ((self self) n))))]) (make make))) (define fact (make-real-fact fact-core)) ;*** (fact 5) and call it `make-recursive`: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define (make-recursive core) ;*** (let ([make (lambda (self) (core (lambda (n) ((self self) n))))]) (make make))) (define fact (make-recursive fact-core)) ;*** (fact 5) We're almost done now --- there's no real need for a separate `fact-core` definition, just use the value for the definition of `fact`: (define (make-recursive core) (let ([make (lambda (self) (core (lambda (n) ((self self) n))))]) (make make))) (define fact (make-recursive (lambda (fact) ;*** (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))))) ;*** (fact 5) turn the `let` into a function form: (define (make-recursive core) ((lambda (make) (make make)) ;*** (lambda (self) ;*** (core (lambda (n) ((self self) n)))))) ;*** (define fact (make-recursive (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))))) (fact 5) do some renamings to make things simpler --- `make` and `self` turn to `x`, and `core` to `f`: (define (make-recursive f) ;*** ((lambda (x) (x x)) ;*** (lambda (x) (f (lambda (n) ((x x) n)))))) ;*** (define fact (make-recursive (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))))) (fact 5) or we can manually expand that first (lambda (x) (x x)) application to make the symmetry more obvious (not really surprising because it started with a `let` whose purpose was to do a self-application): (define (make-recursive f) ((lambda (x) (f (lambda (n) ((x x) n)))) ;*** (lambda (x) (f (lambda (n) ((x x) n)))))) ;*** (define fact (make-recursive (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))))) (fact 5) And we finally got what we were looking for: a general way to define *any* recursive function without any magical `define` tricks. This also work for other recursive functions: #lang pl broken (define (make-recursive f) ((lambda (x) (f (lambda (n) ((x x) n)))) (lambda (x) (f (lambda (n) ((x x) n)))))) (define fact (make-recursive (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))))) (fact 5) (define fib (make-recursive (lambda (fib) (lambda (n) (if (<= n 1) n (+ (fib (- n 1)) (fib (- n 2)))))))) (fib 8) (define length (make-recursive (lambda (length) (lambda (l) (if (null? l) 0 (+ (length (rest l)) 1)))))) (length '(x y z)) A convenient tool that people often use on paper is to perform a kind of a syntactic abstraction: "assume that whenever I write (twice foo) I really meant to write (foo foo)". This can often be done as plain abstractions (that is, using functions), but in some cases --- for example, if we want to abstract over definitions --- we just want such a rewrite rule. (More on this towards the end of the course.) The broken-scope language does provide such a tool --- `rewrite` extends the language with a rewrite rule. Using this, and our `make-recursive`, we can make up a recursive definition form: (rewrite (define/rec (f x) E) => (define f (make-recursive (lambda (f) (lambda (x) E))))) In other words, we've created our own "magical definition" form. The above code can now be written in almost the same way it is written in plain Racket: #lang pl broken (define (make-recursive f) ((lambda (x) (f (lambda (n) ((x x) n)))) (lambda (x) (f (lambda (n) ((x x) n)))))) (rewrite (define/rec (f x) E) => (define f (make-recursive (lambda (f) (lambda (x) E))))) ;; examples (define/rec (fact n) (if (zero? n) 1 (* n (fact (- n 1))))) (fact 5) (define/rec (fib n) (if (<= n 1) n (+ (fib (- n 1)) (fib (- n 2))))) (fib 8) (define/rec (length l) (if (null? l) 0 (+ (length (rest l)) 1))) (length '(x y z)) Finally, note that make-recursive is limited to 1-argument functions only because of the protection from eager evaluation. In any case, it can be used in any way you want, for example, (make-recursive (lambda (f) (lambda (x) f))) is a function that *returns* itself rather than calling itself. Using the rewrite rule, this would be: (define/rec (f x) f) which is the same as: (define (f x) f) in plain Racket. ------------------------------------------------------------------------ # The Core of `make-recursive` [Tuesday, February 11th] As in Racket, being able to express recursive functions is a fundamental property of the language. It means that we can have loops in our language, and that's the essence of making a language powerful enough to be TM-equivalent --- able to express undecidable problems, where we don't know whether there is an answer or not. The core of what makes this possible is the expression that we have seen in our derivation: ((lambda (x) (x x)) (lambda (x) (x x))) which reduces to itself, and therefore has no value: trying to evaluate it gets stuck in an infinite loop. (This expression is often called "Omega".) This is the key for creating a loop --- we use it to make recursion possible. Looking at our final `make-recursive` definition and ignoring for a moment the "protection" that we need against being stuck prematurely in an infinite loop: (define (make-recursive f) ((lambda (x) (x x)) (lambda (x) (f (x x))))) we can see that this is almost the same as the Omega expression --- the only difference is that application of `f`. Indeed, this expression (the result of (make-recursive F) for some `F`) reduces in a similar way to Omega: ((lambda (x) (x x)) (lambda (x) (F (x x)))) ((lambda (x) (F (x x))) (lambda (x) (F (x x)))) (F ((lambda (x) (F (x x))) (lambda (x) (F (x x))))) (F (F ((lambda (x) (F (x x))) (lambda (x) (F (x x)))))) (F (F (F ((lambda (x) (F (x x))) (lambda (x) (F (x x))))))) ... which means that the actual value of this expression is: (F (F (F ...forever...))) This definition would be sufficient if we had a lazy language, but to get things working in a strict one we need to bring back the protection. This makes things a little different --- if we use `(protect f)` to be a shorthand for the protection trick, (rewrite (protect f) => (lambda (x) (f x))) then we have: (define (make-recursive f) ((lambda (x) (x x)) (lambda (x) (f (protect (x x)))))) which makes the (make-recursive F) evaluation reduce to (F (protect (F (protect (F (protect (...forever...))))))) and this is still the same result (as long as `F` is a single-argument function). (Note that `protect` cannot be implemented as a plain function!) ------------------------------------------------------------------------ # Denotational Explanation of Recursion [Tuesday, February 11th] > Note: This explanation is similar to the one you can find in "The > Little Schemer" called "(Y Y) Works!", by Dan Friedman and Matthias > Felleisen. The explanation that we have now for how to derive the `make-recursive` definition is fine --- after all, we did manage to get it working. But this explanation was done from a kind of an operational point of view: we knew a certain trick that can make things work and we pushed things around until we got it working like we wanted. Instead of doing this, we can re-approach the problem from a more declarative point of view. So, start again from the same broken code that we had (using the broken-scope language): (define fact (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))) This is as broken as it was when we started: the occurrence of `fact` in the body of the function is free, which means that this code is meaningless. To avoid the compilation error that we get when we run this code, we can substitute *anything* for that `fact` --- it's even better to use a replacement that will lead to a runtime error: (define fact (lambda (n) (if (zero? n) 1 (* n (777 (- n 1)))))) ;*** This function will not work in a similar way to the original one --- but there is one case where it *does* work: when the input value is `0` (since then we do not reach the bogus application). We note this by calling this function `fact0`: (define fact0 ;*** (lambda (n) (if (zero? n) 1 (* n (777 (- n 1)))))) Now that we have this function defined, we can use it to write `fact1` which is the factorial function for arguments of `0` or `1`: (define fact0 (lambda (n) (if (zero? n) 1 (* n (777 (- n 1)))))) (define fact1 (lambda (n) (if (zero? n) 1 (* n (fact0 (- n 1)))))) And remember that this is actually just shorthand for: (define fact1 (lambda (n) (if (zero? n) 1 (* n ((lambda (n) (if (zero? n) 1 (* n (777 (- n 1))))) (- n 1)))))) We can continue in this way and write `fact2` that will work for n<=2: (define fact2 (lambda (n) (if (zero? n) 1 (* n (fact1 (- n 1)))))) or, in full form: (define fact2 (lambda (n) (if (zero? n) 1 (* n ((lambda (n) (if (zero? n) 1 (* n ((lambda (n) (if (zero? n) 1 (* n (777 (- n 1))))) (- n 1))))) (- n 1)))))) If we continue this way, we *will* get the true factorial function, but the problem is that to handle *any* possible integer argument, it will have to be an infinite definition! Here is what it is supposed to look like: (define fact0 (lambda (n) (if (zero? n) 1 (* n (777 (- n 1)))))) (define fact1 (lambda (n) (if (zero? n) 1 (* n (fact0 (- n 1)))))) (define fact2 (lambda (n) (if (zero? n) 1 (* n (fact1 (- n 1)))))) (define fact3 (lambda (n) (if (zero? n) 1 (* n (fact2 (- n 1)))))) ... The true factorial function is `fact-infinity`, with an infinite size. So, we're back at the original problem... To help make things more concise, we can observe the repeated pattern in the above, and extract a function that abstracts this pattern. This function is the same as the `fact-core` that we have seen previously: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact0 (fact-core 777)) (define fact1 (fact-core fact0)) (define fact2 (fact-core fact1)) (define fact3 (fact-core fact2)) ... which is actually: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact0 (fact-core 777)) (define fact1 (fact-core (fact-core 777))) (define fact2 (fact-core (fact-core (fact-core 777)))) ... (define fact (fact-core (fact-core (fact-core (... (fact-core 777) ...))))) Do this a little differently --- rewrite `fact0` as: (define fact0 ((lambda (mk) (mk 777)) fact-core)) Similarly, `fact1` is written as: (define fact1 ((lambda (mk) (mk (mk 777))) fact-core)) and so on, until the real factorial, which is still infinite at this stage: (define fact ((lambda (mk) (mk (mk (... (mk 777) ...)))) fact-core)) Now, look at that `(lambda (mk) ...)` --- it is an infinite expression, but for every actual application of the resulting factorial function we only need a finite number of `mk` applications. We can guess how many, and as soon as we hit an application of `777` we know that our guess is too small. So instead of `777`, we can try to use the maker function to create and use the next. To make things more explicit, here is the expression that is our `fact0`, without the definition form: ((lambda (mk) (mk 777)) fact-core) This function has a very low guess --- it works for 0, but with 1 it will run into the `777` application. At this point, we want to somehow invoke `mk` again to get the next level --- and since `777` *does* get applied, we can just replace it with `mk`: ((lambda (mk) (mk mk)) fact-core) The resulting function works just the same for an input of `0` because it does not attempt a recursive call --- but if we give it `1`, then instead of running into the error of applying `777`: (* n (777 (- n 1))) we get to apply `fact-core` there: (* n (fact-core (- n 1))) and this is still wrong, because `fact-core` expects a function as an input. To see what happens more clearly, write `fact-core` explicitly: ((lambda (mk) (mk mk)) (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) The problem is in what we're going to pass into `fact-core` --- its `fact` argument will not be the factorial function, but the `mk` function constructor. Renaming the `fact` argument as `mk` will make this more obvious (but not change the meaning): ((lambda (mk) (mk mk)) (lambda (mk) (lambda (n) (if (zero? n) 1 (* n (mk (- n 1))))))) It should now be obvious that this application of `mk` will not work, instead, we need to apply it on some function and *then* apply the result on `(- n 1)`. To get what we had before, we can use `777` as a bogus function: ((lambda (mk) (mk mk)) (lambda (mk) (lambda (n) (if (zero? n) 1 (* n ((mk 777) (- n 1))))))) This will allow one recursive call --- so the definition works for both inputs of `0` and `1` --- but not more. But that `777` is used as a maker function now, so instead, we can just use `mk` itself again: ((lambda (mk) (mk mk)) (lambda (mk) (lambda (n) (if (zero? n) 1 (* n ((mk mk) (- n 1))))))) And this is a *working* version of the real factorial function, so make it into a (non-magical) definition: (define fact ((lambda (mk) (mk mk)) (lambda (mk) (lambda (n) (if (zero? n) 1 (* n ((mk mk) (- n 1)))))))) But we're not done --- we "broke" into the factorial code to insert that `(mk mk)` application --- that's why we dragged in the actual value of `fact-core`. We now need to fix this. The expression on that last line (lambda (n) (if (zero? n) 1 (* n ((mk mk) (- n 1))))) is close enough --- it is `(fact-core (mk mk))`. So we can now try to rewrite our `fact` as: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact ((lambda (mk) (mk mk)) (lambda (mk) (fact-core (mk mk))))) ... and would fail in a familiar way! If it's not familiar enough, just rename all those `mk`s as `x`s: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact ((lambda (x) (x x)) (lambda (x) (fact-core (x x))))) We've run into the eagerness of our language again, as we did before. The solution is the same --- the `(x x)` is the factorial function, so protect it as we did before, and we have a working version: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define fact ((lambda (x) (x x)) (lambda (x) (fact-core (lambda (n) ((x x) n)))))) The rest should not be surprising now... Abstract the recursive making bit in a new `make-recursive` function: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) (define (make-recursive f) ((lambda (x) (x x)) (lambda (x) (f (lambda (n) ((x x) n)))))) (define fact (make-recursive fact-core)) and now we can do the first reduction inside `make-recursive` and write the `fact-core` expression explicitly: #lang pl broken (define (make-recursive f) ((lambda (x) (f (lambda (n) ((x x) n)))) (lambda (x) (f (lambda (n) ((x x) n)))))) (define fact (make-recursive (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))))) and this is the same code we had before. ------------------------------------------------------------------------ # The Y Combinator [Tuesday, February 11th] Our `make-recursive` function is usually called the *fixpoint operator* or the *Y combinator*. It looks really simple when using the lazy version (remember: our version is the eager one): (define Y (lambda (f) ((lambda (x) (f (x x))) (lambda (x) (f (x x)))))) > Note that if we *do* allow a recursive definition for Y itself, then > the definition can follow the definition that we've seen: > > (define (Y f) (f (Y f))) And this all comes from the loop generated by: ((lambda (x) (x x)) (lambda (x) (x x))) This expression, which is also called *Omega* (the `(lambda (x) (x x))` part by itself is usually called *omega* and then `(omega omega)` is *Omega*), is also the idea behind many deep mathematical facts. As an example for what it does, follow the next rule: I will say the next sentence twice: "I will say the next sentence twice". (Note the usage of colon for the first and quotes for the second --- what is the equivalent of that in the lambda expression?) By itself, this just gets you stuck in an infinite loop, as Omega does, and the Y combinator adds `F` to that to get an infinite chain of applications --- which is similar to: I will say the next sentence twice: "I will hop on one foot and then say the next sentence twice". ------------------------------------------------------------------------ ## The main property of Y [Tuesday, February 11th] `fact-core` is a function that given any limited factorial, will generate a factorial that is good for one more integer input. Start with `777`, which is a factorial that is good for nothing (because it's not a function), and you can get `fact0` as fact0 == (fact-core 777) and that's a good factorial function only for an input of `0`. Use that with `fact-core` again, and you get fact1 == (fact-core fact0) == (fact-core (fact-core 777)) which is the factorial function when you only look at input values of `0` or `1`. In a similar way fact2 == (fact-core fact1) is good for `0`...`2` --- and we can continue as much as we want, except that we need to have an infinite number of applications --- in the general case, we have: fact-n == (fact-core (fact-core (fact-core ... 777))) which is good for `0`...`n`. The *real* factorial would be the result of running `fact-core` on itself infinitely, it *is* `fact-infinity`. In other words (here `fact` is the *real* factorial): fact = fact-infinity == (fact-core (fact-core ...infinitely...)) but note that since this is really infinity, then fact = (fact-core (fact-core ...infinitely...)) = (fact-core fact) so we get an equation: fact = (fact-core fact) and a solution for this is going to be the real factorial. The solution is the *fixed-point* of the `fact-core` function, in the same sense that `0` is the fixed point of the `sin` function because 0 = (sin 0) And the Y combinator does just that --- it has this property: (make-recursive f) = (f (make-recursive f)) or, using the more common name: (Y f) = (f (Y f)) This property encapsulates the real magical power of Y. You can see how it works: since `(Y f) = (f (Y f))`, we can add an `f` application to both sides, giving us `(f (Y f)) = (f (f (Y f)))`, so we get: (Y f) = (f (Y f)) = (f (f (Y f))) = (f (f (f (Y f)))) = ... = (f (f (f ...))) and we can conclude that (Y fact-core) = (fact-core (fact-core ...infinitely...)) = fact ------------------------------------------------------------------------ # Yet another explanation for Y [Tuesday, February 11th] Here's another explanation of how the Y combinator works. Remember that our `fact-core` function was actually a function that generates a factorial function based on some input, which is supposed to be the factorial function: (define fact-core (lambda (fact) (lambda (n) (if (zero? n) 1 (* n (fact (- n 1))))))) As we've seen, you can apply this function on a version of factorial that is good for inputs up to some n, and the result will be a factorial that is good for those values up to n+1. The question is *what is the fixpoint of `fact-core`*? And the answer is that if it maps factₙ factorial to factₙ₊₁, then the input will be equal to the output on the *infinitieth* `fact`, which is the *actual* factorial. Since Y is a fixpoint combinator, it gives us exactly that answer: (define the-real-factorial (Y fact-core)) ------------------------------------------------------------------------ # Typing the Y Combinator [Tuesday, February 11th] Typing the Y combinator is a tricky issue. For example, in standard ML you must write a new type definition to do this: datatype 'a t = T of 'a t -> 'a val y = fn f => (fn (T x) => (f (fn a => x (T x) a))) (T (fn (T x) => (f (fn a => x (T x) a)))) > Can you find a pattern in the places where `T` is used? > --- Roughly speaking, that type definition is > > ;; `t' is the type name, `T' is the constructor (aka the variant) > (define-type (RecType a) ; we don't really have polymorphic types > [T ((RecType a) -> a)]) > > First note that the two `fn a => ...` parts are the same as our > protection, so ignoring that we get: > > val y = fn f => (fn (T x) => (f (x (T x)))) > (T (fn (T x) => (f (x (T x))))) > > if you now replace `T` with `Quote`, things make more sense: > > val y = fn f => (fn (Quote x) => (f (x (Quote x)))) > (Quote (fn (Quote x) => (f (x (Quote x))))) > > and with our syntax, this would be: > > (define (Y f) > ((lambda (qx) > (cases qx > [(Quote x) (f (x (Quote x)))])) > (Quote > (lambda (qx) > (cases qx > [(Quote x) (f (x (Quote x)))]))))) > > it's not really quotation --- but the analogy should help: it uses > `Quote` to distinguish functions as values that are applied (the `x`s) > from functions that are passed as arguments. In OCaml, this looks a little different: # type 'a t = T of ('a t -> 'a) ;; type 'a t = T of ('a t -> 'a) # let y f = (fun (T x) -> x (T x)) (T (fun (T x) -> fun z -> f (x (T x)) z)) ;; val y : (('a -> 'b) -> 'a -> 'b) -> 'a -> 'b = # let fact = y (fun fact n -> if n<1 then 1 else n* fact(n-1)) ;; val fact : int -> int = # fact 5 ;; - : int = 120 but OCaml has also a `-rectypes` command line argument, which will make it infer the type by itself: # let y f = (fun x -> x x) (fun x -> fun z -> f (x x) z) ;; val y : (('a -> 'b) -> 'a -> 'b) -> 'a -> 'b = # let fact = y (fun fact n -> if n<1 then 1 else n* fact(n-1)) ;; val fact : int -> int = # fact 5 ;; - : int = 120 It is also possible to write this expression in Typed Racket, but we will need to write a proper type definition. First of all, the type of Y should be straightforward: it is a fixpoint operation, so it takes a `T -> T` function and produces its fixpoint. The fixpoint itself is some `T` (such that applying the function on it results in itself). So this gives us: (: make-recursive : (T -> T) -> T) However, in our case `make-recursive` computes a *functional* fixpoint, for unary `S -> T` functions, so we should narrow down the type (: make-recursive : ((S -> T) -> (S -> T)) -> (S -> T)) Now, in the body of `make-recursive` we need to add a type for the `x` argument which is behaving in a weird way: it is used both as a function and as its own argument. (Remember --- I will say the next sentence twice: "I will say the next sentence twice".) We need a recursive type definition for that: (define-type (Tau S T) = (Rec this (this -> (S -> T)))) This type is tailored for our use of `x`: given a type `T`, `x` is a function that will consume *itself* (hence the `Rec`) and spit out the value that the `f` argument consumes --- an `S -> T` function. The resulting full version of the code: (: make-recursive : (All (S T) ((S -> T) -> (S -> T)) -> (S -> T))) (define-type (Tau S T) = (Rec this (this -> (S -> T)))) (define (make-recursive f) ((lambda ([x : (Tau S T)]) (f (lambda (z) ((x x) z)))) (lambda ([x : (Tau S T)]) (f (lambda (z) ((x x) z)))))) (: fact : Number -> Number) (define fact (make-recursive (lambda ([fact : (Number -> Number)]) (lambda ([n : Number]) (if (zero? n) 1 (* n (fact (- n 1)))))))) (fact 5) ------------------------------------------------------------------------ # Lambda Calculus --- Schlac [Tuesday, February 11th] > [PLAI §22] (we do much more) We know that many constructs that are usually thought of as primitives are not really needed --- we can implement them ourselves given enough tools. The question is how far can we go? The answer: as far as we want. For example: (define foo((lambda(f)((lambda(x)(x x))(lambda(x)(f(x x)))))(lambda( f)(lambda(x)(((x(lambda(x)(lambda(x y)y))(lambda(x y)x))(x(lambda(x) (lambda(x y)y))(lambda(x y)x))(((x(lambda (p)(lambda(s)(s(p(lambda(x y)y))(lambda(f x)(f((p(lambda(x y)y))f x))))))(lambda(s) (s(lambda(f x)x)(lambda(f x)x))))(lambda(x y)x))(lambda(x)(lambda(x y)y))(lambda (x y)x)))(lambda(f x)(f x))((f((x(lambda(p)(lambda(s)(s(p(lambda(x y )y))(lambda(f x)(f((p(lambda(x y)y))f x))))))(lambda(y s)(s(lambda(f x)x)(lambda(f x)x))))(lambda(x y)x)))(lambda(n)(lambda(f x)(f(n f x) )))(f((((x(lambda(p)(lambda(s)(s(p (lambda(x y)y))(lambda(f x)(f((p( lambda(x y)y))f x))))))(lambda(s)(s(lambda(f x) x)(lambda(f x)x))))( lambda(x y)x))(lambda(p)(lambda(s)(s(p(lambda(x y)y))(lambda(f x)(f( (p(lambda(x y)y))f x))))))(lambda(s)(s(lambda(f x)x)(lambda(f x)x))) )(lambda(x y)x))))))))) We begin with a very minimal language, which is based on the Lambda Calculus. In this language we get a very minimal set of constructs and values. In DrRacket, this we will use the Schlac language level (stands for "SchemeRacket as Lambda Calculus"). This language has a Racket-like syntax, but don't be confused --- it is *very* different from Racket. The only constructs that are available in this language are: lambda expressions of at least one argument, function application (again, at least one argument), and simple definition forms which are similar to the ones in the "Broken define" language --- definitions are used as shorthand, and cannot be used for recursive function definition. They're also only allowed at the toplevel --- no local helpers, and a definition is not an expression that can appear anywhere. The BNF is therefore: ::= ... ::= | (define ) ::= | (lambda ( ...) ) | ( ...) Since this language has no primitive values (other than functions), Racket numbers and booleans are also considered identifiers, and have no built-in value that come with the language. In addition, all functions and function calls are curried, so (lambda (x y z) (z y x)) is actually shorthand for (lambda (x) (lambda (y) (lambda (z) ((z y) x)))) The rules for evaluation are simple, there is one very important rule for evaluation which is called "beta reduction": ((lambda (x) E1) E2) --> E1[E2/x] where substitution in this context requires being careful so you won't