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Definitional Interpreters for Higher-Order Programming Languages ^{*}

JOHN C. REYNOLDS **

Systems and Information Science, Syracuse University

Abstract. Higher-order programming languages (i.e., languages in which procedures or labels can occur as values) are usually defined by interpreters that are themselves written in a programming language based on the lambda calculus (i.e., an applicative language such as pure LISP). Examples include McCarthy’s definition of LISP, Landin’s SECD machine, the Vienna definition of PL/I, Reynolds’ definitions of GEDANKEN, and recent unpublished work by L. Morris and C. Wadsworth. Such definitions can be classified according to whether the interpreter contains higher-order functions, and whether the order of application (i.e., call by value versus call by name) in the defined language depends upon the order of application in the defining language. As an example, we consider the definition of a simple applicative programming language by means of an interpreter written in a similar language. Definitions in each of the above classifications are derived from one another by informal but constructive methods. The treatment of imperative features such as jumps and assignment is also discussed.

Keywords: programming language, language definition, interpreter, lambda calculus, applicative language, higher-order function, closure, order of application, continuation, LISP, GEDANKEN, PAL, SECD machine, J-operator, reference.

1. Introduction

An important and frequently used method of defining a programming language is to give an interpreter for the language that is written in a second, hopefully better understood language. (We will call these two languages the defined and defining languages, respectively.) In this paper, we will describe and classify several varieties of such interpreters, and show how they may be derived from one another by informal but constructive methods. Although our approach to “constructive classification” is original, the paper is basically an attempt to review and systematize previous work in the field, and we have tried to make the presentation accessible to readers who are unfamiliar with this previous work.

(Of course, interpretation can provide an implementation as well as a definition, but there are large practical differences between these usages. Definitional interpreters often achieve clarity by sacrificing all semblance of efficiency.)

We begin by noting some salient characteristics of programming languages themselves. The features of these languages can be divided usefully into two categories: applicative features, such as expression evaluation and the definition and application

* Work supported by Rome Air Force Development Center Contract No. 30602-72-C-0281 and ARPA Contract No. DAHC04-72-C-0003. This paper originally appeared in the Proceedings of the ACM National Conference, volume 2, August, 1972, ACM, New York, pages 717–740. ** Current address: Department of Computer Science, Carnegie Mellon University, Pittsburgh, PA 15213, USA. e-mail: John.Reynolds@cs.cmu.edu

of functions, and imperative features, such as statement sequencing, labels, jumps, assignment, and procedural side-effects. Most user-oriented languages provide features in both categories. Although machine languages are usually purely imperative, there are few “higher-level” languages that are purely imperative. (IPL/V might be an example.) On the other hand, there is at least one well-known example of a purely applicative language: LISP (i.e., the language defined in McCarthy’s original paper [1]; most LISP implementations provide an extended language including imperative features). There are also several more recent, rather theoretical languages (ISWIM [2], PAL [3], and GEDANKEN [4]) that have been designed by starting with an applicative language and adding imperative extensions.

Purely applicative languages are often said to be based on a logical system called the lambda calculus [5, 6], or even to be “syntactically sugared” versions of the lambda calculus. In particular, Landin [7] has shown that such languages can be reduced to the lambda calculus by treating each type of expression as an abbreviation for some expression of the lambda calculus. Indeed, this kind of reducibility could be taken as a precise definition of the notion of “purely applicative.” However, as we will see, although an unsugared applicative language is syntactically equivalent to the lambda calculus, there is a subtle semantic difference. Essentially, the semantics of the “real” lambda calculus implies a different “order of application” (i.e., normal-order evaluation) than most applicative programming languages.

A second useful characterization is the notion of a higher-order programming language. In analogy with mathematical logic, we will say that a programming language is higher-order if procedures or labels can occur as data, i.e., if these entities can be used as arguments to procedures, as results of functions, or as values of assignable variables. A language that is not higher-order will be called first-order.

In ALGOL and its various descendents, procedures and labels can be used as procedure arguments, and in more recent languages such as PL/I and ALGOL 68, they may also be used as function results and assignable values, subject to certain “scope” restrictions (which are imposed to preserve a stack discipline for the storage allocation of the representations of functions and labels). However, the unrestricted use of procedures and labels as data is permitted in only a handful of languages which sacrifice efficiency for generality: LISP (in most of its interpretive implementations), ISWIM, PAL, GEDANKEN, and (roughly) POP-2.

With regard to current techniques of language definition, there is a substantial disparity between first-order and higher-order languages. As a result of work by Floyd [8], Manna [9], Hoare [10], and others, most aspects of first-order languages can be defined logically, i.e., one can give an effective method for transforming a program in the defined language into a logical statement of the relation between its inputs and outputs. However, it has not yet been possible to apply this approach to higher-order languages. (Although recent work by Scott [12, 13, 14, 15] and Milner [16] represents a major step in this direction.)

Almost invariably, higher-order languages have been defined by the approach discussed in this paper, i.e., by giving interpreters that are themselves written in a programming language (An apparent exception is definition of ALGOL given by Burstall [17], but this can be characterized as a logical definition of a firstorder interpreter for a higher-order language.) Moreover, even when the defined language contains imperative features, the defining language is usually purely applicative (probably because applicative languages are well suited for computations with symbolic expressions). Examples include McCarthy’s definition of LISP [1], Landin’s SECD machine [7], the Vienna definition of PL/I [18], Reynolds’ definitions of GEDANKEN [19], and recent unpublished work by L. Morris [20] and C. Wadsworth.

(There are a few instances of definitional interpreters that fall outside the conceptual framework developed in this paper. A broader review of the field is given by deBakker [21].)

These examples exhibit considerable variety, ranging from very concise and abstract interpreters to much more elaborate and machine-like ones. To achieve a more precise classification, we will introduce two criteria. First, we ask whether the defining language is higher-order, or more precisely, whether any of the functions that comprise the interpreter either accept or produce values that are themselves functions.

The second criterion involves the notion of order of application. In designing any language that allows the use of procedures or functions, one must choose between two orders of application which are called (following ALGOL terminology) call by value and call by name. Even when the language is purely applicative, this choice will affect the meaning of some, but not all, programs that can be written in the language. Remembering that an interpreter is a specific program, we obtain our second criterion: Does the meaning of the interpreter depend upon the order of application chosen for the defining language?

These two criteria establish four possible classes of interpreters, each of which contains one or more of the examples cited earlier:

Order-of-  Use of higher-order functions:

application

dependence:  yes  no
yes  direct interpreter  McCarthy’s
for GEDANKEN  definition of LISP
no  Morris-Wadsworth SECD machine,
method  Vienna definition

The main goal of this paper is to illustrate and relate these classes of definitional interpreters. In the next section we will introduce a simple applicative language, which we will use as the defining language and also, with several restrictions, as the defined language. Then we will present a simple interpreter that uses higherorder functions and is order-of-application dependent, and we will transform this interpreter into examples of the three remaining classes. Finally, we will consider the problem of adding imperative features to the defined language (while keeping the defining language purely applicative).

2. A Simple Applicative Language

In an applicative language, the meaningful phrases of a program are called expressions, the process of executing or interpreting these expressions is called evaluation, and the result of evaluating an expression is called a value. However, as is evident from a simple arithmetic expression such as x + y, different evaluations of the same expression can produce different values, so that the process of evaluation must depend upon something more than just the expression being evaluated. It is evident that this “something more” must specify a value for every variable that might occur in the expression (more precisely, occur free). We will call such a specification an environment, and say that it binds variables to values.

It is also evident that the evaluation process may involve the creation of new environments from old ones. Suppose x_{1}, …, x_{n} are variables, v_{1}, …, v_{n} are values, and e and e′ are environments. If e′ specifies the value v_{i} for each x_{i}, and behaves the same way as e for all other variables, then we will say that e′ is the extension of e that binds the x_{i}’s to the v_{i}’s.

The simplest expressions in our applicative language are constants and variables. The evaluation of a constant always gives the same value, regardless of the environment. We will not specify the set of constants precisely, but will assume that it contains the integers and the Boolean constants true and false. The evaluation of a variable simply produces the value that is bound to that variable by the environment. In the programs in this paper we will denote variables by alphanumeric strings, with occasional superscripts and subscripts.

If our language is going to involve functions, then we must have a form of expression whose evaluation will cause the application of function to its arguments. If r_{0}, r_{1}, …, r_{n} are expressions, then r_{0}(r_{1}, …, r_{n}) is an application expression, whose operator is r_{0} and whose operands are r_{1}, …, r_{n}. The evaluation of an application expression in an environment proceeds as follows:

1. The subexpressions r_{0}, r_{1}, …, r_{n} are evaluated in the same environment to obtain values f, a_{1}, …, a_{n}.

2. If f is not a function of n arguments, then an error stop occurs.

3. Otherwise, the function f is applied to the arguments a_{1}, …, a_{n}, and if this application produces a result, then the result is the value of the application expression.

There are several assumptions hiding behind this description that need to be made explicit:

1. A “function of n arguments” is a kind of value that can be subjected to the process of being “applied” to a sequence of n values called “arguments”.

2. For some functions and arguments, the process of application may never produce a result, either because the process does not terminate (i.e., it runs on forever), or because it causes an error stop. Similarly, for some expressions and environments, the process of evaluation may never produce a value.

