Scheme: An Interpreter for Extended Lambda Calculus

SCHEME

扩展 Lambda 演算的解释器

Gerald Jay Sussman 和 Guy Lewis Steele Jr.

摘要：

受 ACTORS [Greif and Hewitt] [Smith and Hewitt] 启发，我们实现了一个类似 LISP 的编程语言 SCHEME 的解释器。该编程语言基于 Lambda 演算 [Church]，但扩展了副作用、多进程以及进程同步机制。该实现的目的在于教学。我们希望：

(1) 通过澄清像 LISP 这样的递归宿主语言中非递归控制结构的嵌入方式，减轻由 Micro-PLANNER、CONNIVER 等造成的困惑。

(2) 解释如何使用这些控制结构，而不涉及模式匹配和数据库操作等问题。

(3) 为编程语义和风格中的某些问题提供一个简单具体的实验领域。

本文组织为若干部分。第一部分是一个简短的“参考手册”，包含 SCHEME 所有特殊特性的规范。接下来，我们给出一系列程序示例，用以说明各种程序设计风格以及如何使用它们。这将引出一些语义问题，我们将在第三部分尝试用 Lambda 演算对其进行澄清。在第四部分，我们将对基于 Lambda 演算的编程语言的解释器实现者所面临的问题进行总体讨论。最后，我们将给出一个完全注释的 SCHEME 解释器，该解释器使用 MacLISP [Moon] 编写，以使程序员了解在类似 LISP 这样的递归语言中实现非递归控制结构的技巧。

MASSACHUSETTS INSTITUTE OF TECHNOLOGY ARTIFICIAL INTELLIGENCE LABORATORY

AI Memo No. 349

December 1975

SCHEME

AN INTERPRETER FOR EXTENDED LAMBDA CALCULUS by

Gerald Jay Sussman and Guy Lewis Steele Jr.

Abstract: Inspired by ACTORS [Greif and Hewitt] [Smith and Hewitt], we have implemented an interpreter for a LISP-like language, SCHEME, based on the lambda calculus [Church], but extended for side effects, multiprocessing, and process synchronization. The purpose of this implementation is tutorial. We wish to:

(1) alleviate the confusion caused by Micro-PLANNER, CONNIVER, etc. by clarifying the embedding of non-recursive control structures in a recursive host language like LISP. (2) explain how to use these control structures, independent of such issues as pattern matching and data base manipulation. (3) have a simple concrete experimental domain for certain issues of programming semantics and style.

This paper is organized into sections. The first section is a short "reference manual" containing specifications for all the unusual features of SCHEME. Next, we present a sequence of programming examples which illustrate various programming styles, and how to use them. This will raise certain issues of semantics which we will try to clarify with lambda calculus in the third section. In the fourth section we will give a general discussion of the issues facing an implementor of an interpreter for a language based on lambda calculus. Finally, we will present a completely annotated interpreter for SCHEME, written in MacLISP [Moon], to acquaint programmers with the tricks of the trade of implementing non-recursive control structures in a recursive language like LISP.

This report describes research done at the Artificial Intelligence Laboratory of the Massachusetts Institute of Technology. Support for the laboratory's artificial intelligence research is provided in part by the Advanced Research Projects Agency of the Department of Defense under Office of Naval Research contract N00014-75-C-0643.

Section 1: The SCHEME Reference Manual

SCHEME is essentially a full-funarg LISP. LAMBDA expressions need not be QUOTEd, FUNCTIONed, or *FUNCTIONed when passed as arguments or returned as values; they will evaluate to closures of themselves.

All LISP functions (i.e., EXPRS, SUBRs, and LSUBRs, but not FEXPRS, FSUBRs, or MACROS) are primitive operators in SCHEME, and have the same meaning as they have in LISP. Like LAMBDA expressions, primitive operators and numbers are self-evaluating (they evaluate to trivial closures of themselves).

There are a number of special primitives known as AINTs which are to SCHEME as FSUBRs are to LISP. We will enumerate them here.

IF This is the primitive conditional operator. It takes three arguments. If the first evaluates to non-NIL, it evaluates the second expression, and otherwise the third.

QUOTE As in LISP, this quotes the argument form so that it will be passed verbatim as data. The abbreviation "'FOO" may be used instead of "(QUOTE FOO)".

DEFINE This is analogous to the MacLISP DEFUN primitive (but note that the LAMBDA must appear explicitly!). It is used for defining a function in the "global environment" permanently, as opposed to LABELS (see below), which is used for temporary definitions in a local environment. DEFINE takes a name and a lambda expression; it closes the lambda expression in the global environment and stores the closure in the LISP value cell of the name (which is a LISP atom).