capture names. This requires you to be able to do another kind of transformation which is called "alpha conversion", which basically says that you can rename identifiers as long as you keep the same binding structure (eg, a valid renaming does not change the de-Bruijn form of the expression). There is one more rule that can be used, *eta conversion* which says that `(lambda (x) (f x))` is the same as `f` (we used this rule above when deriving the Y combinator). One last difference between Schlac and Racket is that Schlac is a *lazy* language. This will be important since we do not have any built-in special forms like `if`. Here is a Schlac definition for the identity function: (define identity (lambda (x) x)) and there is not much that we can do with this now: > identity # > (identity identity) # > (identity identity identity) # (In the last expression, note that `(id id id)` is shorthand for `((id id) id)`, and since `(id id)` is the identity, applying that on `id` returns it again.) ------------------------------------------------------------------------ # Church Numerals [Tuesday, February 11th] So far, it seems like it is impossible to do anything useful in this language, since all we have are functions and applications. We know how to write the identity function, but what about other values? For example, can you write code that evaluates to zero? > What's zero? I only know how to write functions! > > (Turing Machine programmer: "What's a function? > --- I only know how to write 0s and 1s!") The first thing we therefore need is to be able to *encode* numbers as functions. For zero, we will use a function of two arguments that simply returns its second value: (define 0 (lambda (f) (lambda (x) x))) or, more concisely (define 0 (lambda (f x) x)) This is the first step in an encoding that is known as *Church Numerals*: an encoding of natural numbers as functions. The number zero is encoded as a function that takes in a function and a second value, and applies the function zero times on the argument (which is really what the above definition is doing). Following this view, the number one is going to be a function of two arguments, that applies the first on the second one time: (define 1 (lambda (f x) (f x))) and note that `1` is just like the identity function (as long as you give it a function as its first input, but this is always the case in Schlac). The next number on the list is two --- which applies the first argument on the second one twice: (define 2 (lambda (f x) (f (f x)))) We can go on doing this, but what we really want is a way to perform arbitrary arithmetic. The first requirement for that is an `add1` function that increments its input (an encoded natural number) by one. To do this, we write a function that expects an encoded number: (define add1 (lambda (n) ...)) and this function is expected to return an encoded number, which is always a function of `f` and `x`: (define add1 (lambda (n) (lambda (f x) ...))) Now, in the body, we need to apply `f` on `x` n+1 times --- but remember that `n` is a function that will do `n` applications of its first argument on its second: (define add1 (lambda (n) (lambda (f x) ... (n f x) ...))) and all we have left to do now is to apply `f` one more time, yielding this definition for `add1`: (define add1 (lambda (n) (lambda (f x) (f (n f x))))) Using this, we can define a few useful numbers: (define 1 (add1 0)) (define 2 (add1 1)) (define 3 (add1 2)) (define 4 (add1 3)) (define 5 (add1 4)) This is all nice theoretically, but how can we make sure that it is correct? Well, Schlac has a few additional special forms that translate Church numerals into Racket numbers. To try our definitions we use the `->nat` (read: to natural number): (->nat 0) (->nat 5) (->nat (add1 (add1 5))) You can now verify that the identity function is really the same as the number 1: (->nat identity) We can even write a test case, since Schlac contains the `test` special form, but we have to be careful in that --- first of all, we cannot test whether functions are equal (why?) so we must use `->nat`, but (test (->nat (add1 (add1 5))) => 7) will not work since `7` is undefined. To overcome this, Schlac has a `back-door` for primitive Racket values --- just use a quote: (test (->nat (add1 (add1 5))) => '7) We can now define natural number addition --- one simple idea is to get two encoded numbers `m` and `n`, then start with `x`, apply `f` on it `n` times by using it as a function, then apply `f` `m` more times on the result in the same way: (define + (lambda (m n) (lambda (f x) (m f (n f x))))) or equivalently: (define + (lambda (m n f x) (m f (n f x)))) Another idea is to use `add1` and increment `n` by `m` using `add1`: (define + (lambda (m n) (m add1 n))) (->nat (+ 4 5)) We can also define multiplication of `m` and `n` quite easily --- begin with addition --- `(lambda (x) (+ n x))` is a function that expects an `x` and returns `(+ x n)` --- it's an increment-by-n function. But since all functions and applications are curried, this is actually the same as `(lambda (x) ((+ n) x))` which is the same as `(+ n)`. Now, what we want to do is repeat this operation `m` times over zero, which will add `n` to zero `m` times, resulting in `m` * `n`. The definition is therefore: (define * (lambda (m n) (m (+ n) 0))) (->nat (* 4 5)) (->nat (+ 4 (* (+ 2 5) 5))) An alternative approach is to consider (lambda (x) (n f x)) for some encoded number `n` and a function `f` --- this function is like `f`^`n` (f composed n times with itself). But remember that this is shorthand for (lambda (x) ((n f) x)) and we know that `(lambda (x) (foo x))` is just like `foo` (if it is a function), so this is equivalent to just (n f) So `(n f)` is `f`^`n`, and in the same way `(m g)` is `g`^`m` --- if we use `(n f)` for `g`, we get `(m (n f))` which is n self-compositions of `f`, self-composed m times. In other words, `(m (n f))` is a function that is like `m`*`n` applications of `f`, so we can define multiplication as: (define * (lambda (m n) (lambda (f) (m (n f))))) which is the same as (define * (lambda (m n f) (m (n f)))) The same principle can be used to define exponentiation (but now we have to be careful with the order since exponentiation is not commutative): (define ^ (lambda (m n) (n (* m) 1))) (->nat (^ 3 4)) And there is a similar alternative here too --- * a Church numeral `m` is the m-self-composition function, * and `(1 m)` is just like `m`^`1` which is the same as `m` (`1`=`identity`) * and `(2 m)` is just like `m`^`2` --- it takes a function `f`, self composes it `m` times, and self composes the result `m` times --- for a total of `f`^`(m*m)` * and `(3 m)` is similarly `f`^`(m*m*m)` * so `(n m)` is `f`^`(m^n)` (note that the first `^` is self-compositions, and the second one is a mathematical exponent) * so `(n m)` is a function that returns `m`^`n` self-compositions of an input function, Which means that `(n m)` is the Church numeral for `m`^`n`, so we get: (define ^ (lambda (m n) (n m))) which basically says that any number encoding `n` is also the `?`^`n` operation. All of this is was not too complicated --- but all so far all we did is write functions that increment their inputs in various ways. What about `sub1`? For that, we need to do some more work --- we will need to encode booleans. ------------------------------------------------------------------------ # More Encodings [Tuesday, February 11th] Our choice of encoding numbers makes sense --- the idea is that the main feature of a natural number is repeating something a number of times. For booleans, the main property we're looking for is choosing between two values. So we can encode true and false by functions of two arguments that return either the first or the second argument: (define #t (lambda (x y) x)) (define #f (lambda (x y) y)) Note that this encoding of `#f` is really the same as the encoding of `0`, so we have to know what type to expect an use the proper operations (this is similar to C, where everything is just integers). Now that we have these two, we can define `if`: (define if (lambda (c t e) (c t e))) it expects a boolean which is a function of two arguments, and passes it the two expressions. The `#t` boolean will simply return the first, and the `#f` boolean will return the second. Strictly speaking, we don't really need this definition, since instead of writing `(if c t e)`, we can simply write `(c t e)`. In any case, we need the language to be lazy for this to work. To demonstrate this, we'll intentionally use the quote back-door to use a non-functional value, using this will normally result in an error: (+ '1 '2) But testing our `if` definition, things work just fine: (if #t (+ 4 5) (+ 1 2)) and we see that DrRacket leaves the second addition expression in red, which indicates that it was not executed. We can also make sure that even when it is defined as a function, it is still working fine because the language is lazy: (if #f ((lambda (x) (x x)) (lambda (x) (x x))) 3) What about `and` and `or`? Simple, `or` takes two arguments, and returns either true or false if one of the inputs is true: (define or (lambda (a b) (if a #t (if b #t #f)))) but `(if b #t #f)` is really the same as just `b` because it must be a boolean (we cannot use more than one "truty" or "falsy" values): (define or (lambda (a b) (if a #t b))) also, if `a` is true, we want to return `#t`, but that is exactly the value of `a`, so: (define or (lambda (a b) (if a a b))) and finally, we can get rid of the `if` (which is actually breaking the `if` abstraction, if we encode booleans in some other way): (define or (lambda (a b) (a a b))) Similarly, you can convince yourself that the definition of `and` is: (define and (lambda (a b) (a b a))) Schlac has to-Racket conversion forms for booleans too: (->bool (or #f #f)) (->bool (or #f #t)) (->bool (or #t #f)) (->bool (or #t #t)) and (->bool (and #f #f)) (->bool (and #f #t)) (->bool (and #t #f)) (->bool (and #t #t)) A `not` function is quite simple --- one alternative is to choose from true and false in the usual way: (define not (lambda (a) (a #f #t))) and another is to return a function that switches the inputs to an input boolean: (define not (lambda (a) (lambda (x y) (a y x)))) which is the same as (define not (lambda (a x y) (a y x))) We can now put numbers and booleans together: we define a `zero?` function. (define zero? (lambda (n) (n (lambda (x) #f) #t))) (test (->bool (and (zero? 0) (not (zero? 3)))) => '#t) (Good question: is this fast?) (Note that it is better to test that the value is explicitly `#t`, if we just use `(test (->bool ...))` then the test will work even if the expression in question evaluated to some bogus value.) The idea is simple --- if `n` is the encoding of zero, it will return it's second argument which is `#t`: (zero? 0) --> ((lambda (f n) n) (lambda (x) #f) #t) -> #t if `n` is an encoding of a bigger number, then it is a self-composition, and the function that we give it is one that always returns `#f`, no matter how many times it is self-composed. Try `2` for example: (zero? 2) --> ((lambda (f n) (f (f n))) (lambda (x) #f) #t) --> ((lambda (x) #f) ((lambda (x) #f) #t)) --> #f Now, how about an encoding for compound values? A minimal approach is what we use in Racket --- a way to generate pairs (`cons`), and encode lists as chains of pairs with a special value at the end (`null`). There is a natural encoding for pairs that we have previously seen --- a pair is a function that expects a selector, and will apply that on the two values: (define cons (lambda (x y) (lambda (s) (s x y)))) Or, equivalently: (define cons (lambda (x y s) (s x y))) To extract the two values from a pair, we need to pass a selector that consumes two values and returns one of them. In our framework, this is exactly what the two boolean values do, so we get: (define car (lambda (x) (x #t))) (define cdr (lambda (x) (x #f))) (->nat (+ (car (cons 2 3)) (cdr (cons 2 3)))) We can even do this: (define 1st car) (define 2nd (lambda (l) (car (cdr l)))) (define 3rd (lambda (l) (car (cdr (cdr l))))) (define 4th (lambda (l) (car (cdr (cdr (cdr l)))))) (define 5th (lambda (l) (car (cdr (cdr (cdr (cdr l))))))) or write a `list-ref` function: (define list-ref (lambda (l n) (car (n cdr l)))) Note that we don't need a recursive function for this: our encoding of natural numbers makes it easy to "iterate N times". What we get with this encoding is essentially free natural-number recursion. We now need a special `null` value to mark list ends. This value should have the same number of arguments as a `cons` value (one: a selector/boolean function), and it should be possible to distinguish it from other values. We choose (define null (lambda (s) #t)) Testing the list encoding: (define l123 (cons 1 (cons 2 (cons 3 null)))) (->nat (2nd l123)) And as with natural numbers and booleans, Schlac has built-in facility to convert encoded lists to Racket values, except that this requires specifying the type of values in a list so it's a higher-order function: ((->listof ->nat) l123) which ("as usual") can be written as (->listof ->nat l123) We can even do this: (->listof (->listof ->nat) (cons l123 (cons l123 null))) Defining `null?