Lambda calculus

The lambda calculus is a formal system designed to investigate function definition, function application, and recursion. It was introduced by Alonzo Church and Stephen Cole Kleene in the 1930s; Church used the lambda calculus in 1936 to give a negative answer to the Entscheidungsproblem. The calculus can be used to cleanly define what a computable function is. The question of whether two lambda calculus expressions are equivalent cannot be solved by a general algorithm, and this was the first question, even before the halting problem, for which undecidability could be proved. Lambda calculus has greatly influenced functional programming languages, especially Lisp (which was "accidentally used to define itself" (Abelson and Sussman)).
The lambda calculus can be called the smallest universal programming language. The lambda calculus consists of a single transformation rule (variable substitution) and a single function definition scheme. The lambda calculus is universal in the sense that any computable function can be expressed and evaluated using this formalism. It is thus equivalent to Turing machines. However, the lambda calculus emphasizes the use of transformation rules, and does not care about the actual machine implementing them. It is an approach more related to software than to hardware.
This article deals with the "untyped lambda calculus" as originally conceived by Church. Since then, some typed lambda calculi have been developed.
Contents 
History
Originally, Church had tried to construct a complete formal system for the foundations of mathematics; when the system turned out to be susceptible to the analog of Russell's paradox, he separated out the lambda calculus and used it to study computability, culminating in his negative answer to the Entscheidungsproblem.
Informal description
In lambda calculus, every expression stands for a function with a single argument; the argument of the function is in turn a function with a single argument, and the value of the function is another function with a single argument. A function is anonymously defined by a lambda expression which expresses the function's action on its argument. For instance, the "addtwo" function f(x) = x + 2 would be expressed in lambda calculus as λ x. x + 2 (or equivalently as λ y. y + 2; the name of the formal argument is immaterial) and the number f(3) would be written as (λ x. x + 2) 3. Function application is left associative: f x y = (f x) y. Consider the function which takes a function as argument and applies it to the argument 3: λ x. x 3. This latter function could be applied to our earlier "add2" function as follows: (λ x. x 3) (λ x. x+2). It is clear that the three expressions
 (λ x. x 3) (λ x. x+2) and (λ x. x + 2) 3 and 3 + 2
are equivalent. A function of two variables is expressed in lambda calculus as a function of one argument which returns a function of one argument (see currying). For instance, the function f(x, y) = x  y would be written as λ x. λ y. x  y. The three expressions
 (λ x. λ y. x  y) 7 2 and (λ y. 7  y) 2 and 7  2
are equivalent. It is this equivalence of lambda expressions which in general can not be decided by an algorithm.
Not every lambda expression can be reduced to a definite value like the ones above; consider for instance
 (λ x. x x) (λ x. x x)
or
 (λ x. x x x) (λ x. x x x)
and try to visualize what happens as you start to apply the first function to its argument. (λ x. x x) is also known as the ω combinator; ((λ x. x x) (λ x. x x)) is known as Ω, ((λ x. x x x) (λ x. x x x)) as Ω_{2}, etc.
While the lambda calculus itself does not contain symbols for integers or addition, these can be defined as abbreviations within the calculus and arithmetic can be expressed as we will see below.
Lambda calculus expressions may contain free variables, i.e. variables not bound by any λ. For example, the variable y is free in the expression (λ x. y), representing a function which always produces the result y. Occasionally, this necessitates the renaming of formal arguments, for instance in order to reduce
 (λ x. λ y. y x) (λ x. y) to λ z. z (λ x. y)
If one only formalizes the notion of function application and does not allow lambda expressions, one obtains combinatory logic.
Formal definition
Formally, we start with a countably infinite set of identifiers, say {a, b, c, ..., x, y, z, x_{1}, x_{2}, ...}. The set of all lambda expressions can then be described by the following contextfree grammar in BNF:
 <expr> ::= <identifier>
 <expr> ::= (λ <identifier> . <expr>)
 <expr> ::= (<expr> <expr>)
The first two rules generate functions, while the third describes the application of a function to an argument. Usually the brackets for lambda abstraction (rule 2) and function application (rule 3) are omitted if there is no ambiguity under the assumptions that (1) function application is leftassociative, and (2) a lambda binds to the entire expression following it. For example, the expression ((λ x. (x x)) (λ y. y)) can be simply written as (λ x. x x) λ y.y.
Lambda expressions such as λ x. (x y) do not define a function because the occurrence of the variable y is free, i.e., it is not bound by any λ in the expression. The binding of occurrences of variables is (with induction upon the structure of the lambda expression) defined by the following rules:
 In an expression of the form V where V is a variable this V is the single free occurrence.
 In an expression of the form λ V. E the free occurrences are the free occurrences in E except those of V. In this case the occurrences of V in E are said to be bound by the λ before V.
 In an expression of the form (E E' ) the free occurrences are the free occurrences in E and E' .