LABELS We have decided not to use the traditional LABEL primitive in this interpreter because it is difficult to define several mutually recursive functions using only LABEL. The solution, which Hewitt [Smith and Hewitt] also uses, is to adopt an ALGOLesque block syntax:

(LABELS <function definition list> <expression>)

This has the effect of evaluating the expression in an environment where all the functions are defined as specified by the definitions list. Furthermore, the functions are themselves closed in that environment, and not in the outer environment; this allows the functions to call themselves and each other recursively. For example, consider a function which counts all the atoms in a list structure recursively to all levels, but which doesn't count the NILs which terminate lists (but NILs in the CAR of some list count). In order to perform this we use two mutually recursive functions, one to count the car and one to count the cdr, as follows:

(DEFINE COUNT (LAMBDA (L) (LABELS ((COUNTCAR (LAMBDA (L) (IF (ATOM L) 1 (+ (COUNTCAR (CAR L)) (COUNTCDR (CDR L))))))) (COUNTCDR (LAMBDA (L) (IF (ATOM L) (IF (NULL L) 0 1) (+ (COUNTCAR (CAR L)) (COUNTCDR (CDR L))))))) (COUNTCDR L))))) ;Note: COUNTCDR is defined here.

ASET This is the side effect primitive. It is analogous to the LISP function SET. For example, to define a cell [Smith and Hewitt], we may use ASET as follows:

(DEFINE CONS-CELL (LAMBDA (CONTENTS) (LABELS ((THE-CELL (LAMBDA (MSG) (IF (EQ MSG 'CONTENTS?) CONTENTS (IF '(EQ MSG 'CELL?) 'YES (IF (EQ (CAR MSG) '<-) (BLOCK (ASET 'CONTENTS (CADR MSG)) THE-CELL) (ERROR '|UNRECOGNIZED MESSAGE - CELL| MSG 'WRNG-TYPE-ARG)))))))) THE-CELL)))

Those of you who may complain about the lack of ASETQ are invited to write (ASET' foo bar) instead of (ASET 'foo bar).

EVALUATE

This is similar to the LISP function EVAL. It evaluates its argument, and then evaluates the resulting s-expression as SCHEME code.

CATCH This is the "escape operator" which gives the user a handle on the control structure of the interpreter. The expression:

(CATCH <identifier> <expression>)

This is the synchronization primitive. It evaluates an expression uninterruptibly; i.e. no other process may run until the expression has returned a value. Note that if a funarg is returned from the scope of an EVALUATE!UNINTERRUPTIBLY, then that funarg will be uninterruptible when it is applied; that is, the uninterruptibility property follows the rules of variable scoping. For example, consider the following function:

(DEFINE SEMGEN (LAMBDA (SEMVAL) (LIST (LAMBDR () (EVALUATE!UNINTERRUPTIBLY (ASET' SEMVAL (+ SEMVAL 1))))) (LABELS (P (LAMBDRA () (EVALUATE!UNINTERRUPTIBLY (IF (PLUSP SEMVAL) (ASET' SEMVAL (- SEMVAL 1)) (P))))))) P)))))

This returns a pair of functions which are V and P operations on a newly created semaphore. The argument to SEMGEN is the initial value for the semaphore. Note that P busy-waits by iterating if necessary; because EVALUATE!UNINTERRUPTIBLY uses variable-scoping rules, other processes have a chance to get in at the beginning of each iteration. This busy-wait can be made much more efficient by replacing the expression (P) in the definition of P with

((LAMBDRA (ME) (BLOCK (START!PROCESS (CREATE!PROCESS '(START!PROCESS ME))) (STOP!PROCESS ME) (P)))) PROCESS

Let's see you figure this one out! Note that a STOP!PROCESS within an EVALUATE!UNINTERRUPTIBLY forces the process to be swapped out even if it is the current one, and so other processes get to run; but as soon as it gets swapped in again, others are locked out as before.

Besides the AINTs, SCHEME has a class of primitives known as AMACROs. These are similar to MacLISP MACROs, in that they are expanded into equivalent code before being executed. Some AMACROs supplied with the SCHEME interpreter:

COND This is like the MacLISP COND statement, except that singleton clauses (where the result of the predicate is the returned value) are not allowed.

AND, OR These are also as in MacLISP.

BLOCK This is like the MacLISP PROGN, but arranges to evaluate its last argument without an extra net control frame (explained later), so that the last argument may involved in an iteration. Note that in SCHEME, unlike MacLISP, the body of a LAMBDA expression is not an implicit PROGN.

DO This is like the MacLISP "new-style" DO; old-style DO is not supported.

AMAPCAR, AMAPLIST These are like MAPCAR and MAPLIST, but they expect a SCHEME lambda closure for the first argument.

To use SCHEME, simply incant at DDT (on MIT-AI):

:LISP LIBLSP;SCHEME which will load up the current version of SCHEME, which will announce itself and give a prompt. If you want to escape to LISP, merely hit ^G. To restart SCHEME, type (SCHEME). Sometimes one does need to use a LISP FSUBR such as UREAD; this may be accomplished by typing, for example,

(EVAL' (UREAD FOO BAR DSK LOSER))

After doing this, typing ^Q will, of course, cause SCHEME to read from the file.

This concludes the SCHEME Reference Manual.

(DEFINE FACT (LAMBDA (N C) (IF (= N 0) (C 1) (FACT (- N 1) (LAMBOA (A) (C (* N A))))))))

Note that we can call this like an ordinary function if we supply (LAMBDA (X) X) as the second argument. For example, (FACT 3 (LAMBDA (X) X)) returns 6.

