So the focus of this paper is to define the way a programming langu...
A nice way to point out the difference between applicative and impe...
This is a useful definition: A purely applicative language is on...
These citations are, I'll recommend Floyd for a future reading grou...
This is the paper we read on week 1!
Notice how each language construct of the object language is implem...
It's known how to do this now, monads
In this section a systematic program transformation technique calle...
In this section the Continuation Passing Style transform is applied.
This is the key part where using CPS lets us express the evaluation...
Higher-Order and Symbolic Computation, 11, 363–397 (1998)
c
°
1998 Kluwer Academic Publishers, Boston. Manufactured in The Netherlands.
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unpublished work by L. Morris and C. Wadsworth.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 forthe 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 triedto 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 of functions,
and imperative features, such as statement sequencing, labels, jumps, assignment, and
*
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
364 REYNOLDS
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 evento be “syntactically sugared” versionsof the lambda calculus.
Inparticular, Landin [7] hasshown that suchlanguagescanbereducedtothe 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 first-order 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
DEFINITIONAL INTERPRETERS 365
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.
Thesecond criterion involvesthenotion oforder ofapplication. In designing anylanguage
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 inter-
preters. 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 higher-order 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).
366 REYNOLDS
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
0
are
environments. If e
0
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
0
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.
DEFINITIONAL INTERPRETERS 367
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
maybe omitted if there is only one formal parameter.) The evaluation of a lambda expression
with n formal parameters always terminates and alwaysproduces 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
).rin 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 extensionof the earlierenvironmentin which the lambda expression was evaluated (rather
than the more recentenvironment in which the application takes place). As aconsequence, if
a lambda expression contains global variables (i.e., variablesthat 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 environ-
ments 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 distinc-
tion 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 oper-
ator 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
368 REYNOLDS
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 al-
ways 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 defin-
ing 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:
DEFINITIONAL INTERPRETERS 369
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 +1and 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 =0then 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, `
1
,...,`
n
are lambda expressions, and r
b
is an expression, then
letrec x
1
= `
1
and ···and x
n
= `
n
in r
b
is a recursive let expression, whose declared variables are x
1
,...,x
n
, whose declaring
expressions are `
1
,...,`
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
0
which satisfies
the following property: e
0
is the extension of e that binds each declared variable x
i
to the
function obtained by evaluating the corresponding declaring lambda expression `
i
in the
environment e
0
.
There is a circularity in the property e
0
is the ... in the environment e
0
that is char-
acteristic of recursion, and that prevents this property from being an explicit definition of
e
0
. 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
370 REYNOLDS
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 per-
form 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 decla-
rations. 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 oper-
ations 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].
DEFINITIONAL INTERPRETERS 371
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
372 REYNOLDS
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),
DEFINITIONAL INTERPRETERS 373
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. Arecursive letexpression(a member of LETREC), which consists of a variablecalled 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.
ForVAR, 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 defining-language 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 theset ENV of environments. Since the purposeof 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 ee
0
ENV
xz VAR cc
0
CONT
` LAMBDA mm
0
m
00
MEM
ab VAL rf REF
f FUNVAL n INTEGER
(The sets CONT, MEM, and REF will be defined later.)
374 REYNOLDS
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
0
= I.8
λx. if x = dvar(r) then evlambda(dexp(r),e
0
)else e(x) I.9
in eval(body(r),e
0
)
¢
I.10
evlambda = λ(`, e). λa. eval
¡
body(`), ext(fp(`),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 ` 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
0
is handled by making e
0
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” λa. succ(a)
,
...
I.14
x = “equal” λa. 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,
“Sch¨onfinkeling” 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).
DEFINITIONAL INTERPRETERS 375
(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 higher-order
functions. For this purpose, we would prefer an interpreter of a higher-order defined
language that was written in a first-order defining language.
2. Changing the order of application used in the defining languageinduces 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)
¢
inthe defining language. Ifcallbyvalueisusedin 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)interpretsthe evaluationof r
1
, andapplying 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 valueof 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
376 REYNOLDS
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 ap-
plication 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 exam-
ination 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
0
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
DEFINITIONAL INTERPRETERS 377
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 ` 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 ` and e:
evlambda = λ(`, e). mk-closr(`, e). I.11
0
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(`), ext(fp(`),a,e)
¢
of the replaced lambda expression in an appropriate environment. This environment must
bindthe formal parameter a of the replaced lambda expressionto the valueof a and must bind
the global variables ` 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,wehave:
apply = λ(f, a).
¡
closr?(f) let a = a and ` = lam(f ) and e = en(f )
in eval
¡
body(`), ext(fp(`),a,e)
¢
,
...
¢
.
378 REYNOLDS
(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
0
x = “equal” mk-eq1()
¢
...
I.15
0
and
apply = λ(f, a).