3. In a purely applicative language, the application of the same function to the same sequence of arguments will always have the same effect, i.e., both the result that is produced, and the prior question of whether any result is produced, depend only upon the function and its arguments. Similarly, the evaluation of the same expression in the same environment will always have the same effect.

4. During the evaluation of an application expression, the application process does not begin until after the operator and all of its operands have been evaluated. This is the call-by-value order of application mentioned in the introduction. In the alternative order of application, known as call by name, the application process would begin as soon as the operator had been evaluated, and each operand would only be evaluated when (and if) the function being applied actually depended upon its value. This distinction will be clarified below.

5. Although we have specified that all of the subexpressions r_{0}, …, r_{n} are to be evaluated before the application process begins we have not specified the relative order in which these subexpressions are to be evaluated. In a purely applicative language, this choice has no effect. (A slight exception occurs if the evaluation of one subexpression never terminates while the evaluation of another gives an error stop.) However, the choice will become significant when we start adding imperative features to the defined language. In anticipation of this extension, we will assume that the subexpressions are evaluated successively from left to right.

Next, we must have a form of expression whose evaluation will produce a function. If x_{1}, …, x_{n} are variables and r is an expression, then λ(x_{1}, …, x_{n}). r is a lambda expression, whose formal parameters are x_{1}, …, x_{n} and whose body is r. (The parentheses may be omitted if there is only one formal parameter.) The evaluation of a lambda expression with n formal parameters always terminates and always produces a function of n arguments. To describe this function, we must specify what will happen when it is applied to its arguments.

Suppose that f is the function obtained by evaluating λ(x_{1}, …, x_{n}). r in an environment e. Then the application of f to the arguments a_{1}, …, a_{n} will cause the evaluation of the body r in the environment that is the extension of e that binds each x_{i} to the corresponding a_{i}. If this evaluation produces a value, then the value becomes the result of the application of f.

The key point is that the environment in which the body is evaluated during application is an extension of the earlier environment in which the lambda expression was evaluated (rather than the more recent environment in which the application takes place). As a consequence, if a lambda expression contains global variables (i.e., variables that are not formal parameters), its evaluation in different environments can produce different functions. For example, the lambda expression λx. x + y can produce an incrementing function, an identity function (for the integers), or a decrementing function, when evaluated in environments that bind y to the values 1, 0, or −1 respectively.

Nowadays, it is generally accepted that this behavior of lambda expressions and environments is a basic characteristic of a well-designed higher-order language. Its importance is that it permits functional data to depend upon the partial results of a program.

Having introduced application and lambda expressions, we may now clarify the distinction between call by value and call by name. Consider the evaluation of an application expression r_{0}(r_{1}, …, r_{n}) in an environment e_{a}, and suppose that the value of the operator r_{0} is a function f that was originally created by evaluating the lambda expression λ(x_{1}, …, x_{n}). r_{λ} in an environment e_{λ}. (Possibly this lambda expression is r_{0} itself, but more generally r_{0} may be a non-lambda expression whose functional value was created earlier in the computation.) When call by value is used, the following steps will occur during the evaluation of the application expression:

1. r_{0} is evaluated in the environment e_{a} to obtain the function value f.

2. r_{1}, …, r_{n} are evaluated in the environment e_{a} to obtain arguments a_{1}, …, a_{n}.

3. r_{λ} is evaluated in the extension of e_{λ} that binds each x_{i} to the corresponding a_{i}, to obtain the value of the application expression.

When call by name is used, the same expressions are evaluated in the same environments. But the evaluations of the operands r_{1}, …, r_{n} will occur at a later time and may occur a different number of times. Specifically, instead of being evaluated before step (3), each operand r_{i} is repeatedly evaluated during step (3), each time that its value a_{i} is actually used (as a function to be applied, a Boolean value determining a branch, or an argument of a primitive operation).

At first sight, since the evaluation of the same expression in the same environment always produces the same effect, it would appear that the result of a program in a purely applicative language should be unaffected by changing the order of application (although it is evident that the repeated evaluation of operands occurring with call by name can be grossly inefficient). But this overlooks the possibility that “repeatedly” may mean “never”. During step (3) of the evaluation of r_{0}(r_{1}, …, r_{n}), it may happen that certain arguments a_{i} are never used, so that the corresponding operands r_{i} will never be evaluated under call by name. Now suppose that the evaluation of one of these r_{i} never terminates (or gives an error stop). Then the evaluation of the original application expression will terminate under call by name but not call by value. In brief, changing the order of application can affect the value of an application expression when the function being applied is independent of some of its arguments and the corresponding operands are nonterminating.

(In ALGOL the distinction between call by value and call by name also involves a change in “coercion conventions”. However, this change is irrelevant in the absence of assignment.)

In the defined language, we will consider only the use of call by value, but in the defining language we will consider both orders of application. In particular, we will inquire whether the above-described situation occurs in our interpreters, so that changing the order of application in the defining language can affect the meaning of the defined language.

We now introduce some additional kinds of expressions. If r_{p}, r_{c} and r_{a} are expressions, then if r_{p} then r_{c} else r_{a} is a simple conditional expression, whose premiss is r_{p}, whose conclusion is r_{c}, and whose alternative is r_{a}. The evaluation of a conditional expression in an environment e begins with the evaluation of its premiss r_{p} in the same environment. Then, depending upon whether the value of the premiss is true or false, the value of the conditional expression is obtained by evaluating either the conclusion r_{c} or the alternative r_{a} in e. Any other value of the premiss causes an error stop.

It is also convenient to use a LISP-like notation for “multiple” conditional expressions. If r_{p1}, …, r_{pn} and r_{c1}, …, r_{cn} are expressions, then

(r_{p1} → r_{c1}, r_{p2}, → r_{c2}, …, r_{pn} → r_{cn})

is a multiple conditional expression, with the same meaning as the following sequence of simple conditional expressions:

if r_{p1} then r_{c1} else if r_{p2} then r_{c2} else ⋯ if r_{pn} then r_{cn} else error.

Next, we introduce a form of expression (due to Landin [7]) that is analogous to the block in ALGOL. If x_{1}, …, x_{n} are variables, and r_{1}, …, r_{n} and r_{b} are expressions, then

let x_{1} = r_{1} and ⋯ and x_{n} = r_{n} in r_{b}

is a let expression, whose declared variables are x_{1}, …, x_{n}, whose declaring expressions are r_{1}, …, r_{n}, and whose body is r_{b}. (We will call each pair x_{i} = r_{i} a declaration.) The evaluation of a let expression in an environment e begins with the evaluation of its declaring expressions r_{i} in the same environment. Then the value of the let expression is obtained by evaluating its body r_{b} in the environment that is the extension of e that binds each declared variable x_{i} to the value of the corresponding declaring expression r_{i}.

It should be noted that the extended environment only affects the evaluation of the body, not the declaring expressions. For example, in an environment that binds x to 4, the value of let x = x + 1 and y = x − 1 in x × y is 15. As a consequence, let expressions cannot be used (at least directly) to define recursive functions. One might expect, for instance, that

let f = λx. if x = 0 then 1 else x × f(x − 1) in ⋯

would create an extended environment in which f was bound to a recursive function (for computing the factorial). But in fact, the occurrence of f inside the declaring expression will not “feel” the binding of f to the value of the declaring expression, so that the resulting function will not call itself recursively.

To overcome this problem, we introduce a second kind of block-like expression. If x_{1}, …, x_{n} are variables, ℓ_{1}, …, ℓ_{n} are lambda expressions, and r_{b} is an expression, then

letrec x_{1} = ℓ_{1} and ⋯ and x_{n} = ℓ_{n} in r_{b}

is a recursive let expression, whose declared variables are x_{1}, …, x_{n}, whose declaring expressions are ℓ_{1}, …, ℓ_{n}, and whose body is r_{b}. The value of a recursive let expression in an environment e is obtained by evaluating its body in an environment e′ which satisfies the following property: e′ is the extension of e that binds each declared variable x_{i} to the function obtained by evaluating the corresponding declaring lambda expression ℓ_{i} in the environment e′.

There is a circularity in the property “e′ is the … in the environment e′” that is characteristic of recursion, and that prevents this property from being an explicit definition of e′. To be rigorous, we would have to show that there actually exists an environment that satisfies this property, and also deal with the possibility that this environment might not be unique. The mathematical techniques needed to achieve this rigor are beyond the scope of this paper [22, 12, 13, 14, 15]. However, we will eventually derive an interpreter that defines recursive let expressions more explicitly.

(It is possible to generalize recursive let expressions by allowing arbitrary declaring expressions. We have chosen not to do so, since the generalization would considerably complicate some of the definitional interpreters, and is not unique.)

To maintain generality, we have avoided specifying the set of data that can occur as the result of expression evaluation (beyond asserting that this set should contain functions and the Boolean values true and false). However, it is evident that our language must contain basic (i.e., built-in) operations and tests for manipulating this data. For example, if integers are to occur as data, we will need at least an incrementing operation and a test for integer equality. More likely, we will want all of the usual arithmetic operations and tests. If some form of structured data is to be used, we will need operations for constructing and analyzing the structures, and tests for classifying them.

Regardless of the specific nature of the data, there are three ways to introduce basic operations and tests into our applicative language:

1. We may introduce constants denoting the basic functions (whose application will perform the basic operations and tests).

2. We may introduce predefined variables denoting the basic functions. These variables differ from constants in that the programmer can redefine them with his own declarations. They are specified by introducing an initial environment, to be used for the evaluation of the entire program, that binds the predefined variables to their functional values.