Apparently "Hairy" Control Structure

A classic problem difficult to solve in most programming languages, including standard (stack-oriented) LISP, is the samefringe problem [Smith and Hewitt]. The problem is to determine whether the fringes of two trees are the same, even if the internal structures of the trees are not. This problem is easy to solve if one merely computes the fringe of each tree separately as a list, and then compares the two lists. We would like to solve the problem so that the fringes are generated and compared incrementally. This is important if the fringes of the trees are very large, but differ, say, in the first position.

Consider the following obscure solution to samefringe, which is in fact isomorphic to the one written by Shrobe and presented by Smith and Hewitt. Note that SCHEME does not have the packagers of PLASMA, and so we were forced to use continuations; rather than using packages and a next operator, we pass a fringe a continuation (called the "getter") which will get the next and the rest of the fringe as its two arguments.

(DEFINE FRINGE (LAMBDA (TREE) (LABELS ((FRINGEN (LAMBDA (NODE ALT) (LAMBDA (GETTER) (IF (ATOM NODE) (GETTER NODE ALT) ((FRINGEN (CAR NODE) (LAMBDA (GETTER1) ((FRINGEN (CDR NODE) ALT) GETTER1))))) GETTER))))) (FRINGEN TREE (LAMBDA (GETTER) (GETTER '(EXHAUSTED) NIL)))))))

(DEFINE SAMEFRINGE (LAMBDA (TREE1 TREE2) (LABELS ((SAME (LAMBDA (S1 S2) (S1 (LAMBDA (X1 R1) (S2 (LAMBDA (X2 R2) (IF (EQUAL X1 X2) (IF (EQUAL X1 '(EXHAUSTED)) T (SAME R1 R2)) NIL))))))))) (SAME (FRINGE TREE1) (FRINGE TREE2)))))

Now let us consider an alternative solution to the samefringe problem. We believe that this solution is clearer for two reasons: (1) the implementation of SAMEFRINGE is more clearly iterative; (2) rather than returning an object which will return both the first and the rest of a fringe to a given continuation, FRINGE returns an object which will deliver up a component in response to a request for that component.

(DEFINE FRINGE (LAMBDA (TREE) (LABELS ((FRINGE1 (LAMBDA (NODE ALT) (IF (ATOM NODE) (LAMBDA (MSG) (IF (EQ MSG 'FIRST) NODE (IF (EQ MSG 'NEXT) (ALT) (ERROR)))) (FRINGE1 (CAR NODE) (LAMBDA () (FRINGE1 (CDR NODE) ALT))))))) (FRINGE1 TREE (LAMBDA () (LAMBDA (MSG) (IF (EQ MSG 'FIRST) 'EOF (ERROR))))))))

(DEFINE SAMEFRINGE (LAMBDA (T1 T2) (DO (((C1 (FRINGE T1) (C1 'NEXT)) (C2 (FRINGE T2) (C2 'NEXT))) ((OR (NOT (EQ (C1 'FIRST) (C2 'FIRST))) (EQ (C1 'FIRST) 'EOF) (EQ (C2 'FIRST) 'EOF)) (EQ (C1 'FIRST) (C2 'FIRST))))))

A much simpler and more probable problem is that of building a pattern matcher with backtracking for segment matches. The matcher presented below is intended for matching single-level list structure patterns against lists of atoms. A pattern is a list containing three types of elements:

(1) constant atoms, which match themselves only. (2) (THV x), which matches any single element in the expression consistently. We may abbreviate this as ?x by means of a LISP reader macro character. (3) (THV* x), which matches any segment of zero or more elements in the expression consistently. We may abbreviate this as !x. The matcher returns either NIL, meaning no match is possible, or a list of two items, an alist specifying the bindings of the match variables, and a continuation to call, if you don't like this particular set of bindings, which will attempt to find another match. Thus, for example, the invocation

(MATCH '(A !B ?C ?C !B !E) '(A X Y Q Q X Y Z Z X Y Q Q X Y R)) would return the result

(((E (Z Z X Y Q Q X Y R)) (C Q) (B X Y)) <continuation1>) where calling <continuation1> as a function of no arguments would produce the result

(((E (R)) (C Z) (B (X Y Q Q X Y))) <continuation2>) where calling <continuation2> would produce NIL.

The MATCH function makes use of two auxiliary functions called NFIRST and NREST. The former returns a list of the first n elements of a given list, while the latter returns the tail of the given list after the first n elements.

(DEFINE NFIRST (LAMBDA (E N) (IF (= N 0) NIL (CONS (CAR E) (NFIRST (CDR E) (- N 1))))))

(DEFINE NREST (LAMBDA (E N) (IF (= N 0) E (NREST (CDR E) (- N 1)))))

The main MATCH function also uses a subfunction called MATCH1 which takes four arguments: the tail of the pattern yet to be matched; the tail of the expression yet to be matched; the alist of match bindings made so far; and a continuation to call if the match fails at this point. A subfunction of MATCH, called MATCH, handles the matching of segments of the expression against THV match variables. It is in the matching of segments that the potential need for backtracking enters, for segments of various lengths may have to be tried. After MATCH* matches a segment, it calls MATCH1 to continue the match, giving it a failure continuation which will back up and try to match a longer segment if possible. A failure can occur if a constant fails to match, or if one or the other of pattern and expression runs out before the other one does.