¡
closr?(f) let ` = lam(f) and e = en(f )
in eval
¡
body(`), ext(fp(`),a,e)
¢
,
sc?(f) succ(a),
eq1?(f) 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 ` = lam(f) and e = en(f )
in eval
¡
body(`), ext(fp(`),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 FUNVALfrom 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 interpretive function get, with
the property that if e
new
represents an environment e
old
and x is a member of VAR, then
DEFINITIONAL INTERPRETERS 379
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
0
.
.
.
λx. if x = dvar(r) then evlambda(dexp(r),e
0
)else get(e, x) I.9
0
.
.
.
ext = λ(z, a, e). λx. if x = z then a else get(e, x)
. I.12
0
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
0
-15
0
none INIT = []
I.12
0
zae SIMP = [bvar: VAR, bval: VAL, old: ENV]
I.9
0
ree
0
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
0
= mk-rec(r, e, e
0
) ··· I.8-9
00
.
.
.
ext = λ(z, a, e). mk-simp(z, a, e) I.12
00
.
.
.
initenv = mk-init(), I.14
00
-15
00
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
0
= new(e)
in if x = dvar(r) then evlambda(dexp(r),e
0
)else get(e, x)
¢
.
But now we are faced with a new problem. By eliminating the lambda expression in I.9
0
,
we have created a recursive let expression
380 REYNOLDS
letrec e
0
= mk-rec(r, e, e
0
) ···
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
0
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
0
= 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
0
= mk-rec(r, e) ··· . I.8
0
-9
000
At this point, once we have collected the bits and pieces produced by the various trans-
formations, 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
0
can be eliminated by replacing occurrences of evlambda by mk-closr.
3. Line I.12
00
can be eliminated by replacing occurrences of ext by mk-simp.
4. Lines I.14
00
-15
00
can be eliminated by replacing occurrences of initenv by mk-init().
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]
DEFINITIONAL INTERPRETERS 381
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 recursivecharacteristic that the declaring subexpressionshould “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
382 REYNOLDS
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 +1in ···.)
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 elimi-
nated 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 declaringsubexpressionsof letexpressions),
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 well-known 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
nontrivialprogram. 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.
DEFINITIONAL INTERPRETERS 383
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 continua-
tions 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 VA L.
(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
0
eval = λ(r, e, c). ··· II.2
0
apply = λ(f, a, c). ··· II.10
0
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:
384 REYNOLDS
eval = λ(r, e, c). II.2
0
¡
const?(r) c(evcon(r)), II.3
0
var?(r) c(get(e, r)), II.4
0
.
.
.
lambda?(r) c(mk-closr(r, e)), ...
¢
II.6
0
apply = λ(f, a, c).
¡
... , II.10
0
sc?(f) c(succ(a)), II.13
0
eq1?(f) c(mk-eq2(a)), II.14
0
eq2?(f) c(equal(arg1(f ),a))
¢
. II.15
0
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
0
.
.
.
¡
closr?(f) II.11
0
eval
¡
body(lam(f)), mk-simp(fp(lam(f)),a,en(f)),c
¢
. II.12
0
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),ef. eval
¡
opnd(r),ea.apply(f, a, c)
¢¢
. II.5
0
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:
DEFINITIONAL INTERPRETERS 385
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
0
λb. if b then eval(conc(r),e,c) else eval(altr(r),e,c)
¢
. II.8
0
Combining the scattered pieces of our transformed interpreter, we have:
interpret = λr. eval(r, mk-init(), λa. a
) II.1
0
eval = λ(r, e, c). II.2
0
¡
const?(r) c(evcon(r)), II.3
0
var?(r) c(get(e, r)), II.4
0
appl?(r) eval
¡
opr(r),ef. eval
¡
opnd(r),ea.apply(f, a, c)
¢¢
, II.5
0
lambda?(r) c(mk-closr(r, e)), II.6
0
cond?(r) eval
¡
prem(r),e, II.7
0
λb. if b then eval(conc(r),e,c) else eval(altr(r),e,c)
¢
, II.8
0
letrec?(r) eval(body(r), mk-rec(r, e),c)
¢
II.9
0
apply = λ(f, a, c). II.10
0
¡
closr?(f) II.11
0
eval
¡
body(lam(f)), mk-simp(fp(lam(f)),a,en(f)),c
¢
, II.12
0
sc?(f) c(succ(a)), II.13
0
eq1?(f) c(mk-eq2(a)), II.14
0
eq2?(f) c(equal(arg1(f ),a))
¢
II.15
0
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
interpretthe newmembersof CONT weintroducea function cont such that ifc
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
0
,4
0
,6
0
,13
0
,14
0
, and 15
0
.
There are four lambda expressions, indicated by solid underlining, that create continu-
ations. The following table gives their locations and global variables, and the equations
defining the new sets of records that will represent their values:
386 REYNOLDS
Location Global Variables New Record Equation
II.1
0
none FIN = []
II.5
0
(outer) rec EVOPN = [ap: APPL, en: ENV, next: CONT]
II.5
0
(inner) fc APFUN = [fun: VAL, next: CONT]
II.8
0
rec 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.