3. We may introduce special expressions whose evaluation will perform the basic operations and tests. Since this approach is used in most programming languages (and in mathematical notation), we will frequently use the common forms of arithmetic and Boolean expressions without explanation.

3. The Defined Language

Although our defining language will use all of the features described in the previous section, along with appropriate basic operations and tests, the defined language will be considerably more limited, in order to avoid complications that would be out of place in an introductory paper. Specifically:

1. Functions will be limited to a single argument. Thus all applicative expressions will have a single operand, and all lambda expressions will have a single formal parameter.

2. Only call by value will be used.

3. Only simple conditional expressions will be used.

4. Nonrecursive let expressions will be excluded.

5. All recursive let expressions will contain a single declaration.

6. Values will be integers, booleans, and functions. The only basic operations and tests will be functions for incrementing integers and for testing integer equality, denoted by the predefined variables succ and equal, respectively.

The reader may accept an assurance that these limitations will eliminate a variety of tedious complications without evading any intellectually significant problems. Indeed, with slight exceptions, the eliminated features can be regarded as syntactic sugar, i.e., they can be defined as abbreviations for expressions in the restricted language [7, 4].

4. Abstract Syntax

We now turn our attention to the defining language. To permit the writing of interpreters, the values used in the defining language must include expressions of the defined language. At first sight, this suggests that we should use character strings as values denoting expressions, but this approach would enmesh us in questions of grammar and parsing that are beyond the scope of this paper. (An excellent review of these matters is contained in Reference [23].)

Instead, we use the approach of abstract syntax, originally suggested by McCarthy [24]. In this approach, it is assumed that programs are “really” abstract, hierarchically structured data objects, and that the character strings that one actually reads into the computer are simply representations of these abstract objects (in the same sense that digit strings are representations of integers). Thus the problems of grammar and parsing can be set aside as “input editing”. (Of course, this does not eliminate these problems, but it separates them clearly from semantic considerations. See, for example, Wozencraft and Evans [25].)

We are left with two closely related problems: how to define sets of abstract expressions (and other structured data to be used by the interpreters), and how to define the basic functions for constructing, analyzing, and classifying these objects. Both problems are solved by introducing three forms of abstract-syntax equations. (A more elaborate defined language would require a more complex treatment of abstract syntax, as given in Reference [18], for example.) Within these equations, upper-case letter strings denote sets, and lower-case letter strings denote basic functions.

Let S_{0}, S_{1}, …, S_{n} be upper-case letter strings and a_{1}, …, a_{n} be lowercase letter strings. Then a record equation of the form

S_{0} = [a_{1} : S_{1}, …, a_{n} : S_{n}]

implies that:

1. S_{0} is a set, disjoint from any other set defined by a record equation, whose members are records with n fields in which the value of the ith field belongs to the set S_{i}. (Mathematically, S_{0} is a disjoint set in one-to-one correspondence with the Cartesian product S_{1} × ⋯ × S_{n}.)

2. Each a_{i} (is a predefined variable which) denotes the selector function that accepts a member of S_{0} and produces its ith field value.

3. Let s_{0} be the string obtained from S_{0} by lowering the case of each character. Then s_{0}? denotes the classifier function that tests whether its argument belong to S_{0}, and mk-s_{0} denotes the constructor function of n arguments (belonging to the sets S_{1}, …, S_{n}) that creates a record in S_{0} from its field values.

For example, the record equation

APPL = [opr: EXP, opnd: EXP]

implies that an application expression (i.e., a member of APPL) is a two-field record whose field values are both expressions (i.e., members of EXP). It also implies that opr and opnd are selector functions that produce the first and second field values of an application expression, that appl? is a classifier function that tests whether a value is an application expression, and that mk-appl is a two-argument constructor function that constructs an application expression from its field values. It is evident that if r_{1} and r_{2} are expressions,

opr(mk-appl(r_{1}, r_{2})) = r_{1}
opnd(mk-appl(r_{1}, r_{2})) = r_{2},

and if appl?(r) is true,

mk-appl(opr(r), opnd(r)) = r.

The remaining forms of abstract syntax equations are the union equation:

S_{0} = S_{1} ∪ ⋯ ∪ S_{n},

which implies that S_{0} is the union of sets S_{1}, …, S_{n}, and the function equation:

S_{0} = S_{1}, …, S_{n} → S_{r},

which implies that S_{0} is the set of n-argument functions that accept arguments in S_{1}, …, S_{n} and produce results in S_{r}. (More precisely, S_{0} is the set of n-argument functions f with the property that if f is applied to arguments in the sets S_{1}, …, S_{n}, and if f terminates without an error stop, then the result of f belongs to S_{r}.)

We may now use these forms of abstract syntax equations to define the principal set of data used by our interpreters, i.e., the set EXP of expressions of the defined language:

EXP = CONST ∪ VAR ∪ APPL ∪ LAMBDA ∪ COND ∪ LETREC APPL = [opr: EXP, opnd: EXP] LAMBDA = [fp: VAR, body: EXP] COND = [prem: EXP, conc: EXP, altr: EXP] LETREC = [dvar: VAR, dexp: LAMBDA, body: EXP].

A cumbersome but fairly accurate translation into English is that an expression (i.e., a member of EXP) is one of the following:

1. A constant (a member of CONST),

2. A variable (a member of VAR),

3. An application expression (a member of APPL), which consists of an expression called its operator (selected by the basic function opr) and an expression called its operand (selected by opnd),

4. A lambda expression (a member of LAMBDA), which consists of a variable called its formal parameter (selected by fp) and an expression called its body (selected by body),

5. A conditional expression (a member of COND), which consists of an expression called its premiss (selected by prem) and an expression called its conclusion (selected by conc) and an expression called its alternative (selected by altr),

6. A recursive let expression (a member of LETREC), which consists of a variable called its declared variable (selected by dvar), a lambda expression called its declaring expression (selected by dexp), and an expression called its body (selected by body).

We have purposely left the sets CONST and VAR unspecified. For CONST, we will assume only that there is a basic function const? which tests whether its argument is a constant, and a basic function evcon which maps each constant into the value that it denotes. For VAR, we will assume that there is a basic function var? which tests whether its argument is a variable, that variables can be tested for equality (of the variables themselves, not their values), and that two particular variables are denoted by the quoted strings “succ” and “equal”.

We must also define the abstract syntax of two other data sets that will be used by our interpreter. The first is the set VAL of values of the defined language:

VAL = INTEGER ∪ BOOLEAN ∪ FUNVAL

FUNVAL = VAL → VAL.

One must be careful not to confuse values in the defined and defining languages. Strictly speaking, VAL is a subset of the values of the defining language whose members represent the values of the defined language. However, since the variety of values provided in the defining language is richer than in the defined language, we have been able to represent each defined-language value by the same defininglanguage value. In our later interpreters this situation will change, and it will become more evident that VAL is a set of value representations.

Finally, we must define the set ENV of environments. Since the purpose of an environment is to specify the value that is bound to each variable, the simplest approach is to assume that an environment is a function from variables to values, i.e.,

ENV = VAR → VAL.

Within the various interpreters that we will present, each variable will range over some set defined by abstract syntax equations. For clarity, we will use different variables for different sets, as summarized in the following table:

Variable Range  Variable Range
r  EXP  e e′  ENV
x z  VAR  c c′  CONT
ℓ  LAMBDA  m m′ m′′ MEM
a b  VAL  rf  REF
f  FUNVAL  n  INTEGER

(The sets CONT, MEM, and REF will be defined later.)

5. A Meta-Circular Interpreter

Our first interpreter is a straightforward transcription of the informal language definition we have already given. Its central component is a function eval that produces the value of an expression r in a environment e:

eval = λ(r, e). I.1 (const?(r) → evcon(r), I.2 var?(r) → e(r), I.3 appl?(r) → (eval(opr(r), e))(eval(opnd(r), e)), I.4 lambda?(r) → evlambda(r, e), I.5 cond?(r) → if eval(prem(r), e) I.6 then eval(conc(r), e) else eval(altr(r), e), I.7 letrec?(r) → letrec e′ = I.8

λx. if x = dvar(r) then evlambda(dexp(r), e′) else e(x) I.9

in eval(body(r), e′)) I.10

evlambda = λ(ℓ, e). \underline{λa. eval(body(ℓ), ext(fp(ℓ), a, e))} I.11

ext = λ(z, a, e). λx. if x = z then a else e(x). I.12

The subsidiary function evlambda produces the value of a lambda expression ℓ in an environment e. (We have extracted it as a separate function since it is called from two places, in lines I.5 and I.9.) The subsidiary function ext produces the extension of an environment e that binds the variable z to the value a. It should be noted that, in the evaluation of a recursive let expression (lines I.8 to I.10), the circularity in the definition of the extended environment e′ is handled by making e′ a recursive function. (However, it is a rather unusual recursive function which, instead of calling itself, calls another function evlambda, to which it provides itself as an argument.)

The function eval does not define the meaning of the predefined variables. For this purpose, we introduce the “main” function interpret, which causes a complete program r to be evaluated in an initial environment initenv that maps each predefined variable into the corresponding basic function:

interpret = λr. eval(r, initenv) I.13 initenv = λx. (x = “succ” → \underline{λa. succ(a)}, … I.14

x = “equal” → \underline{λa. \underline{λb. equal(a, b)}}) . . .. I.15

In the last line we have used a trick called Currying (after the logician H. Curry) to solve the problem of introducing a binary operation into a language where all functions must accept a single argument. (The referee comments that although “Currying” is tastier, “Schönfinkeling” might be more accurate.) In the defined language, equal is a function which accepts a single argument a and returns another function, which in turn accepts a single argument b and returns true or false depending upon whether a = b. Thus in the defined language, one would write (equal(a))(b) instead of equal(a, b).