(DEFINE MATCH (LAMBDA (PATTERN EXPRESSION) (LABELS ((MATCH1 (LAMBDA (P E ALIST LOSE) (IF (NULL P) (IF (NULL E) (LIST ALIST LOSE) (LOSE)) (IF (ATOM (CAR P)) (IF (NULL E) (LOSE) (IF (EQ (CAR E) (CAR P)) (MATCH1 (CDR P) (CDR E) ALIST LOSE) (LOSE)))) (IF (EQ (CAAR P) 'THV) (IF (NULL E) (LOSE) ((LAMBDA (V) (IF V (IF (EQ (CAR E) (CADR V)) (MATCH1 (CDR P) (CDR E) ALIST LOSE) (LOSE))) (MATCH1 (CDR P) (CDR E) (CONS (LIST (CADAR P) (CAR E)) ALIST) LOSE)))) (ASSQ (CADAR P) ALIST)))) (IF (EQ (CAAR P) 'THV) ((LAMBDA (V) (IF V (IF (< (LENGTH E) (LENGTH (CADR V))) (LOSE) (IF (EQUAL (NFIRST E (LENGTH (CADR V))) (CADR V)) (MATCH1 (CDR P) (NREST E (LENGTH (CADR V))) ALIST LOSE) (LOSE)))) (LABELS ((MATCH (LAMBDA (N) (IF (> N (LENGTH E)) (LOSE) (MATCH1 (CDR P) (NREST E N) (CONS (LIST (CADAR P) (NFIRST E N)) ALIST) (LAMBDA () (MATCH* (+ N 1))))))))) (ASSQ (CADAR P) ALIST)) (LOSE))))))))) (MATCH1 PATTERN EXPRESSION NIL (LAMBDA () NIL))))))

A Useless Multiprocessing Example

One thing we might want to use multiprocessing for is to try two things in parallel, and terminate as soon as one succeeds. We can do this with the following function.

(DEFINE TRY!TWO!THINGS!IN!PARALLEL (LAMBDA (F1 F2) (CATCH C ((LAMBDA (P1 P2) ((LAMBDA (F1 F2) (EVALUATE!UNINTERRUPTIBLY (BLOCK (ASET 'P1 (CREATE!PROCESS '(F1))) (ASET 'P2 (CREATE!PROCESS '(F2))) (START!PROCESS P1) (START!PROCESS P2) (STOP!PROCESS PROCESS)))) (LAMBDA () ((LAMBDA (VALUE) (EVALUATE!UNINTERRUPTIBLY (BLOCK (STOP!PROCESS P2) (C VALUE)))) (F1)))) (LAMBDA () ((LAMBDA (VALUE) (EVALUATE!UNINTERRUPTIBLY (BLOCK (STOP!PROCESS P1) (C VALUE)))) (F2))))) NIL NIL))))

TRY!TWO!THINGS!IN!PARALLEL takes two functions of no arguments (in order to pass an unevaluated expression and its environment in for later use, so as to avoid variable conflicts). It creates two processes to run them, and returns the value of whichever completes first.

As an example of how to misuse TRY!TWO!THINGS!IN!PARALLEL, here is a function which determines the sign of an integer using only ADD1, SUB1, and EQUAL.

(DEFINE SIGN (LAMBDA (N) (IF (EQUAL N 0) 'ZERO (TRY!TWO!THINGS!IN!PARALLEL (LAMBDA () (DO ((I 0 (ADD1 I))) ((EQUAL I N) 'POSITIVE))) (LAMBDA () (DO ((I 0 (SUB1 I))) ((EQUAL I N) 'NEGATIVE)))))))

Recursive programs

Now, let's see how lambda calculus may be used (informally) to model a computation. Consider the standard definition of the factorial function:

(DEFINE FACT (LAMBDA (N) (IF (= N 0) 1 (* N (FACT (- N 1))))))

We are being very informal -- lambda calculus as presented by [Church] does not include such constructs as DEFINE, IF, or =, , or even 1! The "usual" lambda calculus construct for defining recursive functions is a rather obscure object called the "fixed-point" operator. We have been lax to avoid the hassle of "rigor mortis" in this tutorial paper. Similarly, IF is the SCHEME conditional construct we will use for convenience, it reduces to its second or third argument depending on whether the first reduces to TRUE or FALSE. The objects , =, 0, 1, etc. may be thought of as abbreviations for complex lambda expressions (such as Church numerals) whose details we are not interested in. On the other hand, we may think of them as primitive expressions, defined by additional axioms; this viewpoint leads to practical interpreter implementations.

Now let's reduce the expression (FACT 3). We will perform the expression reductions, except for the IF primitive, in Applicative Order (call by value), though this is not necessary, as we will discuss later. We display a "trace" of the substitutions:

            
              =>   (FACT 3)
=>   (IF (= 3 0) 1 (* 3 (FACT (- 3 1))))
=>   (* 3 (FACT (- 3 1)))
=>   (* 3 (FACT 2))
=>   (* 3 (IF (= 2 0) 1 (* 2 (FACT (- 2 1)))))
=>   (* 3 (* 2 (FACT (- 2 1))))
=>   (* 3 (* 2 (FACT 1)))
=>   (* 3 (* 2 (IF (= 1 0) 1 (* 1 (FACT (- 1 1)))))) 
=>   (* 3 (* 2 (* 1 (FACT (- 1 1)))))
=>   (* 3 (* 2 (* 1 (FACT 0))))
=>   (* 3 (* 2 (* 1 (IF (= 0 0) 1 (* 0 (FACT (- 0 1)))))))
=>   (* 3 (* 2 (* 1 1)))
=>   (* 3 (* 2 1))
=>   (* 3 2)
=>   6
            
          

You will note that we have calculated (fact 3) by a process wherein each expression is replaced by an expression which is provably equivalent to it via an axiom or which is produced by application of a primitive function.