` is now relatively easy (and it's actually already used by the above `->listof` conversion). The following definition (define null? (lambda (x) (x (lambda (x y) #f)))) works because if `x` is null, then it simply ignores its argument and returns `#t`, and if it's a pair, then it uses the input selector, which always returns `#f` in its turn. Using some arbitrary `A` and `B`: (null? (cons A B)) --> ((lambda (x) (x (lambda (x y) #f))) (lambda (s) (s A B))) --> ((lambda (s) (s A B)) (lambda (x y) #f)) --> ((lambda (x y) #f) A B) --> #f (null? null) --> ((lambda (x) (x (lambda (x y) #f))) (lambda (s) #t)) --> ((lambda (s) #t) (lambda (x y) #f)) --> #t We can use the Y combinator to create recursive functions --- we can even use the rewrite rules facility that Schlac contains (the same one that we have previously seen): (define Y (lambda (f) ((lambda (x) (x x)) (lambda (x) (f (x x)))))) (rewrite (define/rec f E) => (define f (Y (lambda (f) E)))) and using it: (define/rec length (lambda (l) (if (null? l) 0 (add1 (length (cdr l)))))) (->nat (length l123)) And to complete this, um, journey --- we're still missing subtraction. There are many ways to solve the problem of subtraction, and for a challenge try to come up with a solution yourself. One of the clearer solutions uses a simple idea --- begin with a pair of two zeroes `<0,0>`, and repeat this transformation `n` times: `` -> ``. After `n` steps, we will have `` --- so we get: (define inccons (lambda (p) (cons (cdr p) (add1 (cdr p))))) (define sub1 (lambda (n) (car (n inccons (cons 0 0))))) (->nat (sub1 5)) And from this the road is short to general subtraction, `m`-`n` is simply `n` applications of `sub1` on `m`: (define - (lambda (m n) (n sub1 m))) (test (->nat (- 3 2)) => '1) (test (->nat (- (* 4 (* 5 5)) 5)) => '95) We now have a normal-looking language, and we're ready to do anything we want. Here are two popular examples: (define/rec fact (lambda (x) (if (zero? x) 1 (* x (fact (sub1 x)))))) (test (->nat (fact 5)) => '120) (define/rec fib (lambda (x) (if (or (zero? x) (zero? (sub1 x))) 1 (+ (fib (- x 1)) (fib (- x 2)))))) (test (->nat (fib (* 5 2))) => '89) To get generalized arithmetic capability, Schlac has yet another built-in facility for translating Racket natural numbers into Church numerals: (->nat (fib (nat-> '10))) ... and to get to that frightening expression in the beginning, all you need to do is replace all definitions in the `fib` definition over and over again until you're left with nothing but lambda expressions and applications, then reformat the result into some cute shape. For extra fun, you can look for immediate applications of lambda expressions and reduce them manually. All of this is in the following code: ;;; ---<<>>------------------------------------------------- ;; Making Schlac into a practical language (not an interpreter) #lang pl schlac (define identity (lambda (x) x)) ;; Natural numbers (define 0 (lambda (f x) x)) (define add1 (lambda (n) (lambda (f x) (f (n f x))))) ;; same as: ;; (define add1 (lambda (n) (lambda (f x) (n f (f x))))) (define 1 (add1 0)) (define 2 (add1 1)) (define 3 (add1 2)) (define 4 (add1 3)) (define 5 (add1 4)) (test (->nat (add1 (add1 5))) => '7) (define + (lambda (m n) (m add1 n))) (test (->nat (+ 4 5)) => '9) ;; (define * (lambda (m n) (m (+ n) 0))) (define * (lambda (m n f) (m (n f)))) (test (->nat (* 4 5)) => '20) (test (->nat (+ 4 (* (+ 2 5) 5))) => '39) ;; (define ^ (lambda (m n) (n (* m) 1))) (define ^ (lambda (m n) (n m))) (test (->nat (^ 3 4)) => '81) ;; Booleans (define #t (lambda (x y) x)) (define #f (lambda (x y) y)) (define if (lambda (c t e) (c t e))) ; not really needed (test (->nat (if #t 1 2)) => '1) (test (->nat (if #t (+ 4 5) (+ '1 '2))) => '9) (define and (lambda (a b) (a b a))) (define or (lambda (a b) (a a b))) ;; (define not (lambda (a) (a #f #t))) (define not (lambda (a x y) (a y x))) (test (->bool (and #f #f)) => '#f) (test (->bool (and #t #f)) => '#f) (test (->bool (and #f #t)) => '#f) (test (->bool (and #t #t)) => '#t) (test (->bool (or #f #f)) => '#f) (test (->bool (or #t #f)) => '#t) (test (->bool (or #f #t)) => '#t) (test (->bool (or #t #t)) => '#t) (test (->bool (not #f)) => '#t) (test (->bool (not #t)) => '#f) (define zero? (lambda (n) (n (lambda (x) #f) #t))) (test (->bool (and (zero? 0) (not (zero? 3)))) => '#t) ;; Lists (define cons (lambda (x y s) (s x y))) (define car (lambda (x) (x #t))) (define cdr (lambda (x) (x #f))) (test (->nat (+ (car (cons 2 3)) (cdr (cons 2 3)))) => '5) (define 1st car) (define 2nd (lambda (l) (car (cdr l)))) (define 3rd (lambda (l) (car (cdr (cdr l))))) (define 4th (lambda (l) (car (cdr (cdr (cdr l)))))) (define 5th (lambda (l) (car (cdr (cdr (cdr (cdr l))))))) (define null (lambda (s) #t)) (define null? (lambda (x) (x (lambda (x y) #f)))) (define l123 (cons 1 (cons 2 (cons 3 null)))) ;; Note that `->listof' is a H.O. converter (test ((->listof ->nat) l123) => '(1 2 3)) (test (->listof ->nat l123) => '(1 2 3)) ; same as the above (test (->listof (->listof ->nat) (cons l123 (cons l123 null))) => '((1 2 3) (1 2 3))) ;; Subtraction is tricky (define inccons (lambda (p) (cons (cdr p) (add1 (cdr p))))) (define sub1 (lambda (n) (car (n inccons (cons 0 0))))) (test (->nat (sub1 5)) => '4) (define - (lambda (a b) (b sub1 a))) (test (->nat (- 3 2)) => '1) (test (->nat (- (* 4 (* 5 5)) 5)) => '95) (test (->nat (- 2 4)) => '0) ; this is "natural subtraction" ;; Recursive functions (define Y (lambda (f) ((lambda (x) (x x)) (lambda (x) (f (x x)))))) (rewrite (define/rec f E) => (define f (Y (lambda (f) E)))) (define/rec length (lambda (l) (if (null? l) 0 (add1 (length (cdr l)))))) (test (->nat (length l123)) => '3) (define/rec fact (lambda (x) (if (zero? x) 1 (* x (fact (sub1 x)))))) (test (->nat (fact 5)) => '120) (define/rec fib (lambda (x) (if (or (zero? x) (zero? (sub1 x))) 1 (+ (fib (sub1 x)) (fib (sub1 (sub1 x))))))) (test (->nat (fib (* 5 2))) => '89) #| ;; Fully-expanded Fibonacci (define fib ((lambda (f) ((lambda (x) (x x)) (lambda (x) (f (x x))))) (lambda (f) (lambda (x) ((lambda (c t e) (c t e)) ((lambda (a b) (a a b)) ((lambda (n) (n (lambda (x) (lambda (x y) y)) (lambda (x y) x))) x) ((lambda (n) (n (lambda (x) (lambda (x y) y)) (lambda (x y) x))) ((lambda (n) ((lambda (x) (x (lambda (x y) x))) (n (lambda (p) ((lambda (x y s) (s x y)) ((lambda (x) (x (lambda (x y) y))) p) ((lambda (n) (lambda (f x) (f (n f x)))) ((lambda (x) (x (lambda (x y) y))) p)))) ((lambda (x y s) (s x y)) (lambda (f x) x) (lambda (f x) x))))) x))) ((lambda (n) (lambda (f x) (f (n f x)))) (lambda (f x) x)) ((lambda (x y) (x (lambda (n) (lambda (f x) (f (n f x)))) y)) (f ((lambda (n) ((lambda (x) (x (lambda (x y) x))) (n (lambda (p) ((lambda (x y s) (s x y)) ((lambda (x) (x (lambda (x y) y))) p) ((lambda (n) (lambda (f x) (f (n f x)))) ((lambda (x) (x (lambda (x y) y))) p)))) ((lambda (x y s) (s x y)) (lambda (f x) x) (lambda (f x) x))))) x)) (f ((lambda (n) ((lambda (x) (x (lambda (x y) x))) (n (lambda (p) ((lambda (x y s) (s x y)) ((lambda (x) (x (lambda (x y) y))) p) ((lambda (n) (lambda (f x) (f (n f x)))) ((lambda (x) (x (lambda (x y) y))) p)))) ((lambda (x y s) (s x y)) (lambda (f x) x) (lambda (f x) x))))) ((lambda (n) ((lambda (x) (x (lambda (x y) x))) (n (lambda (p) ((lambda (x y s) (s x y)) ((lambda (x) (x (lambda (x y) y))) p) ((lambda (n) (lambda (f x) (f (n f x)))) ((lambda (x) (x (lambda (x y) y))) p)))) ((lambda (x y s) (s x y)) (lambda (f x) x) (lambda (f x) x))))) x))))))))) ;; The same after reducing all immediate function applications (define fib ((lambda (f) ((lambda (x) (x x)) (lambda (x) (f (x x))))) (lambda (f) (lambda (x) (((x (lambda (x) (lambda (x y) y)) (lambda (x y) x)) (x (lambda (x) (lambda (x y) y)) (lambda (x y) x)) (((x (lambda (p) (lambda (s) (s (p (lambda (x y) y)) (lambda (f x) (f ((p (lambda (x y) y)) f x)))))) (lambda (s) (s (lambda (f x) x) (lambda (f x) x)))) (lambda (x y) x)) (lambda (x) (lambda (x y) y)) (lambda (x y) x))) (lambda (f x) (f x)) ((f ((x (lambda (p) (lambda (s) (s (p (lambda (x y) y)) (lambda (f x) (f ((p (lambda (x y) y)) f x)))))) (lambda (y s) (s (lambda (f x) x) (lambda (f x) x)))) (lambda (x y) x))) (lambda (n) (lambda (f x) (f (n f x)))) (f ((((x (lambda (p) (lambda (s) (s (p (lambda (x y) y)) (lambda (f x) (f ((p (lambda (x y) y)) f x)))))) (lambda (s) (s (lambda (f x) x) (lambda (f x) x)))) (lambda (x y) x)) (lambda (p) (lambda (s) (s (p (lambda (x y) y)) (lambda (f x) (f ((p (lambda (x y) y)) f x)))))) (lambda (s) (s (lambda (f x) x) (lambda (f x) x)))) (lambda (x y) x))))))))) ;; Cute reformatting of the above: (define fib((lambda(f)((lambda(x)(x x))(lambda(x)(f(x x)))))(lambda( f)(lambda(x)(((x(lambda(x)(lambda(x y)y))(lambda(x y)x))(x(lambda(x) (lambda(x y)y))(lambda(x y) x))(((x(lambda(p)(lambda(s)(s(p(lambda(x y)y))(lambda(f x)(f((p(lambda(x y)y))f x))))))(lambda(s) (s(lambda(f x)x)(lambda(f x)x))))(lambda(x y)x))(lambda(x)(lambda(x y)y))(lambda (x y)x)))(lambda(f x)(f x))((f((x(lambda(p)(lambda(s)(s(p(lambda(x y )y))(lambda(f x)(f((p(lambda(x y)y))f x))))))(lambda(y s)(s(lambda(f x)x)(lambda(f x)x))))(lambda(x y)x)))(lambda(n)(lambda(f x)(f(n f x) )))(f((((x(lambda(p)(lambda(s)(s(p (lambda(x y)y))(lambda(f x)(f((p( lambda(x y) y))f x))))))(lambda(s)(s(lambda(f x)x)(lambda(f x)x))))( ;; ^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^ ;; `---------------(cons 0 0)---------------' lambda(x y)x))(lambda(p)(lambda(s)(s(p(lambda(x y)y))(lambda(f x)(f( (p(lambda(x y)y))f x))))))(lambda(s)(s(lambda(f x)x)(lambda(f x)x))) )(lambda(x y)x))))))))) And for extra fun: (λ(f)(λ (x)(((x(λ( x)(λ(x y)y) )(λ(x y)x))( x(λ(x)(λ(x y) y))(λ(x y )x))((( x(λ(p)( λ(s)(s (p (λ( x y)y)) (λ(f x )(f((p( λ(x y) y))f x ))))))( λ(s)(s( λ(f x)x) (λ(f x)x) )))(λ(x y) x))(λ(x)(λ( x y)y)) (λ( x y) x)))(λ( f x)(f x))((f ((x(λ(p )(λ (s )(s(p( λ(x y) y))(λ ( f x)(f( (p (λ( x y)y) )f x))) )))(λ( y s)(s (λ (f x )x)(λ( f x)x) )))(λ( x y)x)) )(λ(n) (λ (f x)(f (n f x))) )(f((( (x(λ(p) (λ(s)(s (p( λ( x y )y ))(λ(f x) (f(( p(λ(x y )y)) f x))))) )(λ(s)( s(λ(f x )x)(λ( f x)x) ))) (λ (x y)x ))(λ(p )(λ(s)( s(p(λ( x y)y) )(λ (f x)(f(( p(λ (x y)y)) f x)))))) (λ(s)( s(λ (f x)x)(λ (f x)x) )))(λ( x y)x) )))))) |# ------------------------------------------------------------------------ # Alternative Church Encoding [Tuesday, February 18th] Finally, note that this is just one way to encode things --- other encodings are possible. One alternative encoding is in the following code --- it uses a list of `N` falses as the encoding for `N`. This encoding makes it easier to `add1` (just `cons` another `#f`), and to `sub1` (simply `cdr`). The tradeoff is that some arithmetics operations becomes more complicated, for example, the definition of `+` requires the fixpoint combinator. (As expected, some people want to see what can we do with a language without recursion, so they don't like jumping to Y too fast.) ;;; ---<<>>------------------------------------- ;; An alternative "Church" encoding: use lists to encode numbers #lang pl schlac (define identity (lambda (x) x)) ;; Booleans (same as before) (define #t (lambda (x y) x)) (define #f (lambda (x y) y)) (define if (lambda (c t e) (c t e))) ; not really needed (test (->bool (if #t #f #t)) => '#f) (test (->bool (if #f ((lambda (x) (x x)) (lambda (x) (x x))) #t)) => '#t) (define and (lambda (a b) (a b a))) (define or (lambda (a b) (a a b))) (define not (lambda (a x y) (a y x))) (test (->bool (and #f #f)) => '#f) (test (->bool (and #t #f)) => '#f) (test (->bool (and #f #t)) => '#f) (test (->bool (and #t #t)) => '#t) (test (->bool (or #f #f)) => '#f) (test (->bool (or #t #f)) => '#t) (test (->bool (or #f #t)) => '#t) (test (->bool (or #t #t)) => '#t) (test (->bool (not #f)) => '#t) (test (->bool (not #t)) => '#f) ;; Lists (same as before) (define cons (lambda (x y s) (s x y))) (define car (lambda (x) (x #t))) (define cdr (lambda (x) (x #f))) (define 1st car) (define 2nd (lambda (l) (car (cdr l)))) (define 3rd (lambda (l) (car (cdr (cdr l))))) (define 4th (lambda (l) (car (cdr (cdr (cdr l)))))) (define 5th (lambda (l) (car (cdr (cdr (cdr (cdr l))))))) (define null (lambda (s) #t)) (define null? (lambda (x) (x (lambda (x y) #f)))) ;; Natural numbers (alternate encoding) (define 0 identity) (define add1 (lambda (n) (cons #f n))) (define zero? car) ; tricky (define sub1 cdr) ; this becomes very simple ;; Note that we could have used something more straightforward: ;; (define 0 null) ;; (define add1 (lambda (n) (cons #t n))) ; cons anything ;; (define zero? null?) ;; (define sub1 (lambda (l) (if (zero? l) l (cdr l)))) (define 1 (add1 0)) (define 2 (add1 1)) (define 3 (add1 2)) (define 4 (add1 3)) (define 5 (add1 4)) (test (->nat* (add1 (add1 5))) => '7) (test (->nat* (sub1 (sub1 (add1 (add1 5))))) => '5) (test (->bool (and (zero? 0) (not (zero? 3)))) => '#t) (test (->bool (zero? (sub1 (sub1 (sub1 3))))) => '#t) ;; list-of-numbers tests (define l123 (cons 1 (cons 2 (cons 3 null)))) (test (->listof ->nat* l123) => '(1 2 3)) (test (->listof (->listof ->nat*) (cons l123 (cons l123 null))) => '((1 2 3) (1 2 3))) ;; Recursive functions (define Y (lambda (f) ((lambda (x) (x x)) (lambda (x) (f (x x)))))) (rewrite (define/rec f E) => (define f (Y (lambda (f) E)))) ;; note that this example is doing something silly now (define/rec length (lambda (l) (if (null? l) 0 (add1 (length (cdr l)))))) (test (->nat* (length l123)) => '3) ;; addition becomes hard since it requires a recursive definition ;; (define/rec + ;; (lambda (m n) (if (zero? n) m (+ (add1 m) (sub1 n))))) ;; (test (->nat* (+ 4 5)) => '9) ;; faster alternative: (define/rec + (lambda (m n) (if (zero? m) n (if (zero? n) m (add1 (add1 (+ (sub1 m) (sub1 n)))))))) (test (->nat* (+ 4 5)) => '9) ;; subtraction is similar to addition ;; (define/rec - ;; (lambda (m n) (if (zero? n) m (- (sub1 m) (sub1 n))))) ;; (test (->nat* (- (+ 4 5) 4)) => '5) ;; but this is not "natural subtraction": doesn't work when n>m, ;; because (sub1 0) does not return 0. ;; a solution is like alternative form of +: (define/rec - (lambda (m n) (if (zero? m) 0 (if (zero? n) m (- (sub1 m) (sub1 n)))))) (test (->nat* (- (+ 4 5) 4)) => '5) (test (->nat* (- 2 5)) => '0) ;; alternatively, could change sub1 above: ;; (define sub1 (lambda (n) (if (zero? n) n (cdr n)))) ;; we can do multiplication in a similar way (define/rec * (lambda (m n) (if (zero? m) 0 (+ n (* (sub1 m) n))))) (test (->nat* (* 4 5)) => '20) (test (->nat* (+ 4 (* (+ 2 5) 5))) => '39) ;; and the rest of the examples (define/rec fact (lambda (x) (if (zero? x) 1 (* x (fact (sub1 x)))))) (test (->nat* (fact 5)) => '120) (define/rec fib (lambda (x) (if (or (zero? x) (zero? (sub1 x))) 1 (+ (fib (sub1 x)) (fib (sub1 (sub1 x))))))) (test (->nat* (fib (* 5 2))) => '89) #| ;; Fully-expanded Fibonacci (note: much shorter than the previous ;; encoding, but see how Y appears twice -- two "((lambda" pairs) (define fib((lambda(f)((lambda(x)(x x))(lambda(x)(f(x x)))))(lambda( f)(lambda(x)(((((x(lambda(x y)x))(x(lambda(x y)x)))((x(lambda(x y)y) )(lambda(x y)x)))(lambda(s)(s(lambda(x y)y)(lambda(x)x))))((((lambda (f)((lambda(x)(x x))(lambda(x)(f(x x))))) (lambda(f)(lambda(m n)((m( lambda(x y)x))n (((n(lambda(x y)x)) m)(lambda(s)((s (lambda(x y)y))( lambda(s)((s (lambda(x y)y))((f(m(lambda(x y)y)))(n(lambda(x y)y)))) ))))))))(f(x(lambda(x y)y))))(f((x(lambda(x y)y))(lambda(x y)y)))))) ))) |# Another interesting way to implement lists follows the pattern matching approach, where both pairs and the null value are represented by a function that serves as a kind of a `match` dispatcher. This function takes in two inputs --- if it is the representation of null then it will return the first input, and if it is a pair, then it will apply the second input on the two parts of the pair. This is implemented as follows (with type comments to make it clear): ;; null : List (define null (lambda (n p) n)) ;; cons : A List -> List (define cons (lambda (x y) (lambda (n p) (p x y)))) This might seem awkward, but it follows the intended use of pairs and null as a match-like construct. Here is an example, with the equivalent Racket code on the side: ;; Sums up a list of numbers (define (sum l) (l ; (match l 0 ; ['() 0] (lambda (x xs) ; [(cons x xs) (+ x (sum xs))))) ; (+ x (sum xs))]) In fact, it's easy to implement our selectors and predicate using this: (define null? (lambda (l) (l #t (lambda (x xs) #f)))) (define car (lambda (l) (l #f (lambda (x y) x)))) (define cdr (lambda (l) (l #f (lambda (x y) y)))) ;; in the above `#f' is really any value, since it ;; should be an error alternatively: (define car (lambda (l) (l ((lambda (x) (x x)) (lambda (x) (x x))) ; "error" (lambda (x y) x)))) The same approach can be used to define any kind of new data type in a way that looks like our own `define-type` definitions. For example, consider a much-simplified definition of the AE type we've seen early in the semester, and a matching `eval` definition as an example for using `cases`: (define-type AE [Num Number] [Add AE AE]) (: eval : AE -> Number) (define (eval expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))])) We can follow the above approach now to write Schlac code that more than being equivalent, is also very similar in nature. Note that the type definition is replaced by two definitions for the two constructors: (define Num (lambda (n) (lambda (num add) (num n )))) (define Add (lambda (l r) (lambda (num add) (add l r)))) (define/rec eval (lambda (expr) ; `expr` is always a (lambda (num add) ...), and it ; expects a unary `num` argument and a binary `add` (expr (lambda (n) n) (lambda (l r) (+ (eval l) (eval r)))))) (test (->nat (eval (Add (Num 1) (Num 2)))) => '3) We can even take this further: the translations from `define-type` and `cases` are mechanical enough that we could implement them almost exactly via rewrites (there are a subtle change in that we're list field names rather than types): (rewrite (define-type _ignored_ [Variant arg ...] ...) => (define Variant (lambda (arg ...) (lambda (Variant ...) (Variant arg ...)))) ...) (rewrite (cases value [(Variant arg ...) result] ...) => (value (lambda (arg ...) result) ...)) And using that, an evluator is simple: (define-type AE [Num n] [Add l r] [Sub l r] [Mul l r]) (define/rec eval (lambda (expr) (cases expr [(Num n) n] [(Add l r) (+ (eval l) (eval r))] [(Sub l r) (- (eval l) (eval r))] [(Mul l r) (* (eval l) (eval r))]))) (test (->nat (eval (Mul (Add (Num 1) (Num 2)) (Sub (Num 4) (Num 2))))) => '6) ------------------------------------------------------------------------ # Recursive Environments [Tuesday, February 18th] > [PLAI §11.5] What we really need for recursion, is a special kind of an environment, one that can refer to itself. So instead of doing (note: `call`s removed for readability): {with {fact {fun {n} {if {zero? n} 1 {* n {fact {- n 1}}}}}} {fact 5}} which does not work for the usual reasons, we want to use some {rec {fact {fun {n} {if {zero? n} 1 {* n {fact {- n 1}}}}}} {fact 5}} that will do the necessary magic. One way to achieve this is using the Y combinator as we have seen --- a kind of a "constructor" for recursive functions. We can do that in a similar way to the `rewrite` rule that we have seen in Schlac --- translate the above expression to: {with {fact {make-rec {fun {fact} {fun {n} {if {zero? n} 1 {* n {fact {- n 1}}}}}}}} {fact 5}} or even: {with {fact {{fun {f} {{fun {x} {f {x x}}} {fun {x} {f {x x}}}}} {fun {fact} {fun {n} {if {zero? n} 1 {* n {fact {- n 1}}}}}}}} {fact 5}} Now, we will see how it can be used in *our* code to implement a recursive environment. If we look at what `with` does in {with {fact {fun {n} {if {zero? n} 1 {* n {call fact {- n 1}}}}}} {call fact 5}} then we can say that to evaluate this expression, we evaluate the body expression in an extended environment that contains `fact`, even if a bogus one that is good for `0` only --- the new environment is created with something like this: extend("fact", make-fact-closure(), env) so we can take this whole thing as an operation over `env` add-fact(env) := extend("fact", make-fact-closure(), env) This gives us the first-level fact. But `fact` itself is still undefined in `env`, so it cannot call itself. We can try this: add-fact(add-fact(env)) but that still doesn't work, and it will never work no matter how far we go: add-fact(add-fact(add-fact(add-fact(add-fact(...env...))))) What we really want is infinity: a place where add-fact works and the result is the same as what we've started with --- we want to create a "magical" environment that makes this possible: let magic-env = ??? such that: add-fact(magic-env) = magic-env which basically gives us the illusion of being at the infinity point. This magic-env thing is exactly the *fixed-point* of the `add-fact` operation. We can use: magic-env = rec(add-fact) and following the main property of the Y combinator, we know that: magic-env = rec(add-fact) ; def. of magic-env = add-fact(rec(add-fact)) ; Y(f) = f(Y(f)) = add-fact(magic-env) ; def. of magic-env What does all this mean? It means that if we have a fixed-point operator at the level of the implementation of our environments, then we can use it to implement a recursive binder. In our case, this means that a fixpoint in Racket can be used to implement a recursive language. But we have that --- Racket does have recursive functions, so we should be able to use that to implement our recursive binder. There are two ways that make it possible to write recursive functions in Racket. One is to define a function, and use its name to do a recursive call --- using the Racket formal rules, we can see that we said that we mark that we now *know* that a variable is bound to a value. This is essentially a side-effect --- we modify what we know, which corresponds to modifying the global environment. The second way is a new form: `letrec`. This form is similar to `let`, except that the scope that is established includes the named expressions --- it is exactly what we want `rec` to do. A third way is using recursive local definitions, but that is equivalent to using `letrec`, more on this soon. ------------------------------------------------------------------------ # Recursion: Racket's `letrec` [Tuesday, February 18th] So we want to add recursion to our language, practically. We already know that Racket makes it possible to write recursive functions, which is possible because of the way it implements its "global environment": our evaluator can only *extend* an environment, while Racket *modifies* its global environment. This means that whenever a function is defined in the global environment, the resulting closure will have it as its environment "pointer", but the global environment was not extended --- it stays the same, and was just modified with one additional binding. But Racket has another, a bit more organized way of using recursion: there is a special local-binding construct that is similar to `let`, but allows a function to refer to itself. It is called `letrec`: (letrec ([fact (lambda (n) (if (zero? n) 1 (* n (fact (- n 1)))))]) (fact 5)) Some people may remember that there was a third way for creating recursive functions: using local definition in function bodies. For example, we have seen things like: (define (length list) (define (helper list len) (if (null? list) len (helper (rest list) (+ len 1)))) (helper list 0)) This looks like the same kind of environment magic that happens with a global `define` --- but actually, Racket defines the meaning of internal definitions using `letrec` --- so the above code is exactly the same as: (define (length list) (letrec ([helper (lambda (list len) (if (null? list) len (helper (rest list) (+ len 1))))]) (helper list 0))) The scoping rules for a `letrec` is that the scope of the bound name covers both the body *and* the named expression. Furthermore, multiple names can be bound to multiple expressions, and the scope of each name covers all named expression as well as the body. This makes it easy to define mutually recursive functions, such as: (letrec ([even? (lambda (n) (if (zero? n) #t (odd? (- n 1))))] [odd? (lambda (n) (if (zero? n) #f (even? (- n 1))))]) (even? 99)) But it is not a required functionality --- it could be done with a single recursive binding that contains several functions: (letrec ([even+odd (list (lambda (n) (if (zero? n) #t ((second even+odd) (- n 1)))) (lambda (n) (if (zero? n) #f ((first even+odd) (- n 1)))))]) ((first even+odd) 99)) This is basically the same problem we face