(Each of our interpreters will consist of a sequence of function declarations. We will assume that these are implicitly embedded in a recursive let expression whose body is interpret(R), where R is the program to be interpreted.)

We have coined the word “meta-circular” to indicate the basic character of this interpreter: It defines each feature of the defined language by using the corresponding feature of the defining language. For example, when eval is applied to an application expression (lambda expression, conditional expression, recursive let expression) of the defined language, it evaluates an application expression (lambda expression, conditional expression, recursive let expression) in the defining language. Similarly, the initial environment defines the basic functions of the defined language in terms of the same functions in the defining language.

In one sense, this situation is not undesirable. For the reader who already has a thorough and correct understanding of the defining language, a meta-circular definition will provide a concise and complete description of the defined language. (Of course this is a rather vacuous accomplishment when the defined language is a subset of the defining language.) The problem is that any misunderstandings about the defining language are likely to be carried over to the defined language intact. For example, if we were to assume that in the defining language, the function succ decreases an integer by one, or that a conditional expression gives the same result when the value of its premiss is non-Boolean as when it is false, the above interpreter would lead us to the same assumptions about the defined language.

These particular difficulties are easily overcome; we could define functions such as succ in terms of elementary mathematics, and we could insert explicit tests for erroneous values. But there are three objections to meta-circularity that are much more serious:

1. The meta-circular interpreter does not shed much light on the nature of higherorder functions. For this purpose, we would prefer an interpreter of a higherorder defined language that was written in a first-order defining language.

2. Changing the order of application used in the defining language induces a similar change in the defined language. To see this, suppose that eval is applied to an application expression r_{0}(r_{1}) of the defined language. Then the result of eval will be obtained by evaluating the application expression (line I.4)

(eval(r_{0}, e))(eval(r_{1}, e))

in the defining language. If call by value is used in the defining language, then eval(r_{1}, e) will be evaluated before the functional value of eval(r_{0}, e) is applied. But evaluating eval(r_{1}, e) interprets the evaluation of r_{1}, and applying the value of eval(r_{0}, e) interprets the application of the value of r_{0}. Thus in terms of the defined language, r_{1} will be evaluated before the value of r_{0} is applied, i.e., call by value will be used in the defined language.

On the other hand, if call by name is used in the defining language, then the application of the functional value of eval(r_{0}, e) will begin as soon as eval(r_{0}, e) has been evaluated, and the operand eval(r_{1}, e) will only be evaluated when and if the function being applied depends upon its value. In terms of the defined language, the application of the value of r_{0} will begin as soon as r_{0} has been evaluated, and the operand r_{1} will only be evaluated when and if the function being applied depends upon its value, i.e., call by name will be used in the defined language.

3. Suppose we wish to extend the defined language by introducing the imperative features of labels and jumps (including jumps out of blocks). As far as is known, it is impossible to extend the meta-circular definition straightforwardly to accommodate these features (without introducing similar features into the defining language).

In the following sections we will develop transformations of the meta-circular interpreter that will meet the first two of these objections. Then we will find that the transformation designed to meet the second objection also meets the third.

It should be emphasized that, although these transformations are motivated by their application to interpreters, they are actually applicable to any program written in the defining language, and their validity depends entirely upon the properties of the defining language.

6. Elimination of Higher-Order Functions

Our first task is to modify the meta-circular interpreter so that none of the functions that comprise this interpreter accept arguments or produce results that are functions. An examination of the abstract syntax shows that this goal will be met if we can replace the two sets FUNVAL and ENV by sets of values that are not functions. Specifically, the new members of these sets will be records that represent functions.

We first consider the set FUNVAL. Since the new members of this set are to be records rather than functions, we can no longer apply these members directly to arguments. Instead we will introduce a new function apply that will “interpret” the new members of FUNVAL. Specifically, if f_{new} is a record in FUNVAL that represents a function f_{old} and if a is any member of VAL, then apply(f_{new}, a) will produce the same result as f_{old}(a). Assuming for the moment that we will be able to define apply, we must replace each application of a member of FUNVAL (to an argument a) by an application of apply (to the member of FUNVAL and the argument a). In fact, the only such application occurs in line I.4, which must become

appl?(r) → apply(eval(opr(r), e), eval(opnd(r), e)). I.4′

To decide upon the form of the new members of FUNVAL, we recall that whenever a function is obtained by evaluating a lambda expression, the function will be determined by two items of information: (1) the lambda expression itself, and (2) the values that were bound to the global variables of the lambda expression at the time of its evaluation. It is evident that these items of information will be sufficient to represent the function. This suggests that the new set FUNVAL should be a union of disjoint sets of records, one set for each lambda expression whose value belonged to the old FUNVAL, and that the fields of each record should contain values of the global variables of the corresponding lambda expression.

In fact, the meta-circular interpreter contains four lambda expressions (indicated by solid underlining) that produce members of FUNVAL. The following table gives their locations and global variables, and the equations defining the new sets of records that will represent their values. (The connotations of the set and selector names we have chosen will become apparent when we discuss the role of these entities in the interpretation of the defined language.)

Location  Global Variables New Record Equation
I.11  ℓ e  CLOSR = [lam: LAMBDA, en: ENV]
I.14  none  SC = [ ]
I.15 (outer) none  EQ1 = [ ]
I.15 (inner) a  EQ2 = [arg1 : VAL]

Thus the new set FUNVAL will be

FUNVAL = CLOSR ∪ SC ∪ EQ1 ∪ EQ2, and the overall structure of apply will be:

apply = λ(f, a).

(closr?(f) → ⋯,
sc?(f) → ⋯,
eq1?(f) → ⋯,
eq2?(f) → ⋯).

Our remaining task is to replace each of the four solidly underlined lambda expressions by appropriate record-creation operations, and to insert expressions in the branches of apply that will interpret the corresponding records. The lambda expression in line I.11 must be replaced by an expression that creates a CLOSR-record containing the value of the global variables ℓ and e:

evlambda = λ(ℓ, e). mk-closr(ℓ, e). I.11′

Now apply(f, a) must produce the result of applying the function represented by f to the argument a. When f is a CLOSR-record, this result may be obtained by evaluating the body

eval(body(ℓ), ext(fp(ℓ), a, e))

of the replaced lambda expression in an appropriate environment. This environment must bind the formal parameter a of the replaced lambda expression to the value of a and must bind the global variables ℓ and e of the lambda expression to the same value as the environment in which the CLOSR-record f was created. Since the latter values are stored in the fields of f, we have:

apply = λ(f, a).

(closr?(f) → let a = a and ℓ = lam(f) and e = en(f)

in eval(body(ℓ), ext(fp(ℓ), a, e)), …).

(In this particular case, but not in general, the declaration a = a is unnecessary, since the formal parameter of the replaced lambda expression and the second formal parameter of apply are the same variable. From now on, we will omit such vacuous declarations.)

A similar treatment (somewhat simplified since there are no global variables) of the lambda expression in I.14 and the outer lambda expression in I.15 gives:

initenv = λx. (x = “succ” → mk-sc(), … I.14′

x = “equal” → mk-eq1 ()) … I.15′

and

apply = λ(f, a).

(closr?(f) → let ℓ = lam(f) and e = en(f) in eval(body(ℓ), ext(fp(ℓ), a, e)),

sc?(f) → succ(a),

eq1?(f) → \underline{λb. equal(a, b)},

eq2?(f) → ⋯).

Finally, we must replace the lambda expression that originally occurred as the inner expression in I.15. Although we have already moved this expression into the body of apply (since it was the body of a previously replaced lambda expression), the same basic treatment can be applied to the new occurrence, giving:

apply = λ(f, a).

(closr?(f) → let ℓ = lam(f) and e = en(f) in eval(body(ℓ), ext(fp(ℓ), a, e)),

sc?(f) → succ(a), eq1?(f) → mk-eq2 (a), eq2?(f) → let b = a and a = arg1 (f) in equal(a, b)).

(Note that the declaration relating formal parameters is not vacuous in this case.)

The entire transformation that converts FUNVAL from a set of functions to a set of records has been informally justified by appealing to an understanding of the defining language, without regard to the meaning or use of the particular program being transformed. But now it is illuminating to examine the different kinds of records in FUNVAL in terms of their role in the interpretation of the defined language. The records in the set CLOSR represent functional values that are produced by evaluating the lambda expressions occurring in the defined language programs. They are equivalent to the objects called FUNARG triplets in LISP and closures in the work of Landin [7]. The unique records in the one-element sets SC and EQ1 represent the basic functions succ and equal. Finally, the records in EQ2 represent the functions that are created by applying equal to one argument.