Now, What About Iteration?

Consider the "iterative" definition of FACT. Although it appears to be recursive, since it "calls itself", we will see that it captures the essence of our intuitive notion of iteration.

(DEFINE FACT (LAMBDA (N) (LABELS ((FACT1 (LAMBDA (M ANS) (IF (= M 0) ANS (FACT1 (- M 1) (* M ANS))))))) (FACT1 N 1))))

Let us now compute (fact 3).

            
              =>   (FACT 3)
=>   (FACT1 3 1)
=>   (IF (= 3 0) 1
        (FACT1 (- 3 1) (* 3 1)))
=>   (FACT1 (- 3 1) (* 3 1))
=>   (FACT1 2 (* 3 1))
=>   (FACT1 2 3)
=>   (IF (= 2 0) 3
        (FACT1 (- 2 1) (* 2 3)))
=>   (FACT1 (- 2 1) (* 2 3))
=>   (FACT1 1 (* 2 3))
=>   (FACT1 1 6)
=>   (IF (= 1 0) 6
        (FACT1 (- 1 1) (* 1 6)))
=>   (FACT1 (- 1 1) (* 1 6))
=>   (FACT1 0 (* 1 6))
=>   (FACT1 0 6)
=>   (IF (= 0 0) 6
        (FACT1 (- 0 1) (* 0 6)))
=>   6
            
          

Notice that the expressions involved have a fixed maximum size independent of the argument to FACT! In fact, as Marvin Minsky pointed out, successive reductions produce a cycle of expressions which are identical except for the numerical quantities involved. Looking back, we may note by way of comparison that the recursive version caused creation of expressions proportional in size to the argument. This is why we think that this version of FACT is iterative rather than recursive. At each stage of the iterative version the "state" of the computation is summarized in two variables, the counter and the answer accumulator, while at each stage of the recursive version the "state" contains a chain of pieces each of which contains a component of the state. In the recursive version of FACT, for example, the state contains the sequence of multiplications to be performed upon return from the bottom. It is true that the iterative factorial also can produce expressions of arbitrary size, since the number of bits needed to express factorial of n grows with n; but this is a property of the numbers calculated by the function which is implemented in iterative style, and not of the iterative control structure itself. A recursive control structure inherently creates expressions of unbounded size as a function of the recursion depth, while an iterative control structure produces a cycle of equivalent expressions, and so the expressions are of approximately the same size no matter how many iteration steps are taken. This is the essence of the difference between the notions of iteration and recursion. Hewitt [MAC, p. 234] made a similar observation in passing, expressing the difference in terms of storage used in program execution rather than in terms of intermediate expressions produced by substitution semantics.

Continuation Passing Recursion

Remember the other way to compute factorials?

(DEFINE FACT (LAMBDA (N C) (IF (= N 0) (C 1) (FACT (- N 1) (LAMBDA (A) (C (* N A)))))))

This looks iterative on the surface! but in fact it is recursive. Let's compute (FACT 3 ANSWER), where ANSWER is a continuation which is to receive the result of FACT applied to 3; that is, the last thing FACT should do is apply the continuation ANSWER to its result.

=> (FACT 3 ANSWER) => (IF (= 3 0) (ANSWER 1) (FACT (- 3 1) (LAMBDA (A) (ANSWER ( 3 A))))) => (FACT (- 3 1) (LAMBDA (A) (ANSWER ( 3 A)))) => (FACT 2 (LAMBDA (A) (ANSWER ( 3 A)))) => (IF (= 2 0) ((LAMBDA (A) (ANSWER ( 3 A))) 1) (FACT (- 2 1) (LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A)))))) => (FACT (- 2 1) (LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A)))) => (FACT 1 (LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))))

=> (IF (= 1 0) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A)))) 1) (FACT (- 1 1) (LAMBDA (A) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) ( 1 A))))) => (FACT (- 1 1) (LAMBDA (A) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) ( 1 A))))) => (FACT 0 (LAMBDA (A) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) ( 1 A))))) => (IF (= 0 0) ((LAMBDA (A) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) ( 1 A)))) 1) (FACT (- 0 1) (LAMBDA (A) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) ( 1 A))) ( 0 A))))) => ((LAMBDA (A) ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) (* 1 A))))

=> ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) ( 1 1)) => ((LAMBDA (A) ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 A))) 1) => ((LAMBDA (A) (ANSWER ( 3 A))) ( 2 1)) => ((LAMBDA (A) (ANSWER ( 3 A))) 2) => (ANSWER (* 3 2)) => (ANSWER 6) Whew!