DEFINITIONAL INTERPRETERS 387
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 appli-
cation 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 consid-
erable 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 com-
plexity, 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
388 REYNOLDS
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),ef. eval
¡
opnd(r), e, λa. c(f (a))
¢¢
.
Unfortunately, the subexpression c(f (a)) is not independent of the order-of-application,
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 nonter-
minating 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 Inter-
preter I follows the basic lines described in the previous section. We omit the details and
state the final result:
DEFINITIONAL INTERPRETERS 389
VAL = INTEGER BOOLEAN FUNVAL
FUNVAL = VAL, CONT VAL
ENV = VAR VA L
CONT = VAL VA L
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),ef. 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
0
=
λx. if x = dvar(r) then evlambda(dexp(r),e
0
)else e(x)
in eval(body(r),e
0
,c)
¢
evlambda = λ(`, e)(a, c). eval
¡
body(`), ext(fp(`),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
0
).c
0
(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 accommo-
date 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 languagewithout 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
390 REYNOLDS
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)(a, c
0
). ...,e),c
¢¢
,
where the value of λ(a, c
0
). ...must be the member of FUNVAL representing the escape
function.
DEFINITIONAL INTERPRETERS 391
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
0
, 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)(a, c
0
).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 signif-
icant 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
392 REYNOLDS
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 mecha-
nisms 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 `
1
,...,`
n
are labels, then a block
of the form
begin s
0
,`
1
:s
1
; ··· ;`
n
:s
n
end
can be regarded as an abbreviation for
escape g in letrec `
1
= λx. g(s
1
; ··· ;s
n
)and `
2
= λx. g(s
2
; ··· ;s
n
)
and ··· and `
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
DEFINITIONAL INTERPRETERS 393
operations that are performed during the evaluation of a program will occur in a definite
temporal order. Some of these operationswill assign valuesto “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 refer-
ence 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)+1then 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 )).
394 REYNOLDS
Our next task is to introduce memories into our interpreters. Although any of our inter-
preters 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
0
. Thus the
function eval will accept m as an additional argument. However, because of the use of
continuations, m
0
will not be part of the result of eval. Instead, m
0
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, vari-
ables, 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 pro-
duce 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 VA L
CONT = MEM, VAL VAL
interpret = λr. eval(r, initenv, initmem(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
0
,f). eval(opnd(r),e,m
0
,
λ(m
00
,a).f(a, m
00
,c))
¢
,
lambda?(r) c(m, evlambda(r, e)),
cond?(r) eval
¡
prem(r),e,m,
λ(m
0
,b).if b then eval(conc(r),e,m
0
,c)else eval(altr(r),e,m
0
,c)
¢
,
letrec?(r) letrec e
0
=
λx. if x = dvar(r) then evlambda(dexp(r),e
0
)else e(x)
in eval(body(r),e
0
,m,c),
escp?(r) eval
¡
body(r), ext(escv(r)(a, m
0
,c
0
).c(m
0
,a),e),m,c
¢¢
evlambda = λ(`, e)(a, m, c). eval
¡
body(`), ext(fp(`),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
0
,c
0
).c
0
(m
0
,equal(a, b))
¢¢
.
DEFINITIONAL INTERPRETERS 395
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
0
,c
0
).c
0
(update(m
0
, 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.
396 REYNOLDS
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. Itwould still be verydesirable to be able to define higher-orderlanguages 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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So the focus of this paper is to define the way a programming languages works, using an interpreter in a programming language. It's known how to do this now, monads This is a useful definition: A purely applicative language is one which can be entirely desugared into lambda calculus. There's a great blog post about doing this for lisp: http://matt.might.net/articles/compiling-up-to-lambda-calculus/ These citations are, I'll recommend Floyd for a future reading group week - [8] Floyd, Robert W. Assigning meanings to programs. In J. T. Schwartz, editor, Mathematical Aspects of Computer Science, volume 19 of Proceedings of Symposia in Applied Mathematics, pages 19–32, New York - [9] City, April 5–7, 1966. American Mathematical Society, Providence, Rhode Island, 1967.9. Manna, Zohar. The correctness of programs. Journal of Computer and System Sciences, 3(2):119–127,May 1969. - [10] Hoare, C. A. R. An axiomatic basis for computer programming. Communications of the ACM, 12(10):576– 580 and 583, October 1969. Reprinted in [11]. Notice how each language construct of the object language is implemented by interpreting it into the host language version of that construct. e.g. a lambda expression gets turned into an actual lambda. A conditional is implemented using 'if'. And so on. In this section the Continuation Passing Style transform is applied. This is the key part where using CPS lets us express the evaluation order of a function application A nice way to point out the difference between applicative and imperative languages could be if, in lisp we can do 1 + (if x y z) in C we can't, we'd need to do something like this if(x) { return 1+y; } else { return 1+z; } This is the paper we read on week 1! In this section a systematic program transformation technique called "defunctionalization" is performed.