A similar transformation can be used to “defunctionalize” the set ENV of environments. To interpret the new members of ENV, we will introduce a function get, with the property that if e_{new} represents an environment e_{old} and x is a member of VAR, then get(e_{new}, x) = e_{old}(x). Applications of get must be inserted at the three points (in lines I.3, I.9, and I.12) in the interpreter where environments are applied to variables:

var?(r) → get(e, r), I.3′

⋮

λx. if x = dvar(r) then evlambda(dexp(r), e′) else get(e, x) I.9′

⋮

ext = λ(z, a, e). λx. if x = z then a else get(e, x). I.12′

Next, there are three lambda expressions that produce environments; they are indicated by broken underlining which we have carefully preserved during the previous transformations. The following table gives their locations and global variables, and the equations defining the new sets of records that will represent their values:

Location Global Variables New Record Equation
I.14′-15′ none  INIT = [ ]
I.12′  z a e  SIMP = [bvar: VAR, bval: VAL, old: ENV]
I.9′  r e e′  REC = [letx: LETREC, old: ENV, new: ENV]

Thus the new set of environment representations is:

ENV = INIT ∪ SIMP ∪ REC.

Replacement of the three environment-producing lambda expressions gives:

letrec?(r) → letrec e′ = mk-rec(r, e, e′) ⋯ I.8-9′′

⋮

ext = λ(z, a, e). mk-simp(z, a, e) I.12′′

⋮

initenv = mk-init(), I.14′′-15′′

and the environment-interpreting function is:

get = λ(e, x).

(init?(e) → (x = “succ” → mk-sc(), x = “equal” → mk-eq1 ()), simp?(e) → let z = bvar(e) and a = bval(e) and e = old(e)

in if x = z then a else get(e, x), rec?(e) → let r = letx(e) and e = old(e) and e′ = new(e)

in if x = dvar(r) then evlambda(dexp(r), e′) else get(e, x)).

But now we are faced with a new problem. By eliminating the lambda expression in I.9′, we have created a recursive let expression

letrec e′ = mk-rec(r, e, e′) ⋯

that violates the structure of the defining language, since its declaring subexpression is no longer a lambda expression. However, there is still an obvious intuitive interpretation of this illicit construction: it binds e′ to a “cyclic” record, whose last field is (a pointer to) the record itself.

If we accept this interpretation, then whenever e is a member of REC, we will have new(e) = e. This allows us to replace the only occurrence of new(e) by e, so that the penultimate line of get becomes:

rec?(e) → let r = letx(e) and e = old(e) and e′ = e · · · .

But now our program no longer contains any references to the cyclic new fields, so that these fields can be deleted from the records in REC. Thus the record equation for REC is reduced to:

REC = [letx: LETREC, old: ENV],

and the offending recursive let expression becomes:

letrec?(r) → let e′ = mk-rec(r, e) · · · . I.8′-9′′′

At this point, once we have collected the bits and pieces produced by the various transformations, we will have obtained an interpreter that no longer contains any higher-order functions. However, it is convenient to make a few simplications:

1. let expressions can be eliminated by substituting the declaring expressions for each occurrence of the corresponding declared variables in the body.

2. Line I.11′ can be eliminated by replacing occurrences of evlambda by mk-closr.

3. Line I.12′′ can be eliminated by replacing occurrences of ext by mk-simp.

4. Lines I.14′′-15′′ can be eliminated by replacing occurrences of initenv by mkinit().

Thus we obtain our second interpreter:

FUNVAL = CLOSR ∪ SC ∪ EQ1 ∪ EQ2 CLOSR = [lam: LAMBDA, en: ENV] SC = [ ] EQ1 = [ ] EQ2 = [arg1 : VAL] ENV = INIT ∪ SIMP ∪ REC INIT = [ ] SIMP = [bvar: VAR, bval: VAL, old: ENV] REC = [letx: LETREC, old: ENV] interpret = λr. eval(r, mk-init()) II.1 eval = λ(r, e). II.2

(const?(r) → evcon(r), II.3 var?(r) → get(e, r), II.4 appl?(r) → apply(eval(opr(r), e), eval(opnd(r), e)), II.5 lambda?(r) → mk-closr(r, e), II.6 cond?(r) → if eval(prem(r), e) II.7 then eval(conc(r), e) else eval(altr(r), e), II.8 letrec?(r) → eval(body(r), mk-rec(r, e))) II.9 apply = λ(f, a). II.10

(closr?(f) → II.11

eval(body(lam(f)), mk-simp(fp(lam(f)), a, en(f))), II.12 sc?(f) → succ(a), II.13 eq1?(f) → mk-eq2 (a), II.14 eq2?(f) → equal(arg1 (f), a)) II.15 get = λ(e, x). II.16 (init?(e) → (x = “succ” → mk-sc(), x = “equal” → mk-eq1 ()), II.17 simp?(e) → if x = bvar(e) then bval(e) else get(old(e), x), II.18 rec?(e) → if x = dvar(letx(e)) II.19 then mk-closr(dexp(letx(e)), e) else get(old(e), x)). II.20

Just as with FUNVAL, we may examine the different kinds of records in ENV with regard to their role in the interpretation of the defined language. The unique record in INIT has no subfields, while the records in SIMP and REC each have one field (selected by old) that is another member of ENV. Thus environments in our second interpreter are linear lists (in which each element specifies the binding of a single variable), and the unique record in INIT serves as the empty list.

It is easily seen that get(e, x) searches such a list to find the binding of the variable x. When get encounters a record in SIMP, it compares x with the bvar field, and if a match occurs, it returns the value stored in the bval field. When get encounters a record in REC, it compares x with dvar(letx(e)) (the declared variable of the recursive let expression that created the binding), and if a match occurs, it returns the value obtained by evaluating dexp(letx(e)) (the declaring subexpression of the same recursive let expression) in the environment e. The fact that e includes the very binding that is being “looked up” reflects the essential recursive characteristic that the declaring subexpression should “feel” the effect of the declaration in which it is embedded. When get encounters the empty list, it compares x with each of the predefined variables, and if a match is found, it returns the appropriate value.

The definition of get reveals the consequences of our restricting recursive let expressions by requiring that their declaring subexpressions should be lambda expressions. Because of this restriction, the declaring subexpressions are always evaluated by the trivial operation of forming a closure. Therefore, the function get always terminates, since it never calls any other recursive function, and can never call itself more times than the length of the list that it is searching. (On the other hand, if we had permitted arbitrary declaring subexpressions, line II.20 would contain eval(dexp(letx(e)), e) instead of mk-closr(dexp(letx(e)), e). This seemingly slight modification would convert get into a function that might run on forever, as for example, when looking up the variable k in an environment created by the defined-language construction letrec k = k + 1 in · · ·.)

The second interpreter is similar in style, and in many details, to McCarthy’s definition of LISP [1]. The main differences arise from our insistence upon FUNARG binding, the use of recursive let expressions instead of label expressions, and the use of predefined variables instead of variables with flagged property lists.

7. Continuations

The transition from the meta-circular interpreter to our second interpreter has not eliminated order-of-application dependence. It can easily be seen that a change in the order of application used in the defining-language expression (in II.5)

apply(eval(opr(r), e), eval(opnd(r), e))

will cause a similar change for all application expressions of the defined language.

To eliminate this dependence, we must first identify the circumstances under which an arbitrary program in the defining language will be affected by the order of application. The essential effect of switching from call by value to call by name is to postpone the evaluation of the operands of application expressions (and declaring subexpressions of let expressions), and to alter the number of times these operands are evaluated. We have already seen that in a purely applicative language, the only way in which this change can affect the meaning of a program is to avoid the evaluation of a nonterminating operand. Now suppose we define an expression to be serious if there is any possibility that its evaluation might not terminate. Then a sufficient condition for order-of-application independence is that a program should contain no serious operands or declaring expressions.

Next, suppose that we can divide the functions that may be applied by our program into serious functions, whose application may sometimes run on forever, and trivial functions, whose application will always terminate. (Of course, it is wellknown that one cannot effectively decide whether an arbitrary function will always terminate, but one can still establish this classification in a “fail-safe” manner, i.e., classify a function as serious unless it can be shown to terminate for all arguments.) Then an expression will only be serious if its evaluation can cause the application of a serious function, and a program will be independent of order-of-application if no operand or declaring expression can cause such an application.

At first sight, this condition appears to be so restrictive that it could not be met in a nontrivial program. As can be seen with a little thought, the condition implies that whenever some function calls a serious function, the calling function must return the same result as the called function, without performing any further computation. But any function that calls a serious function must be serious itself. Thus by induction, as soon as any serious function returns a result, every function must immediately return the same result, which must therefore be the final result of the entire program.

Nevertheless, there is a method for transforming an arbitrary program into one that meets our apparently restrictive condition. The underlying idea has appeared in a variety of contexts [26, 27, 28], but its application to definitional interpreters is due to L. Morris [20] and Wadsworth. Basically, one replaces each serious function f_{old} (except the main program) by a new serious function f_{new} that accepts an additional argument c called a continuation. The continuation will be a function itself, and f_{new} is expected to compute the same result as f_{old}, apply the continuation to this result, and then return the result of the continuation, i.e.,

f_{new}(x_{1}, …, x_{n}, c) = c(f_{old}(x_{1}, …, x_{n})).

This introduction of continuations provides an additional “degree of freedom” that can be used to meet the condition of order-of-application independence. Essentially, instead of performing further actions after a serious function has returned, one embeds the further actions in the continuation that is passed to the serious function.

To transform our second interpreter, we must first classify its functions. Since the defined language contains expressions and functions whose evaluation and application may never terminate, the defining-language functions eval and apply are serious and must be altered to accept continuations. On the other hand, since we have seen that get always terminates, it is trivial and will not be altered. (Note that this situation would change if the defined language permitted recursive let expressions with arbitrary declaring subexpressions.)