Note that we have computed factorial of 3 (and are about to give this result to the continuation), but in the process no combination with FACT in the first position has ever been reduced except as the outermost expression. If we think of the computation in terms of evaluation rather than substitution, this means that we never returned a value from any application of the function FACT! It is always possible, if we are willing to specify explicitly what to do with the answer, to perform any calculation in this way: rather than reducing to its value, it reduces to an application of a continuation to its value (cf. [Fischer]). That is, in this continuation-passing programming style, a function always "returns" its result by "sending" it to another function. This is the key idea.

We also note that by our previous observation, this program is essentially recursive in that the expressions produced as intermediate results of the substitution semantics grow to a size proportional to the depth. In fact, the same information is being stored in the nested continuations produced by this program as in the nested products produced by the traditional recursion -- what to do with the result.

One might object that this FACT is not the same kind of object as the previous definition, since we can't use it as a function in the same manner. Note, however, that if we supply the continuation (LAMBDA (X) X), the resulting combination (FACT 3 (LAMBDA (X) X)) will reduce to 6, just as with traditional recursion.

One might also object that we are using function values -- the primitives =, -, and are functions which return values, for example. But this is just a property of the primitives; consider a new set of primitives ==, --, and * which accept continuations (indeed, let == take two continuations: if the predicate is TRUE call the first, otherwise call the second). We can then define fact as follows:

(DEFINE FACT (LAMBDA (N C) (= N 0 (LAMBDA () (C 1)) (LAMBDA () (-- N 1 (LAMBDA (M) (FACT M (LAMBDA (A) (** A N C)))))))))

We can see here that no functional application returns a value in a computation of factorial in this situation. We believe that functional usage, where applicable (pun intended), is more perspicuous than continuation-passing. We shall therefore use functional primitives such as * rather than wherever possible, keeping in mind that we could use instead if we wished.

(DEFINE CONS (LAMBDA (A B) (LAMBDA (M) (IF (EQ M 'FIRST?) A (IF (EQ M 'REST?) B Y (IF (EQ M 'LIST?) 'YES (ERROR '|UNRECOGNIZED MESSAGE - CONS| M 'WRNG-TYPE-ARG)))))))

Note that we have used explicit eq tests on the incoming message rather than the implicit pattern-matching of the cases statement, but the underlying semantics of the approach are the same.

There are several important consequences of closing every lambda expression in the environment from which it is passed (i.e., in its "lexical" or "static" environment). First, the axioms of lambda calculus are automatically preserved. Thus, referential transparency is enforced. This in turn implies that there are no "fluid" variable bindings (as there are in standard stack implementations of LISP such as MacLISP). Second, the upward funarg problem [Moses] requires that the environment structure be potentially tree-like. Finally, the environment at any point in a computation can never be deeper than the lexical depth of the expression being evaluated at that time; i.e., the environment contains bindings only for variables bound in lambdas lexically surrounding the expression being evaluated. This is true even if recursive functions are involved. It follows that if list structures are used to implement environments, the time to look up a variable is proportional only to the lexical distance from the reference to the binding and not to the depth of recursion or any other dynamic parameter. Therefore it is not necessarily as expensive as many people have been led to believe. Furthermore, it is not even necessary to scan the environment for the variable, since its value must be in a known position relative to the top of the environment structure; this position can be computed by a compiler at compile time on the basis of lexical scope. The tree-like structure of an environment prevents one from merely indexing into the it, however; it is necessary to cdr down it. (Some ALGOL compilers use a similar technique involving base registers pointing to sets of variables for each level of block nesting; it is necessary to determine the base pointer for the block desired for a variable reference, but then the variable is at a known offset from the base address.) It also follows that an iterative programming style will lead to no net accumulation of environment structures in the interpreter. The recursive style, with or without continuation-passing, will lead to accumulation of environment structures as a function of the recursion depth, not because any environment becomes arbitrarily deep, but rather because at each level of recursion it is necessary to save the environment at that point. It is saved by the interpreter in the case of traditional recursion, so that computation can continue in the correct environment on return from the recursive call; it is saved as part of the continuation closure in continuation-passing.

Another problem is concerned with control. This is a consequence of the functional interpretation of the lambda calculus, i.e. the view that a "expression" (combination) represents a value to be "returned" (to replace the combination) to its "caller" (the process evaluating the combination containing the original one). The interpreter must provide for correctly resuming the caller when the callee has returned its value. The state of the computation at the time of the call must therefore be preserved. We see, then, that part of the state of the computation must be (a pointer to) the preserved state of its caller; we will call this component of the state the clink [McDermott and Sussman] [Bobrow and Wegbreit]. Just before the evaluation of a subexpression, the state of the current computation, including the clink, must be gathered together into a single data structure, which we will call a frame; the clink is then altered to point to this new frame. The evaluation of the subexpression then returns by restoring the state of the process from the current clink. Note that the value of the subexpression had better not be part of the state, for otherwise it would be lost by the state restoration. Thus, we only build a new frame if further computation would result in losing information which might be necessary. This only occurs if we must somehow return to that state. This in turn can only occur if we must evaluate an expression whose value must be obtained in order to continue computation in the current state.