if we want to use the Y combinator for mutually recursive bindings. The above solution is inconvenient, but it can be improved using more `let`s to have easier name access. For example: (letrec ([even+odd (list (lambda (n) (let ([even? (first even+odd)] [odd? (second even+odd)]) (if (zero? n) #t (odd? (- n 1))))) (lambda (n) (let ([even? (first even+odd)] [odd? (second even+odd)]) (if (zero? n) #f (even? (- n 1))))))]) (let ([even? (first even+odd)] [odd? (second even+odd)]) (even? 99))) ------------------------------------------------------------------------ # Implementing Recursion using `letrec` [Tuesday, February 18th] We will see how to add a similar construct to our language --- for simplicity, we will add a `rec` form that handles a single binding: {rec {fact {fun {n} {if {= 0 n} 1 {* n {fact {- n 1}}}}}} {fact 5}} Using this, things can get a little tricky. What should we get if we do: {rec {x x} x} ? Currently, it seems like there is no point in using any expression except for a *function* expression in a `rec` expression, so we will handle only these cases. (BTW, under what circumstances would non-function values be useful in a letrec?) ------------------------------------------------------------------------ One way to achieve this is to use the same trick that we have recently seen: instead of re-implementing language features, we can use existing features in our own language, which hopefully has the right functionality in a form that can be re-used to in our evaluator. Previously, we have seen a way to implement environments using Racket closures: ;; Define a type for functional environments (define-type ENV = Symbol -> VAL) (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error 'lookup "no binding for ~s" id))) (: lookup : Symbol ENV -> VAL) (define (lookup name env) (env name)) (: Extend : Symbol VAL ENV -> ENV) (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val (rest-env name)))) We can use this implementation, and create circular environments using Racket's `letrec`. The code for handling a `with` expressions is: [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] It looks like we should be able to handle `rec` in a similar way (the AST constructor name is `WRec` ("with-rec") so it doesn't collide with TR's `Rec` constructor for recursive types): [(WRec bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] but this won't work because the named expression is evaluated prematurely, in the previous environment. Instead, we will move everything that needs to be done, including evaluation, to a separate `extend-rec` function: [(WRec bound-id named-expr bound-body) (eval bound-body (extend-rec bound-id named-expr env))] Now, the `extend-rec` function needs to provide the new, "magically circular" environment. Following what we know about the arguments to `extend-rec`, and the fact that it returns a new environment (= a lookup function), we can sketch a rough definition: (: extend-rec : Symbol FLANG ENV -> ENV) ; FLANG, not VAL! ;; extend an environment with a new binding that is the result of ;; evaluating an expression in the same environment as the extended ;; result (define (extend-rec id expr rest-env) (lambda (name) (if (eq? name id) ... something that uses expr to get a value ... (rest-env name)))) What should the missing expression be? It can simply evaluate the object given itself: (define (extend-rec id expr rest-env) (lambda (name) (if (eq? name id) (eval expr ...this environment...) (rest-env name)))) But how do we get *this environment*, before it is defined? Well, the environment is itself a Racket *function*, so we can use Racket's `letrec` to make the function refer to itself recursively: (define (extend-rec id expr rest-env) (letrec ([rec-env (lambda (name) (if (eq? name id) (eval expr rec-env) (rest-env name)))]) rec-env)) It's a little more convenient to use an internal definition, and add a type for clarity: (define (extend-rec id expr rest-env) (: rec-env : Symbol -> VAL) (define (rec-env name) (if (eq? name id) (eval expr rec-env) (rest-env name))) rec-env) This works, but there are several problems: 1. First, we no longer do a simple lookup in the new environment. Instead, we evaluate the expression on *every* such lookup. This seems like a technical point, because we do not have side-effects in our language (also because we said that we want to handle only function expressions). Still, it wastes space since each evaluation will allocate a new closure. 2. Second, a related problem --- what happens if we try to run this: {rec {x x} x} ? Well, we do that stuff to extend the current environment, then evaluate the body in the new environment, this body is a single variable reference: (eval (Id 'x) the-new-env) so we look up the value: (lookup 'x the-new-env) which is: (the-new-env 'x) which goes into the function which implements this environment, there we see that `name` is the same as `name1`, so we return: (eval expr rec-env) but the `expr` here is the original named-expression which is itself `(Id 'x)`, and we're in an infinite loop. We can try to get over these problems using another binding. Racket allows several bindings in a single `letrec` expression or multiple internal function definitions, so we change `extend-rec` to use the newly-created environment: (define (extend-rec id expr rest-env) (: rec-env : Symbol -> VAL) (define (rec-env name) (if (eq? name id) val (rest-env name))) (: val : VAL) (define val (eval expr rec-env)) rec-env) This runs into an interesting type error, which complains about possibly getting some `Undefined` value. It does work if we switch to the untyped language for now (using `#lang pl untyped`) --- and it seems to run fine too. But it raises more questions, beginning with: what is the meaning of: (letrec ([x ...] [y ...x...]) ...) or equivalently, an internal block of (define x ...) (define y ...x...) ? Well, DrRacket seems to do the "right thing" in this case, but what about: (letrec ([y ...x...] [x ...]) ...) ? As a hint, see what happens when we now try to evaluate the problematic {rec {x x} x} expression, and compare that with the result that you'd get from Racket. This also clarifies the type error that we received. It should be clear now why we want to restrict usage to just binding recursive functions. There are no problems with such definitions because when we evaluate a `fun` expression, there is no evaluation of the body, which is the only place where there are potential references to the same function that is defined --- a function's body is *delayed*, and executed only when the function is applied later. But the biggest question that is still open: we just implemented a circular environment using Racket's own circular environment implementation, and that does not explain how they are actually implemented. The cycle of pointers that we've implemented depends on the cycle of pointers that Racket uses, and that is a black box we want to open up. For reference, the complete code is below. #lang pl #| The grammar: ::= | { + } | { - } | { * } | { / } | { with { } } | { rec { } } | | { fun { } } | { call } Evaluation rules: eval(N,env) = N eval({+ E1 E2},env) = eval(E1,env) + eval(E2,env) eval({- E1 E2},env) = eval(E1,env) - eval(E2,env) eval({* E1 E2},env) = eval(E1,env) * eval(E2,env) eval({/ E1 E2},env) = eval(E1,env) / eval(E2,env) eval(x,env) = lookup(x,env) eval({with {x E1} E2},env) = eval(E2,extend(x,eval(E1,env),env)) eval({rec {x E1} E2},env) = ??? eval({fun {x} E},env) = <{fun {x} E}, env> eval({call E1 E2},env1) = eval(Ef,extend(x,eval(E2,env1),env2)) if eval(E1,env1) = <{fun {x} Ef}, env2> = error! otherwise |# (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [WRec Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons 'with more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `with' syntax in ~s" sexpr)])] [(cons 'rec more) (match sexpr [(list 'rec (list (symbol: name) named) body) (WRec name (parse-sexpr named) (parse-sexpr body))] [else (error 'parse-sexpr "bad `rec' syntax in ~s" sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; Types for environments, values, and a lookup function (define-type VAL [NumV Number] [FunV Symbol FLANG ENV]) ;; Define a type for functional environments (define-type ENV = Symbol -> VAL) (: EmptyEnv : -> ENV) (define (EmptyEnv) (lambda (id) (error 'lookup "no binding for ~s" id))) (: lookup : Symbol ENV -> VAL) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (env name)) (: Extend : Symbol VAL ENV -> ENV) ;; extend a given environment cache with a new binding (define (Extend id val rest-env) (lambda (name) (if (eq? name id) val (rest-env name)))) (: extend-rec : Symbol FLANG ENV -> ENV) ;; extend an environment with a new binding that is the result of ;; evaluating an expression in the same environment as the extended ;; result (define (extend-rec id expr rest-env) (: rec-env : Symbol -> VAL) (define (rec-env name) (if (eq? name id) val (rest-env name))) (: val : VAL) (define val (eval expr rec-env)) rec-env) (: NumV->number : VAL -> Number) ;; convert a FLANG runtime numeric value to a Racket one (define (NumV->number val) (cases val [(NumV n) n] [else (error 'arith-op "expected a number, got: ~s" val)])) (: arith-op : (Number Number -> Number) VAL VAL -> VAL) ;; gets a Racket numeric binary operator, and uses it within a NumV ;; wrapper (define (arith-op op val1 val2) (NumV (op (NumV->number val1) (NumV->number val2)))) (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) (NumV n)] [(Add l r) (arith-op + (eval l env) (eval r env))] [(Sub l r) (arith-op - (eval l env) (eval r env))] [(Mul l r) (arith-op * (eval l env) (eval r env))] [(Div l r) (arith-op / (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (eval named-expr env) env))] [(WRec bound-id named-expr bound-body) (eval bound-body (extend-rec bound-id named-expr env))] [(Id name) (lookup name env)] [(Fun bound-id bound-body) (FunV bound-id bound-body env)] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV bound-id bound-body f-env) (eval bound-body (Extend bound-id (eval arg-expr env) f-env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) (EmptyEnv))]) (cases result [(NumV n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => 7) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Implementing `rec` Using Cyclic Structures [Tuesday, February 18th] > [PLAI §10] Looking at the arrows in the environment diagrams, what we're really looking for is a closure that has an environment pointer which is the same environment in which it was defined. This will make it possible for `fact` to be bound to a closure that can refer to *itself* since its environment is the same one in which it is defined. However, so far we have no tools that makes it possible to do this. What we need is to create a "cycle of pointers", and so far we do not have a way of achieving that: when we create a closure, we begin with an environment which is saved in the slot's environment slot, but we want that closure to be the value of a binding in that same environment. ------------------------------------------------------------------------ # Boxes and Mutation [Tuesday, February 18th] To actually implement a circular structure, we will now use *side-effects*, using a new kind of Racket value which supports mutation: a box. A box value is built with the `box` constructor: (define my-thing (box 7)) the value is retrieved with the `unbox' function, (* 6 (unbox my-thing)) and finally, the value can be changed with the `set-box!` function. (set-box! my-thing 17) (* 6 (unbox my-thing)) An important thing to note is that `set-box!` is much like `display` etc, it returns a value that is not printed in the Racket REPL, because there is no point in using the result of a `set-box!