Both eval and apply produce results in the set VAL, so that the arguments of continuations will belong to this set. The result of a continuation will always be the value of the entire program being interpreted, which will also belong to the set VAL. Thus the set of continuations is:

CONT = VAL → VAL.

(In a more complicated interpreter in which different serious functions produced different kinds of results, we would introduce different kinds of continuations.)

The overall form of our transformed interpreter will be:

interpret = λr. eval(r, mk-init(), λa. a) II.1′ eval = λ(r, e, c). · · · II.2′ apply = λ(f, a, c). · · · II.10′ get = same as in Interpreter II. II.16–20

Note that the “main level” call of eval by interpret provides an identity function as the initial continuation.

We must now alter each branch of eval and apply to apply the continuation c to the former results of these functions. In lines II.3, 4, 6, 13, 14, and 15, the branches evaluate expressions which are not serious, and which are therefore permissible operands. Thus in these cases, we may simply apply the continuation c to each expression:

eval = λ(r, e, c). II.2′ (const?(r) → c(evcon(r)), II.3′ var?(r) → c(get(e, r)), II.4′

⋮

lambda?(r) → c(mk-closr(r, e)), …) II.6′ apply = λ(f, a, c). ( …, II.10′ sc?(f) → c(succ(a)), II.13′ eq1?(f) → c(mk-eq2 (a)), II.14′ eq2?(f) → c(equal(arg1 (f), a))). II.15′

In lines II.9 and II.12, the branches evaluate expressions that are serious themselves but contain no serious operands. By themselves, these expressions are permissible, but they must not be used as operands in applications of the continuation. The solution is straightforward; instead of applying the continuation c to the result of eval, we pass c as an argument to eval, i.e., we “instruct” eval to apply c before returning its result:

letrec?(r) → eval(body(r), mk-rec(r, e), c)) II.9′

⋮ (closr?(f) → II.11′ eval(body(lam(f)), mk-simp(fp(lam(f)), a, en(f)), c). II.12′

The most complex part of our transformation occurs in the branch of eval that evaluates application expressions in line II.5. Here we must perform four serious operations:

1. Evaluate the operator.

2. Evaluate the operand.

3. Apply the value of the operator to the value of the operand.

4. Apply the continuation c to the result of (3).

Moreover, we must specify explicitly that these operations are to be done in the above order. This will insure that the defined language uses call by value, and also that the subexpressions of an application expression are evaluated from left to right (operator before operand).

The solution is to call eval to perform operation (1), to give this call of eval a continuation that will call eval to perform operation (2), to give the second call of eval a continuation that will call apply to perform (3), and to give apply a continuation (the original continuation c) that will perform (4). Thus we have:

appl?(r) → eval(opr(r), e, λf. eval(opnd(r), e, λa. apply(f, a, c))). II.5′

A similar approach handles the branch that evaluates conditional expressions in lines II.7 and 8. Here there are three serious operations to be performed successively:

1. Evaluate the premiss.

2. Evaluate the conclusion or the alternative, depending on the result of (1).

3. Apply the continuation c to the result of (2).

The transformed branch is:

cond?(r) → eval(prem(r), e, II.7′ λb. if b then eval(conc(r), e, c) else eval(altr(r), e, c)). II.8′ Combining the scattered pieces of our transformed interpreter, we have:

interpret = λr. eval(r, mk-init(), \underline{λa. a}) II.1′ eval = λ(r, e, c). II.2′ (const?(r) → c(evcon(r)), II.3′ var?(r) → c(get(e, r)), II.4′ appl?(r) → eval(opr(r), e, \underline{λf. eval(opnd(r), e, \underline{λa. apply(f, a, c)})}), II.5′

lambda?(r) → c(mk-closr(r, e)), II.6′ cond?(r) → eval(prem(r), e, II.7′ \underline{λb. if b then eval(conc(r), e, c) else eval(altr(r), e, c)}), II.8′

letrec?(r) → eval(body(r), mk-rec(r, e), c)) II.9′ apply = λ(f, a, c). II.10′ (closr?(f) → II.11′ eval(body(lam(f)), mk-simp(fp(lam(f)), a, en(f)), c), II.12′ sc?(f) → c(succ(a)), II.13′ eq1?(f) → c(mk-eq2 (a)), II.14′ eq2?(f) → c(equal(arg1 (f), a))) II.15′ get = same as in Interpreter II. II.16–20

At this stage, since continuations are functional arguments, we have achieved order-of-application independence at the price of re-introducing higher-order functions. Fortunately, we can now “defunctionalize” the set CONT in the same way as FUNVAL and ENV. To interpret the new members of CONT we introduce a function cont such that if c_{new} represents the continuation c_{old} and a is a member of VAL then cont(c_{new}, a) = c_{old}(a). The application of cont must be introduced at each point in eval and apply where a continuation is applied to a value, i.e., in lines II.3′, 4′, 6′, 13′, 14′, and 15′.

There are four lambda expressions, indicated by solid underlining, that create continuations. The following table gives their locations and global variables, and the equations defining the new sets of records that will represent their values:

Location  Global Variables New Record Equation
II.1′  none  FIN = [ ]
II.5′ (outer) r e c  EVOPN = [ap: APPL, en: ENV, next: CONT]
II.5′ (inner) f c  APFUN = [fun: VAL, next: CONT]
II.8′  r e c  BRANCH = [cn: COND, en: ENV, next: CONT]

By replacing these lambda expressions by record-creation operations and moving their bodies into the new function cont (within let expressions that rebind their formal parameters and global variables appropriately), we obtain a third interpreter, which is independent of order-of-application and does not use higher-order functions:

CONT = FIN ∪ EVOPN ∪ APFUN ∪ BRANCH FIN = [ ] EVOPN = [ap: APPL, en: ENV, next: CONT] APFUN = [fun: VAL, next: CONT] BRANCH = [cn: COND, en: ENV, next: CONT] FUNVAL, ENV, etc. = same as in Interpreter II. interpret = λr. eval(r, mk-init(), mk-fin()) eval = λ(r, e, c).

(const?(r) → cont(c, evcon(r)), var?(r) → cont(c, get(e, r)), appl?(r) → eval(opr(r), e, mk-evopn(r, e, c)), lambda?(r) → cont(c, mk-closr(r, e)), cond?(r) → eval(prem(r), e, mk-branch(r, e, c)), letrec?(r) → eval(body(r), mk-rec(r, e), c)) III

apply = λ(f, a, c).

(closr?(f) →

eval(body(lam(f)), mk-simp(fp(lam(f)), a, en(f)), c), sc?(f) → cont(c, succ(a)), eq1?(f) → cont(c, mk-eq2 (a)), eq2?(f) → cont(c, equal(arg1 (f), a)))

cont = λ(c, a).

(fin?(c) → a, evopn?(c) → let f = a and r = ap(c) and e = en(c) and c = next(c)

in eval(opnd(r), e, mk-apfun(f, c)), apfun?(c) → let f = fun(c) and c = next(c) in apply(f, a, c), branch?(c) → let b = a and r = cn(c) and e = en(c) and c = next(c)

in if b then eval(conc(r), e, c) else eval(altr(r), e, c))

get = same as in Interpreter II.

From their abstract syntax, it is evident that continuations in our third interpreter are linear lists, with the unique record in FIN acting as the empty list, and the next fields in the other records acting as link fields. In effect, a continuation is a list of instructions to be interpreted by the function cont. Each instruction accepts a “current value” (the second argument of cont) and produces a new value that will be given to the next instruction. The following list gives approximate meanings for each type of instruction:

FIN: The current value is the final value of the program. Halt.

EVOPN: The current value is the value of an operator. Evaluate the operand of the application expression in the ap field, using the environment in the en field. Then obtain a new value by applying the current value to the value of the operand.

APFUN: The current value is the value of an operand. Obtain a new value by applying the function stored in the fun field to the current value.

BRANCH: The current value is the value of a premiss. If it is true (false) obtain a new value by evaluating the conclusion (alternative) of the conditional expression stored in the cn field, using the environment in the en field.

Each of the three serious functions, eval, apply, and cont, does a branch on the form of its first argument, performs trivial operations such as field selection, record creation, and environment lookup, and then calls another serious function. Thus our third interpreter is actually a state-transition machine, whose states each consist of the name of a serious function plus a list of its arguments.

This interpreter is similar in style to Landin’s SECD machine [7], though there is considerable difference in detailed mechanisms. (Very roughly, one can construct the continuation by merging Landin’s stack and control and concatenating this merged stack with the dump.)

8. Continuations with Higher-Order Functions

In transforming Interpreter I into Interpreter III, we have moved from a concise, abstract definition to a more complex machine-like one. If clarity consists of the avoidance of subtle characteristics of the defining language, then Interpreter III is certainly clearer than Interpreter I. But if clarity consists of conciseness and the absence of unnecessary complexity, then the reverse is true. The machine-like character of Interpreter III includes a variety of “cogs and wheels” that are quite arbitrary, i.e., one can easily construct equivalent interpreters (such as the SECD machine) with different cogs and wheels.

In fact, these “cogs and wheels” were introduced when we defunctionalized the sets FUNVAL, ENV, and CONT, since we replaced the functions in these sets by representations that were correct, but not unique. Had we chosen different representations, we would have obtained an equivalent but quite different interpreter.