This implies that no frame need be created in order to apply a lambda expression to its arguments. This in turn implies that the iterative and continuation-passing styles lead to no net creation of frames, because they are implemented only in terms of explicit lambda applications, whereas the recursive style leads to the creation of one net frame per level of recursive depth, because the recursive invocation involves the evaluation of a expression containing the recursive lambda application as a subexpression.

A clink in a lambda calculus-based interpreter is in fact equivalent to a low-level default continuation as created by the PLASMA interpreter. Such a continuation is a (closed) lambda expression of one argument whose script will carry on the computation after receiving the value of the subexpression. The clink mechanism is therefore not necessary, if we are willing to transform all our programs into pure continuation-passing style. We could do this explicitly, by requiring the user to write his programs in this form; or implicitly, as PLASMA does, by creating these one-argument continuations as necessary, passing them as hidden extra arguments to lambda expressions which behave like functions. On the other hand, we may think of a clink as a highly optimized continuation, whose "script" is that carefully coded portion of the lambda calculus interpreter which restores the frame and then carries on. We find this notion useful in defining a primitive, CATCH (named for the CATCH construct in MacLISP [Moon]), for "hairy control structure", similar to Reynolds' ESCAPE operator [Reynolds], which makes these low-level continuations available to the user. Note that PLASMA has a similar facility for getting hold of the low-level continuations, namely the ">>>" receiver construct.

Another problem for the implementor of an interpreter of a lambda calculus based language is the order in which to perform reductions. There are two standard orders of evaluation (and several other semi-standard ones, which we will not consider here). The first is Normal Order, which corresponds roughly to ALGOL's "call by name", and the second is Applicative Order, which corresponds roughly to ALGOL's "call by value" or to LISP functional application.

Under a call-by-name implementation, the <args>* mentioned above are in fact the actual argument expressions, each paired with the environment E. The evaluator has two additional rules: (1) when a variable x is to be evaluated in environment E1, then its associated expression-environment pair [A,E2] (which is equivalent to an ALGOL thunk) is looked up in E1, and then A is evaluated in E2. (2) when a "primitive operator" is to be applied, its arguments must be evaluated at that time, and then the operator applied in a call-by-value manner.

Under a call-by-value implementation, the <args>* are the values of the argument expressions; i.e., the argument expressions are evaluated in environment E, and only then is the lambda expression applied. Note that this leads to trouble in defining conditionals. Under call-by-name one may define predicates to return (LAMBDA (X Y) X) for TRUE and (LAMBDA (X Y) Y) for FALSE, and then one may simply write

((= A B) <do this if TRUE> <do this if FALSE>)

This trick depends implicitly on the order of evaluation. It will not work under call-by-value, nor in general under any other reductive order except Normal Order. It is therefore necessary to introduce a special primitive operator (such as "if") which is applied in a call-by-name manner. This leads us to the interesting conclusion that a practical lambda calculus interpreter cannot be purely call-by-name or call-by-value; it is necessary to have at least a little of each.

There is a fundamental problem, however, with using Normal Order evaluation in a lambda calculus interpreter, which is brought out by the iterative programming style. We already know that no net frames are created by iterative programs, and that no net environment structures are created either. The problem is that under a call-by-name implementation there may be a net thunk structure created proportional in size to the number of iteration steps. This problem is inherent in Normal Order, because Normal Order substitution semantics exhibit the same phenomenon of increasing expression size. Therefore iteration cannot be effectively modeled in a call-by-name interpreter. An alternative view is that a call-by-name interpreter remembers more than is logically necessary to perform the computations indicated by the original expressions. This is indicated by the fact that the Applicative Order substitution semantics lead to expressions of fixed maximum size independent of the number of iteration steps.

It turns out that this conflict between call-by-name and iteration is resolved by the use of continuation-passing. If we use a pure continuation-passing programming style, then Normal Order and Applicative Order are the same order! In pure continuation-passing no combination is ever a subcombination of another combination. (This is the justification for the fact mentioned above that no clinks are needed if pure continuation-passing style is used.) Thus, if we wish to model iteration in pure lambda calculus without even an if primitive, we can use Normal Order substitutions and express the iteration in the continuation-passing style.

Under any reductive order, whether Normal Order, Applicative Order, or any other order, it is in practice convenient to introduce a means of terminating the evaluation process on a given form; in order to do this we introduce three different and equally useful notions. The first is the primitive operator such as +; the evaluator can apply such an operator directly, without substituting a lambda expression for the operator and reducing the resulting form. The second is the self-evaluating constant; this is used for primitive objects such as numbers, which effectively behave as if always "bound to themselves" in any environment. The third is the quoting function, which protects its argument from reductions so that it is returned as is; this is used for treating forms as data in the usual LISP manner.

These three ideas are not logically necessary, since the evaluation process will (eventually) terminate when no reductions can be made, but they are a great convenience for introducing various functions and data into the lambda calculus. Note too that some are easily defined in terms of the others; for example, instead of letting 3 be a self-evaluating constant, we could let 3 be a primitive operator of no arguments which returned 3, or we could merely quote it; similarly, instead of quoting forms we could let forms be a self-evaluating data type, as in MDL [Galley and Pfister] (better known as MUDDLE), written with different parentheses. Because, as we have said, these constructs are all strictly for convenience, we will not strive for any kind of minimality, but will continue to use all three notions in our interpreter, as we already have in our examples. We provide an interface so that all MacLISP subrs may be used as primitive operators; we define numbers to be self-evaluating; and we will use QUOTE to quote forms as in LISP (and thus we may use the '" character as an abbreviation).