`, it is called for the side-effect it generates. (Languages like C blur this distinction between returning a value and a side-effect with its assignment statement.) As a side note, we now have side effects of two kinds: mutation of state, and I/O (at least the O part). (Actually, there is also infinite looping that can be viewed as another form of a side effect.) This means that we're now in a completely different world, and lots of new things can make sense now. A few things that you should know about: * We never used more than one expression in a function body because there was no point in it, but now there is. To evaluate a sequence of Racket expressions, you wrap them in a `begin` expression. * In most places you don't actually need to use `begin` --- these are places that are said to have an *implicit* `begin`: the body of a function (or any lambda expression), the body of a `let` (and `let`-relatives), the consequence positions in `cond`, `match`, and `cases` clauses and more. One of the common places where a `begin` is used is in an `if` expression (and some people prefer using `cond` instead when there is more than a single expression). * `cond` without an `else` in the end can make sense, if all you're using it it for is side-effects. * `if` could get a single expression which is executed when the condition is true (and an unspecified value is used otherwise), but our language (as well as the default Racket language) always forbids this --- there are convenient special forms for a one-sided `if`s: `when` & `unless`, and they can have any number of expressions (they have an implicit `begin`). They have an advantage of saying "this code does some side-effects here" more explicit. * There is a function called `for-each` which is just like `map`, except that it doesn't collect the list of results, it is used only for performing side effects. When any one of these things is used (in Racket or other languages), you can tell that side-effects are involved, because there is no point in any of them otherwise. In addition, any name that ends with a `!` ("bang") is used to mark a function that changes state (usually a function that only changes state). So how do we create a cycle? Simple, boxes can have any value, and they can be put in other values like lists, so we can do this: #lang pl untyped (define foo (list 1 (box 3))) (set-box! (second foo) foo) and we get a circular value. (Note how it is printed.) And with types: #lang pl (: foo : (List Number (Boxof Any))) (define foo (list 1 (box 3))) (set-box! (second foo) foo) ------------------------------------------------------------------------ # Types for Boxes [Tuesday, February 18th] Obviously, `Any` is not too great --- it is the most generic type, so it provides the least information. For example, notice that (unbox (second foo)) returns the right list, which is equal to `foo` itself --- but if we try to grab some part of the resulting list: (second (unbox (second foo))) we get a type error, because the result of the `unbox` is `Any`, so Typed Racket knows nothing about it, and won't allow you to treat it as a list. It is not too surprising that the type constructor that can help in this case is `Rec` which we have already seen --- it allows a type that can refer to itself: #lang pl (: foo : (Rec this (List Number (Boxof (U #f this))))) (define foo (list 1 (box #f))) (set-box! (second foo) foo) Note that either `foo` or the value in the box are both printed with a `Rec` type --- the value in the box can't just have a `(U #f this)` type, since `this` doesn't mean anything in there, so the whole type needs to still be present. There is another issue to be aware of with `Boxof` types. For most type constructors (like `Listof`), if `T1` is a subtype of `T2`, then we also know that`(Listof T1)` is a subtype of `(Listof T2)`. This makes the following code typecheck: #lang pl (: foo : (Listof Number) -> Number) (define (foo l) (first l)) (: bar : Integer -> Number) (define (bar x) (foo (list x))) Since the `(Listof Integer)` is a subtype of the `(Listof Number)` input for `foo`, the application typechecks. But this is *not* the same for the output type, for example --- if we change the `bar` type to: (: bar : Integer -> Integer) we get a type error since `Number` is not a subtype of `Integer`. So subtypes are required to "go up" on the input side and "down" on the other. So, in a sense, the fact that boxes are mutable means that their contents can be considered to be on the other side of the arrow, which is why for such `T1` subtype of `T2`, it is `(Boxof T2)` that is a subtype of `(Boxof T1)`, instead of the usual. For example, this doesn't work: #lang pl (: foo : (Boxof Number) -> Number) (define (foo b) (unbox b)) (: bar : Integer -> Number) (define (bar x) (: b : (Boxof Integer)) (define b (box x)) (foo b)) And you can see why this is the case --- the marked line is fine given a `Number` contents, so if the type checker allows passing in a box holding an integer, then that expression would mutate the contents and make it an invalid value. However, boxes are not only mutable, they hold a value that can be read too, which means that they're on *both* sides of the arrow, and this means that `(Boxof T1)` is a subtype of `(Boxof T2)` if `T2` is a subtype of `T1` *and* `T1` is a subtype of `T2` --- in other words, this happens only when `T1` and `T2` are the same type. (See below for an extended demonstration of all of this.) Note also that this demonstration requires that extra `b` definition, if it's skipped: (define (bar x) (foo (box x))) then this will typecheck again --- Typed Racket will just consider the context that requires a box holding a `Number`, and it is still fine to initialize such a box with an `Integer` value. > As a side comment, this didn't always work. Earlier in its existence, > Typed Racket would always choose a specific type for values, which > would lead to confusing errors with boxes. For example, the above > would need to be written as > > (define (bar x) > (foo (box (ann x : Number)))) > > to prevent Typed Racket from inferring a specific type. This is no > longer the case, but there can still be some surprises. A similar > annotation was needed in the case of a list holding a self-referential > box, to avoid the initial `#f` from getting a specific-but-wrong type. ------------------------------------------------------------------------ # `Boxof`'s Lack of Subtyping [Tuesday, February 18th] The lack of any subtype relations between `(Boxof T)` and `(Boxof S)` regardless of `S` and `T` can roughly be explained as follows. First, a box is a container that you can pull a value out of --- which makes it similar to lists. In the case of lists, we have: if: S subtype-of T then: (Listof S) subtype-of (Listof T) This is true for all such containers that you can pull a value out of: if you expect to pull a `T` but you're given a container of a subtype `S`, then things are still fine. Such "containers" include functions that produce a value --- for example: if: S subtype-of T then: Q -> S subtype-of Q -> T However, functions also have the other side, where things are different --- instead of a side of some *produced* value, it's the side of the *consumed* value. We get the opposite rule there: if: T subtype-of S then: S -> Q subtype-of T -> Q To see why this is right, use `Number` and `Integer` for `S` and `T`: if: Integer subtype-of Number then: Number -> Q subtype-of Integer -> Q so --- if you expect a function that takes a number is a *subtype* of one that takes an integer; in other words, every function that takes a number is also a function that takes an integer, but not the other way. To summarize all of this, when you make the output type of a function "smaller" (more constrained), the resulting type is smaller, but on the input side things are flipped --- a bigger input type means a more constrained function. Now, a `(Boxof T)` is a producer of `T` when you pull a value out of the box, but it's also a consumer of `T` when you put such a value in it. This means that --- using the above analogy --- the `T` is on both sides of the arrow. This means that if: S subtype-of T *and* T subtype-of S then: (Boxof S) subtype-of (Boxof T) which is actually: if: S is-the-same-type-as T then: (Boxof S) is-the-same-type-as (Boxof T) A different way to look at this conclusion is to consider the function type of `(A -> A)`: when is it a subtype of some other `(B -> B)`? Only when `A` is a subtype of `B` and `B` is a subtype of `A`, which means that this happens only when `A` and `B` are the same type. (Side note: this is related to the fact that in logic, `P => Q` is roughly equivalent to `not(P) or Q` --- the left side, `P`, is inside a negation. It also explains why in `((S -> T) -> Q)` the `S` obeys the first rule, as if it was on the right side --- because it's negated twice.) The following piece of code makes the analogy to function types more formally. Boxes behave as if their contents is on both sides of a function arrow --- on the right because they're readable, and on the left because they're writable, which the conclusion that a `(Boxof A)` type is a subtype of itself and no other `(Boxof B)`. #lang pl ;; a type for a "read-only" box (define-type (Boxof/R A) = (-> A)) ;; Boxof/R constructor (: box/r : (All (A) A -> (Boxof/R A))) (define (box/r x) (lambda () x)) ;; we can see that (Boxof/R T1) is a subtype of (Boxof/R T2) ;; if T1 is a subtype of T2 (this is not surprising, since ;; these boxes are similar to any other container, like lists): (: foo1 : Integer -> (Boxof/R Integer)) (define (foo1 b) (box/r b)) (: bar1 : (Boxof/R Number) -> Number) (define (bar1 b) (b)) (test (bar1 (foo1 123)) => 123) ;; a type for a "write-only" box (define-type (Boxof/W A) = (A -> Void)) ;; Boxof/W constructor (: box/w : (All (A) A -> (Boxof/W A))) (define (box/w x) (lambda (new) (set! x new))) ;; in contrast to the above, (Boxof/W T1) is a subtype of ;; (Boxof/W T2) if T2 is a subtype of T1, *not* the other way ;; (and note how this is related to A being on the *left* side ;; of the arrow in the `Boxof/W' type): (: foo2 : Number -> (Boxof/W Number)) (define (foo2 b) (box/w b)) (: bar2 : (Boxof/W Integer) Integer -> Void) (define (bar2 b new) (b new)) (test (bar2 (foo2 123) 456)) ;; combining the above two into a type for a "read/write" box (define-type (Boxof/RW A) = (A -> A)) ;; Boxof/RW constructor (: box/rw : (All (A) A -> (Boxof/RW A))) (define (box/rw x) (lambda (new) (let ([old x]) (set! x new) old))) ;; this combines the above two: `A' appears on both sides of the ;; arrow, so (Boxof/RW T1) is a subtype of (Boxof/RW T2) if T1 ;; is a subtype of T2 (because there's an A on the right) *and* ;; if T2 is a subtype of T1 (because there's another A on the ;; left) -- and that can happen only when T1 and T2 are the same ;; type. So this is a type error: ;; (: foo3 : Integer -> (Boxof/RW Integer)) ;; (define (foo3 b) (box/rw b)) ;; (: bar3 : (Boxof/RW Number) Number -> Number) ;; (define (bar3 b new) (b new)) ;; (test (bar3 (foo3 123) 456) => 123) ;; ** Expected (Number -> Number), but got (Integer -> Integer) ;; And this a type error too: ;; (: foo3 : Number -> (Boxof/RW Number)) ;; (define (foo3 b) (box/rw b)) ;; (: bar3 : (Boxof/RW Integer) Integer -> Integer) ;; (define (bar3 b new) (b new)) ;; (test (bar3 (foo3 123) 456) => 123) ;; ** Expected (Integer -> Integer), but got (Number -> Number) ;; The two types must be the same for this to work: (: foo3 : Integer -> (Boxof/RW Integer)) (define (foo3 b) (box/rw b)) (: bar3 : (Boxof/RW Integer) Integer -> Integer) (define (bar3 b new) (b new)) (test (bar3 (foo3 123) 456) => 123) ------------------------------------------------------------------------ # Implementing a Circular Environment [Tuesday, February 18th] We now use this to implement `rec` in the following way: 1. Change environments so that instead of values they hold boxes of values: `(Boxof VAL)` instead of `VAL`, and whenever `lookup` is used, the resulting boxed value is unboxed, 2. In the `WRec` case, create the new environment with some temporary binding for the identifier --- any value will do since it should not be used (when named expressions are always `fun` expressions), 3. Evaluate the expression in the new environment, 4. Change the binding of the identifier (the box) to the result of this evaluation. The resulting definition is: (: extend-rec : Symbol FLANG ENV -> ENV) ;; extend an environment with a new binding that is the result of ;; evaluating an expression in the same environment as the extended ;; result (define (extend-rec id expr rest-env) (let ([new-cell (box (NumV 42))]) (let ([new-env (Extend id new-cell rest-env)]) (let ([value (eval expr new-env)]) (set-box! new-cell value) new-env)))) Racket has another `let` relative for such cases of multiple-nested `let`s --- `let*`. This form is a derived form --- it is defined as a shorthand for using nested `let`s. The above is therefore exactly the same as this code: (: extend-rec : Symbol FLANG ENV -> ENV) ;; extend an environment with a new binding that is the result of ;; evaluating an expression in the same environment as the extended ;; result (define (extend-rec id expr rest-env) (let* ([new-cell (box (NumV 42))] [new-env (Extend id new-cell rest-env)] [value (eval expr new-env)]) (set-box! new-cell value) new-env)) This `let*` form can be read almost as a C/Java-ish kind of code: fun extend_rec(id, expr, rest_env) { new_cell = new NumV(42); new_env = Extend(id, new_cell, rest_env); value = eval(expr, new_env); *new_cell = value; return new_env; } The code can be simpler if we fold the evaluation into the `set-box!