This suggests the desirability of retaining the use of higher-order functions, providing these entities can be given a mathematically rigorous definition that is independent of any specific representation. Fortunately, such a definition has recently been provided by D. Scott’s new theory of computation [12, 13, 14, 15], which is based on concepts of lattice theory and topology. (The central technical problem that Scott has solved is to define functions that are not only higher-order, but also typeless, so that any function may be applied to any other function, including itself.) Although a description of this work would be beyond the scope of this paper, we may summarize its main implication for definitional interpreters: Scott has developed a mathematical model of the lambda calculus, which is thereby a model for a purely applicative higher-order defining language. But the defining language modelled by Scott uses call by name rather than call by value. (In terms of the lambda calculus, it uses normal order of evaluation.) Thus to apply Scott’s work to a defined language that uses call by value, we need a definitional interpreter that retains higher-order functions but is order-of-application independent.

An obvious approach to this goal is to introduce continuations directly into the meta-circular interpreter. At first sight, this appears to be straightforward. Referring back to Interpreter I, we see that the function eval is obviously serious, while evlambda, ext and initenv are trivial. (evlambda is trivial since the evaluation of lambda expressions always terminates.) Apparently eval is the only function that must accept continuations.

But when we transform the branch of eval that evaluates application expressions, the construction described in the previous section seems to give:

appl?(r) → eval(opr(r), e, λf. eval(opnd(r), e, λa. c(f(a)))).

Unfortunately, the subexpression c(f(a)) is not independent of the order-of-appli-cation, since the evaluation of the operand f(a) may never terminate, while the function c may be independent of its argument.

The difficulty is that the class of serious functions must include every potentially nonterminating function that may be applied during the execution of the interpreter; in addition to eval, this class contains the members of the set FUNVAL of defined-language functional values. Thus we must modify the functions in FUNVAL to accept continuations:

FUNVAL = VAL, CONT → VAL,

replacing each function f_{old} by an f_{new} such that f_{new}(a, c) = c(f_{old}(a)). This allows us to replace the order-dependent expression c(f(a)) by the order-independent expression f(a, c). Of course, we must add continuations as an extra formal parameter to each lambda expression that creates a member of FUNVAL.

(A similar modification of the functions in ENV is unnecessary, since it can be shown that the functions in this set always terminate. Just as with get, this depends on the exclusion of recursive let expressions with arbitrary declaring subexpressions.)

Once the necessity of altering FUNVAL has been realized, the transformation of Interpreter I follows the basic lines described in the previous section. We omit the details and state the final result:

VAL = INTEGER ∪ BOOLEAN ∪ FUNVAL FUNVAL = VAL, CONT → VAL ENV = VAR → VAL CONT = VAL → VAL interpret = λr. eval(r, initenv, λa. a)

eval = λ(r, e, c).

(const?(r) → c(evcon(r)),

var?(r) → c(e(r)),

appl?(r) → eval(opr(r), e, λf. eval(opnd(r), e, λa. f(a, c))),

lambda?(r) → c(evlambda(r, e)), IV cond?(r) → eval(prem(r), e,

λb. if b then eval(conc(r), e, c) else eval(altr(r), e, c)), letrec?(r) → letrec e′ =

λx. if x = dvar(r) then evlambda(dexp(r), e′) else e(x) in eval(body(r), e′, c))

evlambda = λ(ℓ, e). λ(a, c). eval(body(ℓ), ext(fp(ℓ), a, e), c)

ext = λ(z, a, e). λx. if x = z then a else e(x) initenv = λx. (x = “succ” → λ(a, c). c(succ(a)),

x = “equal” → λ(a, c). c(λ(b, c′). c′(equal(a, b)))).

This is basically the form of interpreter devised by L. Morris [20] and Wadsworth. It is almost as concise as the meta-circular interpreter, yet it offers the advantages of order-of-application independence and, as we will see in the next section, extensibility to accommodate imperative control features.

(The zealous reader may wish to verify that defunctionalization and the introduction of continuations are commutative, i.e., by replacing FUNVAL, ENV, and CONT by appropriate nonfunctional representations, one can transform Interpreter IV into Interpreter III.)

9. Escape Expressions

We now turn to the problem of adding imperative features to the defined language (while keeping the defining language purely applicative). These features may be divided into two classes:

1. Imperative control mechanisms, e.g., statement sequencing, labels and jumps.

2. Assignment.

We will first introduce control mechanisms and then consider assignment.

At first sight, this order of presentation seems facetious. In a language without assignment, it seems pointless to jump to a label, since there is no significant way for the part of the computation before the jump to influence the part afterwards. However, in Reference [29], Landin introduced an imperative control mechanism that is more general than labels and jumps, and that significantly enhances the power of a language without assignment. The specific mechanism that he introduced was called a J-operator, but in this paper we will develop a slightly simpler mechanism called an escape expression.

If (in the defined language) x is a variable and r is an expression, then

escape x in r

is an escape expression, whose escape variable is x and whose body is r. The evaluation of an escape expression in an environment e proceeds as follows:

1. The body r is evaluated in the environment that is the extension of e that binds x to a function called the escape function.

2. If the escape function is never applied during the evaluation of r, then the value of r becomes the value of the escape expression.

3. If the escape function is applied to an argument a, then the evaluation of the body r is aborted, and a immediately becomes the value of the escape expression.

Essentially, an escape function is a kind of label, and its application is a kind of jump. The greater generality lies in the ability to pass arguments while jumping.

(Landin’s J-operator can be defined in terms of the escape expression by regarding let g = J λx. r_{1} in r_{0} as an abbreviation for escape h in let g = λx. h(r_{1}) in r_{0}, where h is a new variable not occurring in r_{0} or r_{1}. Conversely, one can regard escape g in r as an abbreviation for let g = J λx. x in r.)

In order to extend our interpreters to handle escape expressions, we begin by extending the abstract syntax of expressions appropriately:

EXP = … ∪ ESCP

ESCP = [escv: VAR, body: EXP].

It is evident that in each interpreter we must add a branch to eval that evaluates the new kind of expression.

First consider Interpreter IV. Since an escape expression is evaluated by evaluating its body in an extended environment that binds the escape variable to the escape function, and since the escape function must be represented by a member of the set FUNVAL = VAL, CONT → VAL, we have

eval = λ(r, e, c). ( …,
escp?(r) → eval(body(r), ext(escv(r), λ(a, c′)…. , e), c)),

where the value of λ(a, c′)…. must be the member of FUNVAL representing the escape function.

Since eval is a serious function, its result, which is obtained by applying the continuation c to the value of the escape expression, must be the final result of the entire program being interpreted. This means that c itself must be a function that will accept the value of the escape expression and carry out the interpretation of the remainder of the program. But the member of FUNVAL representing the escape function is also serious, and must therefore also produce the final result of the entire program. Thus to abort the evaluation of the body and treat the argument a as the value of the escape expression, it is only necessary for the escape function ignore its own continuation c′, and to apply the higher-level continuation c to a. Thus we have:

eval = λ(r, e, c). ( …,

escp?(r) → eval(body(r), ext(escv(r), λ(a, c′). c(a), e), c)).

The extension of Interpreter III is essentially similar. In this case, we must add to the set FUNVAL a new kind of record that represents escape functions:

FUNVAL = … ∪ ESCF ESCF = [cn: CONT].

These records are created in the new branch of eval:

eval = λ(r, e, c). ( …, escp?(r) → eval(body(r), mk-simp(escv(r), mk-escf (c), e), c)),

and are interpreted by a new branch of apply:

apply = λ(f, a, c). ( …,

escf?(f) → cont(cn(f), a)).

From the viewpoint of this interpreter, it is clear that the escape expression is a significant extension of the defined language, since it introduces the possibility of embedding continuations in values.

(The reader should be warned that either of the above interpreters is a more precise definition of the escape expression than the informal English description given beforehand. For example, it is possible that the evaluation of the body of an escape expression may not cause the application of the escape function, but may produce the escape function (or some function that can call the escape function) as its value. It is difficult to infer the consequences of such a situation from our informal description, but it is precisely defined by either of the interpreters. In fact, the possibility that an escape function may propagate outside of the expression that created it is a powerful facility that can be used to construct control-flow mechanisms such as coroutines and nondeterministic algorithms.)

When we consider Interpreters I and II, we find an entirely different situation. The ability to “jump” by switching continuations is no longer possible. An escape function must still be represented by a member of FUNVAL, but now this implies that, if the function terminates without an error stop, then its result must become the value of the application expression that applied the function. As far as is known, there is no way to define the escape expression by adding branches to Interpreter I or II (except by the “cheat” of adding imperative control mechanisms to the defining language, as in Reference [19]). The essential problem is that the information that was explicitly available in the continuations of Interpreters III and IV is implicit in the recursive structure of Interpreters I and II, and in this form it cannot be manipulated with sufficient flexibility.

We have asserted that the escape mechanism encompasses less general control mechanisms such as labels and jumps. The following description outlines the way in which these more specialized operations can be expressed in terms of the escape expression. (A more detailed exposition is given in Reference [29].)

1. In the next section we will introduce assignment in such a way that assignments can be executed during the evaluation of expressions. In this situation it is unnecessary to make a semantic distinction between expressions and statements; any statement can be regarded as an expression whose evaluation produces a dummy value.

2. A label-free sequence of statements s_{1}; ⋯; s_{n} can be regarded as an abbreviation for the expression

( ⋯ ((λx_{1}…. λx_{n}. x_{n})(s_{1})) ⋯ (s_{n})).