One final issue which the implementor of a lambda calculus based interpreter should consider is that of extensions to the language, such as primitives for side effects, multiprocessing, and synchronization of processes. Note that these are ideas which are very hard, if not impossible, to model using the substitution semantics of the lambda calculus, but which are easily incorporated in other semantic models, including the environment interpreter and, perhaps more notably, the ACTORS model [Greif and Hewitt]. The fundamental problem with modelling such concepts using substitution semantics is that substitution produces copies of expressions, and so cannot model the notion of sharing very well. In an interpreter which uses environments, all instances of a variable scoped in a given environment refer to the same virtual substitution contained in that environment, and so may be thought of as sharing a value cell in that environment. We can take advantage of this sharing by introducing a primitive operator which modifies the contents of a value cell; since all occurrences refer to the same value cell, changing the contents of that value cell will change the result of future references to that value cell (i.e., occurrences of the variable which invoke the virtual substitution mechanism). Such a primitive operator would then be similar to the SET function of LISP, or the := of ALGOL. We include such an operator, ASET, in our interpreter.

Introducing multiprocessing into the interpreter is fairly straightforward; all that is necessary is to introduce a mechanism for timeslicing the interpreter among several processes. One can even model this in substitution semantics by supposing that there can be more than one expression, and at each step an expression is randomly chosen to perform a reduction within. (On the other hand, synchronizing of the processes is very hard to model using substitution semantics!)

Since our value cells effectively solve the readers and writers problem (i.e. reads and writes of variables are indivisible) no more than our side effect primitive is necessary to synchronize our processes [Dijkstra] [Knuth] [Lamport]. However, the techniques for achieving synchronization using only := are quite cumbersome and opaque. It behooves the implementor to make things easier for the user by introducing a more tractable synchronization primitive (e.g. P+V or monitors or path expressions or ...). Machine language programmers have long known that the easiest way to synchronize processes is to turn off the scheduling clock during the execution of critical code. We have introduced such a primitive, EVALUATE!UNINTERRUPTIBLY, (which is a sort of "over-anxious serializer", because it locks out the whole world) into our interpreter.

Section 5: The Implementation of the Interpreter

Here we present a real live SCHEME interpreter. This particular version was written primarily for expository purposes; it works, but not as efficiently as possible. The "production version" of SCHEME is coded somewhat more intricately, and runs about twice as fast as the interpreter presented below.

The basic idea behind the implementation is think machine language. In particular, we must not use recursion in the implementation language to implement recursion in the language being interpreted. This is a crucial mistake which has screwed many language implementations (e.g. Micro-PLANNER [Sussman]). The reason for this is that if the implementation language does not support certain kinds of control structures, then we will not be able to effectively interpret them. Thus, for example, if the control frame structure in the implementation language is constrained to be stack-like, then modelling more general control structures in the interpreted language will be very difficult unless we divorce ourselves from the constrained structures at the outset.

It will be convenient to think of an implementation machine which has certain operations, which are "micro-coded" in LISP; these are used to operate on various "registers", which are represented as free LISP variables. These registers are:

EXP The expression currently being evaluated.

ENV A pointer to the environment in which to evaluate EXP.

CLINK A pointer to the frame for the computation of which the current one is a subcomputation.

PC The "program counter". As each "instruction" is executed, it updates PC to point to the next instruction to be executed.

VAL The returned value of a subcomputation. This register is not saved and restored in CLINK frames; in fact, its sole purpose is to pass values back safely across the restoration of a frame.

UNEVLIS, EVLIS These are utility registers which are part of the state of the interpreter (they are saved in CLINK frames). They are used primarily for evaluation of components of combinations, but may be used for other purposes also.

TEM A super-temporary register, used for random purposes, and not saved in CLINK frames or across interrupts. It therefore may not be used to pass information between "instructions" of the "machine", and so is best thought of as an internal hardware register.

QUEUE A list of all processes other than the one currently being interpreted.

TICK A magic register which a "hardware clock" sets to T every so often, used to drive the scheduler.

PROCESS This register always contains the name of the process currently swapped in and running.

The following declarations and macros are present only to make the compiler happy, and to make the version number of the SCHEME implementation available in the global variable VERSION.

(DECLARE (SPECIAL EXP UNEVLIS ENV EVLIS PC CLINK VAL TEM TOP QUEUE TICK PROCESS QUANTUM VERSION LISPVERSION))

(DEFUN VERSION MACRO (X) (COND (COMPILER-STATE (LIST 'QUOTE (STATUS UREAD))) (T (RPLACA X 'QUOTE) (RPLACD X (LIST VERSION)) (LIST 'QUOTE VERSION))))

(DECLARE (READ))

This section is missing here.

(setq Version ((lambd0 (compiler-state) (version)) T))

This section is missing here.

(setq **quantum** 1000000. Alarmclock 'settick)

This section is missing here.

(defun Primop (x) (getl X '(subr Expr Lsubr)))

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(deffprop Labels Labels Aint)

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Acknowledgments