` (since `value` is used just there), and if use `lookup` to do the mutation --- since this way there is no need to hold onto the box. This is a bit more expensive, but since the binding is guaranteed to be the first one in the environment, the addition is just one quick step. The only binding that we need is the one for the new environment, which we can do as an internal definition, leaving us with: (: extend-rec : Symbol FLANG ENV -> ENV) (define (extend-rec id expr rest-env) (define new-env (Extend id (box (NumV 42)) rest-env)) (set-box! (lookup id new-env) (eval expr new-env)) new-env) A complete rehacked version of FLANG with a `rec` binding follows: ;;; ---<<>>---------------------------------------------- #lang pl (define-type FLANG [Num Number] [Add FLANG FLANG] [Sub FLANG FLANG] [Mul FLANG FLANG] [Div FLANG FLANG] [Id Symbol] [With Symbol FLANG FLANG] [WRec Symbol FLANG FLANG] [Fun Symbol FLANG] [Call FLANG FLANG]) (: parse-sexpr : Sexpr -> FLANG) ;; parses s-expressions into FLANGs (define (parse-sexpr sexpr) (match sexpr [(number: n) (Num n)] [(symbol: name) (Id name)] [(cons (or 'with 'rec) more) (match sexpr [(list 'with (list (symbol: name) named) body) (With name (parse-sexpr named) (parse-sexpr body))] [(list 'rec (list (symbol: name) named) body) (WRec name (parse-sexpr named) (parse-sexpr body))] [(cons x more) (error 'parse-sexpr "bad `~s' syntax in ~s" x sexpr)])] [(cons 'fun more) (match sexpr [(list 'fun (list (symbol: name)) body) (Fun name (parse-sexpr body))] [else (error 'parse-sexpr "bad `fun' syntax in ~s" sexpr)])] [(list '+ lhs rhs) (Add (parse-sexpr lhs) (parse-sexpr rhs))] [(list '- lhs rhs) (Sub (parse-sexpr lhs) (parse-sexpr rhs))] [(list '* lhs rhs) (Mul (parse-sexpr lhs) (parse-sexpr rhs))] [(list '/ lhs rhs) (Div (parse-sexpr lhs) (parse-sexpr rhs))] [(list 'call fun arg) (Call (parse-sexpr fun) (parse-sexpr arg))] [else (error 'parse-sexpr "bad syntax in ~s" sexpr)])) (: parse : String -> FLANG) ;; parses a string containing a FLANG expression to a FLANG AST (define (parse str) (parse-sexpr (string->sexpr str))) ;; Types for environments, values, and a lookup function (define-type ENV [EmptyEnv] [Extend Symbol (Boxof VAL) ENV]) (define-type VAL [NumV Number] [FunV Symbol FLANG ENV]) (: lookup : Symbol ENV -> (Boxof VAL)) ;; lookup a symbol in an environment, return its value or throw an ;; error if it isn't bound (define (lookup name env) (cases env [(EmptyEnv) (error 'lookup "no binding for ~s" name)] [(Extend id boxed-val rest-env) (if (eq? id name) boxed-val (lookup name rest-env))])) (: extend-rec : Symbol FLANG ENV -> ENV) ;; extend an environment with a new binding that is the result of ;; evaluating an expression in the same environment as the extended ;; result (define (extend-rec id expr rest-env) (define new-env (Extend id (box (NumV 42)) rest-env)) (set-box! (lookup id new-env) (eval expr new-env)) new-env) (: NumV->number : VAL -> Number) ;; convert a FLANG runtime numeric value to a Racket one (define (NumV->number val) (cases val [(NumV n) n] [else (error 'arith-op "expected a number, got: ~s" val)])) (: arith-op : (Number Number -> Number) VAL VAL -> VAL) ;; gets a Racket numeric binary operator, and uses it within a NumV ;; wrapper (define (arith-op op val1 val2) (NumV (op (NumV->number val1) (NumV->number val2)))) (: eval : FLANG ENV -> VAL) ;; evaluates FLANG expressions by reducing them to values (define (eval expr env) (cases expr [(Num n) (NumV n)] [(Add l r) (arith-op + (eval l env) (eval r env))] [(Sub l r) (arith-op - (eval l env) (eval r env))] [(Mul l r) (arith-op * (eval l env) (eval r env))] [(Div l r) (arith-op / (eval l env) (eval r env))] [(With bound-id named-expr bound-body) (eval bound-body (Extend bound-id (box (eval named-expr env)) env))] [(WRec bound-id named-expr bound-body) (eval bound-body (extend-rec bound-id named-expr env))] [(Id name) (unbox (lookup name env))] [(Fun bound-id bound-body) (FunV bound-id bound-body env)] [(Call fun-expr arg-expr) (let ([fval (eval fun-expr env)]) (cases fval [(FunV bound-id bound-body f-env) (eval bound-body (Extend bound-id (box (eval arg-expr env)) f-env))] [else (error 'eval "`call' expects a function, got: ~s" fval)]))])) (: run : String -> Number) ;; evaluate a FLANG program contained in a string (define (run str) (let ([result (eval (parse str) (EmptyEnv))]) (cases result [(NumV n) n] [else (error 'run "evaluation returned a non-number: ~s" result)]))) ;; tests (test (run "{call {fun {x} {+ x 1}} 4}") => 5) (test (run "{with {add3 {fun {x} {+ x 3}}} {call add3 1}}") => 4) (test (run "{with {add3 {fun {x} {+ x 3}}} {with {add1 {fun {x} {+ x 1}}} {with {x 3} {call add1 {call add3 x}}}}}") => 7) (test (run "{with {identity {fun {x} x}} {with {foo {fun {x} {+ x 1}}} {call {call identity foo} 123}}}") => 124) (test (run "{with {x 3} {with {f {fun {y} {+ x y}}} {with {x 5} {call f 4}}}}") => 7) (test (run "{call {with {x 3} {fun {y} {+ x y}}} 4}") => 7) (test (run "{with {f {with {x 3} {fun {y} {+ x y}}}} {with {x 100} {call f 4}}}") => 7) (test (run "{call {call {fun {x} {call x 1}} {fun {x} {fun {y} {+ x y}}}} 123}") => 124) ------------------------------------------------------------------------ # Variable Mutation [Tuesday, February 18th] > [PLAI §12] and [PLAI §13] (different: adds boxes to the language) > > [PLAI §14] (that's what we do) The code that we now have implements recursion by *changing* bindings, and to make that possible we made environments hold boxes for all bindings, therefore bindings are *all* mutable now. We can use this to add more functionality to our evaluator, by allowing changing any variable --- we can add a `set!` form: {set! } to the evaluator that will modify the value of a variable. To implement this functionality, all we need to do is to use `lookup` to retrieve some box, then evaluate the expression and put the result in that box. The actual implementation is left as a homework exercise. One thing that should be considered here is --- all of the expressions in our language evaluate to some value, the question is what should be the value of a `set!` expression? There are three obvious choices: 1. return some bogus value, 2. return the value that was assigned, 3. return the value that was previously in the box. Each one of these has its own advantage --- for example, C uses the second option to `chain` assignments (eg, `x = y = 0`) and to allow side effects where an expression is expected (eg, `while (x = x-1) ...`). The third one is useful in cases where you might use the old value that is overwritten --- for example, if C had this behavior, you could `pop` a value from a linked list using something like: first(stack = rest(stack)); because the argument to `first` will be the old value of `stack`, before it changed to be its `rest`. You could also swap two variables in a single expression: `x = y = x`. (Note that the expression `x = x + 1` has the meaning of C's `++x` when option (2) is used, and `x++` when option (3) is used.) Racket chooses the first option, and we will do the same in our language. The advantage here is that you get no discounts, therefore you must be explicit about what values you want to return in situations where there is no obvious choice. This leads to more robust programs since you do not get other programmers that will rely on a feature of your code that you did not plan on. In any case, the modification that introduces mutation is small, but it has a tremendous effect on our language: it was true for Racket, and it is true for FLANG. We have seen how mutation affects the language subset that we use, and in the extension of our FLANG the effect is even stronger: since *any* variable can change (no need for explicit `box`es). In other words, a binding is not always the same --- in can change as a result of a `set!` expression. Of course, we could extend our language with boxes (using Racket boxes to implement FLANG boxes), but that will be a little more verbose. > Note that Racket does have a `set!` form, and in addition, fields in > structs can be made modifiable. However, we do not use any of these. > At least not for now. ------------------------------------------------------------------------ # State and Environments [Tuesday, February 18th] A quick example of how mutation can be used: (define counter (let ([counter (box 0)]) (lambda () (set-box! counter (+ 1 (unbox counter))) (unbox counter)))) and compare that to: (define (make-counter) (let ([counter (box 0)]) (lambda () (set-box! counter (+ 1 (unbox counter))) (unbox counter)))) It is a good idea if you follow the exact evaluation of (define foo (make-counter)) (define bar (make-counter)) and see how both bindings have separate environment so each one gets its own private state. The equivalent code in the homework interpreter extended with `set!` doesn't need boxes: {with {make-counter {fun {} {with {counter 0} {fun {} {set! counter {+ counter 1}} counter}}}} {with {foo {call make-counter}} {with {bar {call make-counter}} ...}}} (To see multiple values from a single expression you can extend the language with a `list` binding.) Note that we cannot describe this behavior with substitution rules! We now use the environments to make it possible to change bindings --- so finally an environment is actually an environment rather than a substitution cache. When you look at the above, note that we still use lexical scope --- in fact, the local binding is actually a private state that nobody can access. For example, if we write this: (define counter (let ([counter (box 0)]) (lambda () (set-box! counter (+ 1 (unbox counter))) (if (zero? (modulo (unbox counter) 4)) 'tock 'tick)))) then the resulting function that us bound to `counter` keeps a local integer state which no other code can access --- you cannot modify it, reset it, or even know if it is really an integer that is used in there. ------------------------------------------------------------------------ # Implementing Objects with State [Tuesday, February 18th] We have already seen how several pieces of information can be encapsulate in a Racket closure that keeps them all; now we can do a little more --- we can actually have mutable state, which leads to a natural way to implement objects. For example: (define (make-point x y) (let ([xb (box x)] [yb (box y)]) (lambda (msg) (match msg ['getx (unbox xb)] ['gety (unbox yb)] ['incx (set-box! xb (add1 (unbox xb)))])))) implements a constructor for `point` objects which keep two values and can move one of them. Note that the messages act as a form of methods, and that the values themselves are hidden and are accessible only through the interface that these messages make. For example, if these points correspond to some graphic object on the screen, we can easily incorporate a necessary screen update: (define (make-point x y) (let ([xb (box x)] [yb (box y)]) (lambda (msg) (match msg ['getx (unbox xb)] ['gety (unbox yb)] ['incx (set-box! xb (add1 (unbox xb))) (update-screen)])))) and be sure that this is always done when the value changes --- since there is no way to change the value except through this interface. A more complete example would define functions that actually send these messages --- here is a better implementation of a point object and the corresponding accessors and mutators: (define (make-point x y) (let ([xb (box x)] [yb (box y)]) (lambda (msg) (match msg ['getx (unbox xb)] ['gety (unbox yb)] [(list 'setx newx) (set-box! xb newx) (update-screen)] [(list 'sety newy) (set-box! yb newy) (update-screen)])))) (define (point-x p) (p 'getx)) (define (point-y p) (p 'gety)) (define (set-point-x! p x) (p (list 'setx x))) (define (set-point-y! p y) (p (list 'sety y))) And a quick imitation of inheritance can be achieved using delegation to an instance of the super-class: (define (make-colored-point x y color) (let ([p (make-point x y)]) (lambda (msg) (match msg ['getcolor color] [else (p msg)])))) You can see how all of these could come from some preprocessing of a more normal-looking class definition form, like: (defclass point (x y) (public (getx) x) (public (gety) y) (public (setx new) (set! x newx)) (public (setx new) (set! x newx))) (defclass colored-point point (c) (public (getcolor) c))