The effect is to evaluate the statements sequentially from left to right, ignoring the value of all but the last.

3. If s_{0}, …, s_{n} are label-free statement sequences, and ℓ_{1}, …, ℓ_{n} are labels, then a block of the form

begin s_{0}, ℓ_{1}: s_{1}; ⋯; ℓ_{n}: s_{n} end

can be regarded as an abbreviation for

escape g in letrec ℓ_{1} = λx. g(s_{1}; ⋯; s_{n}) and ℓ_{2} = λx. g(s_{2}; ⋯; s_{n})

and ⋯ and ℓ_{n} = λx. g(s_{n}) in (s_{0}; ⋯; s_{n})

(where g and x are new variables not occurring in the original block). The effect is that each label denotes a function that ignores its argument, evaluates the appropriate sequence of statements, and then escapes out of the enclosing block.

4. An expression of the form goto r can be regarded as an abbreviation for r(0), i.e., a jump to a label becomes an application of the function denoted by the label to a dummy argument.

10. Assignment

Although the basic concept of assignment is well understood by any competent programmer, a surprising degree of care is needed to combine this concept with the language features we have discussed previously. Intuitively, the notion of assignment presupposes that the operations that are performed during the evaluation of a program will occur in a definite temporal order. Some of these operations will assign values to “variables”. Other operations may be affected by these assignments; specifically, an operation may depend upon the value most recently assigned to each “variable”, which we will call the value currently possessed by the “variable”.

This suggests that for each instant during program execution, there should be an entity which specifies the set of “variables” that are present and the values that they currently possess. We will call such an entity a memory, and denote the set of possible memories by MEM.

The main subtlety is to realize that the “variables” discussed here are distinct from the variables used in previous sections. This is necessitated by the fact that most programming languages permit situations (such as might arise from the use of “call by address”) in which several variables denote the same “variable”, in the sense that assignment to one of them will change the value possessed by all. This suggests that a “variable” is actually a new kind of object to which a variable can be bound. Henceforth, we will call these new objects references rather than “variables”. (Other terms used commonly in the literature are L-value and name.) We will denote the set of references by REF.

Abstractly, the nature of references and memories can be characterized by specifying an initial memory and four functions:

initmem: Contains no references.

nextref (m): Produces a reference not contained in the memory m.

augment(m, a): Produces a memory containing the new reference nextref (m) plus the references already in m. The new reference possesses the value a, while the remaining references possess the same values as in m.

update(m, rf , a): Produces a memory containing the same references as m. The reference rf (assuming it is present) possesses the value a, while the remaining references possess the same value as in m.

lookup(m, rf ): Produces the value possessed by the reference rf in memory m.

A simple “implementation” can be obtained by numbering references in the order of their creation [25]:

REF = [number: INTEGER]

MEM = [count: INTEGER, possess: INTEGER → VAL] initmem = mk-mem(0, λn. 0) nextref = λm. mk-ref (count(m) + 1) augment = λ(m, a). mk-mem(count(m) + 1,

λn. if n = count(m) + 1 then a else (possess(m))(n)) update = λ(m, rf , a). mk-mem(count(m),

λn. if n = number(rf ) then a else (possess(m))(n)) lookup = λ(m, rf ). (possess(m))(number(rf )).

Our next task is to introduce memories into our interpreters. Although any of our interpreters could be so extended, we will only consider Interpreter IV.

It is evident that the operation of evaluating a defined-language expression will now depend upon a memory m and will produce a (possibly) altered memory m′. Thus the function eval will accept m as an additional argument. However, because of the use of continuations, m′ will not be part of the result of eval. Instead, m′ will be passed on as an additional argument to the continuation that is applied by eval to perform the remainder of program execution.

In a similar manner, the application of a defined-language function will depend upon and produce memories. Thus each function in the set FUNVAL will accept a memory as an additional argument, and will also pass on a memory to its continuation.

On the other hand, there are particular kinds of expressions, specifically constants, variables, and lambda expressions, whose evaluation cannot cause assignments. For this reason, the functions evcon and evlambda, and the functions in the set ENV, will not accept or produce memories.

These considerations lead to the following interpreter, in which memories propagate through the various operations in a manner that correctly reflects the temporal order of execution:

VAL = INTEGER ∪ BOOLEAN ∪ FUNVAL FUNVAL = VAL, MEM, CONT → VAL ENV = VAR → VAL CONT = MEM, VAL → VAL interpret = λr. eval(r, initenv, initmem, λ(m, a). a) eval = λ(r, e, m, c).

(const?(r) → c(m, evcon(r)),

var?(r) → c(m, e(r)),

appl?(r) → eval(opr(r), e, m,

λ(m′, f). eval(opnd(r), e, m′,
λ(m′′, a). f(a, m′′, c))),

lambda?(r) → c(m, evlambda(r, e)), cond?(r) → eval(prem(r), e, m,

λ(m′, b). if b then eval(conc(r), e, m′, c) else eval(altr(r), e, m′, c)), letrec?(r) → letrec e′ =

λx. if x = dvar(r) then evlambda(dexp(r), e′) else e(x) in eval(body(r), e′, m, c),

escp?(r) → eval(body(r), ext(escv(r), λ(a, m′, c′). c(m′, a), e), m, c))

evlambda = λ(ℓ, e). λ(a, m, c). eval(body(ℓ), ext(fp(ℓ), a, e), m, c)
ext = λ(z, a, e). λx. if x = z then a else e(x)
initenv = λx. (x = “succ” → λ(a, m, c). c(m, succ(a)),

x = “equal” → λ(a, m, c). c(m, λ(b, m′, c′). c′(m′, equal(a, b)))).

At this stage, although we have “threaded” memories through the operations of our interpreter, we have not yet introduced references, nor any operations that alter or depend upon memories. To proceed further, however, we must distinguish between two approaches to assignment, each of which characterizes certain programming languages.

In the “L-value” approach, in each context of the evaluation process where a value would occur, a reference (i.e., L-value) possessing that value occurs instead. Thus, for example, expressions evaluate to references, functional arguments and results are references, and environments bind variables to references. (In richer languages, references would occur instead of values in still other contexts, such as array elements.) This approach is used in the languages PAL [3] and ISWIM [2], and in somewhat modified form (i.e., references always occur in certain kinds of contexts, while values always occur in others) in such languages as FORTRAN, ALGOL 60, and PL/I. Its formalization is due to Strachey [30], and is used extensively in the Vienna definition of PL/I [18].

In the “reference” approach, references are introduced as a new kind of value, so that either references or “normal” values can occur in any meaningful context. This approach is used in ALGOL 68 [31], BASIL [32], and GEDANKEN [4].

The relative merits of these approaches are discussed briefly in Reference [4]. Although either approach can be accommodated by the various styles of interpreter discussed in this paper, we will limit ourselves to incorporating the reference approach into the above extension of Interpreter IV. We first augment the set of values appropriately:

VAL = INTEGER ∪ BOOLEAN ∪ FUNVAL ∪ REF.

Next we introduce basic operations for creating, assigning, and evaluating references. For simplicity, we will make these operations basic functions, denoted by the predefined variables ref, set, and val. The following is an informal description:

ref (a): Accepts a value a and returns a new reference initialized to possess a.

(set(rf ))(a): Accepts a reference rf and a value a. The value a is assigned to rf and also returned as the result. (Because of our restriction to functions of a single argument, this function is Curried, i.e., set accepts rf and returns a function that accepts a.)

val(rf ): Accepts a reference rf and returns its currently possessed value.

To introduce these new functions into our interpreter, we extend the initial environment as follows:

initenv = λx. ( ⋯
x = “ref” → λ(a, m, c). c(augment(m, a), nextref (m)),
x = “set” → λ(rf , m, c). c(m, λ(a, m′, c′). c′(update(m′, rf , a), a)),
x = “val” → λ(rf , m, c). c(m, lookup(m, rf ))).

The main shortcoming of the reference approach is the incessant necessity of using the function val. This problem can be alleviated by introducing coercion conventions, as discussed in Reference [4], that cause references to be replaced by their possessed values in appropriate contexts. However, since these conventions can be treated as abbreviations, they do not affect the basic structure of the definitional interpreters.

11. Directions Of Future Research

Within this paper we have tried to present a systematic, self-contained, and reasonably complete description of the current state of the art of definitional interpreters. We conclude with a brief (and hopeful) list of possible future developments:

1. It would still be very desirable to be able to define higher-order languages logically rather than interpretively, particularly if such an approach can lead to practical correctness proofs for programs. A major step in this direction, based on the work of Scott [12, 13, 14, 15], has been taken by R. Milner [16]. However, Milner’s work essentially treats a language using call by name rather than call by value.

2. It should be possible to treat languages with multiprocessing features, or other features that involve “controlled ambiguity”. An initial step is the work of the IBM Vienna Laboratory [18], using a nondeterministic state-transition machine.

3. It should also be possible to define languages, such as ALGOL 68 [31], with a highly refined syntactic type structure. Ideally, such a treatment should be meta-circular, in the sense that the type structure used in the defined language should be adequate for the defining language.

4. The conciseness of definitional interpreters makes them powerful tools for language design, particularly when one wishes to add new capabilities to a language with a minimum of increased complexity. Of particular interest (at least to the author) are the problems of devising better type systems and of generalizing assignment (for example, by permitting memories to be embedded in values.)

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