Over the set of lambda expressions an equivalence relation (here denoted as ==) is defined that captures the intuition that two expressions denote the same function. This equivalence relation is defined by the socalled alphaconversion rule and the betareduction rule.
αconversion
The alphaconversion rule is intended to express the idea that the names of the bound variables are unimportant; for example that λx.x and λy.y are the same function. However, the rule is not as simple as it first appears. There are a number of restrictions on when one bound variable may be replaced with another.
The alphaconversion rule states that if V and W are variables, E is a lambda expression, and E[V := W] means the expression E with every free occurrence of V in E replaced with W, then
 λ V. E == λ W. E[V := W]
if W does not appear freely in E and W is not bound by a λ in E whenever it replaces a V. This rule tells us for example that λ x. (λ x. x) x is the same as λ y. (λ x. x) y.
βreduction
The betareduction rule expresses the idea of function application. It states that
 ((λ V. E ) E' ) == E [V := E' ]
if all free occurrences in E' remain free in E [V := E' ].
The relation == is then defined as the smallest equivalence relation that satisfies these two rules.
A more operational definition of the equivalence relation can be given by applying the rules only from left to right. A lambda expression which does not allow any beta reduction, i.e., has no subexpression of the form ((λ V. E) E' ), is called a normal form. Not every lambda expression is equivalent to a normal form, but if it is, then the normal form is unique up to naming of the formal arguments. Furthermore, there is an algorithm for computing normal forms: keep replacing the first (leftmost) formal argument with its corresponding concrete argument, until no further reduction is possible. This algorithm halts if and only if the lambda expression has a normal form. The ChurchRosser theorem then states that two expressions result in the same normal form up to renaming of the formal arguments if and only if they are equivalent.
ηconversion
There is a third rule, etaconversion, which may be added to these two to form a new equivalence relation. Etaconversion expresses the idea of extensionality, which in this context is that two functions are the same iff they give the same result for all arguments. Etaconversion converts between λ x . f x and f, whenever x does not appear free in f. This can be seen to be equivalent to extensionality as follows:
If f and g are extensionally equivalent, i.e. if f a == g a for all lambda expressions a, then in particular by taking a to be a variable x not appearing free in f we have f x == g x and hence λ x . f x == λ x . g x, and so by etaconversion f == g. So if we take etaconversion to be valid, we find extensionality is valid.
Conversely if extensionality is taken to be valid, then since by betareduction for all y we have (λ x . f x) y == f y, we have λ x . f x == f  i.e. etaconversion is found to be valid.
Arithmetic in lambda calculus
There are several possible ways to define the natural numbers in lambda calculus, but by far the most common are the Church integers, which can be defined as follows:
 0 := λ f. λ x. x
 1 := λ f. λ x. f x
 2 := λ f. λ x. f (f x)
 3 := λ f. λ x. f (f (f x))
and so on. Intuitively, the number n in lambda calculus is a function that takes a function f as argument and returns the nth power of f. That is to say, a Church integer is a higherorder function  it takes a singleargument function f, and returns another singleargument function.
(Note that in Church's original lambda calculus, the formal parameter of a lambda expression was required to occur at least once in the function body, which made the above definition of 0 impossible.) Given this definition of the Church integers, we can define a successor function, which takes a number n and returns n + 1:
 SUCC := λ n. λ f. λ x. f (n f x)
Addition is defined as follows:
 PLUS := λ m. λ n. λ f. λ x. m f (n f x)
PLUS can be thought of as a function taking two natural numbers as arguments and returning a natural number; it is fun to verify that
 PLUS 2 3 and 5
are equivalent lambda expressions. Multiplication can then be defined as
 MULT := λ m. λ n. m (PLUS n) 0,
the idea being that multiplying m and n is the same as m times adding n to zero. Alternatively
 MULT = λ m. λ n. λ f. m (n f)
The predecessor PRED n = n  1 of a positive integer n is more difficult:
 PRED := λ n. λ f. λ x. n (λ g. λ h. h (g f)) (λ u. x) (λ u. u)
or alternatively
 PRED = λ n. n (λ g. λ k. (g 1) (λ u. PLUS (g k) 1) k) (λ l. 0) 0
Note the trick (g 1) (λ u. PLUS (g k) 1) k which evaluates to k if g(1) is zero and to g(k) + 1 otherwise.
Logic and predicates
By convention, the following two definitions (known as Church booleans) are used for the boolean values TRUE and FALSE:
 TRUE := λ u. λ v. u
 FALSE := λ u. λ v. v = 0
Then, with these two λterms, we can define some logic operators:
 AND := λ uv.u v FALSE
 OR := λ uv.u TRUE v
 NOT := λ b u v.b v u
 IFTHENELSE := λ b u v.b u v
We are now able to compute some logic functions, as for example:
 <math>\mbox{AND TRUE FALSE}<math>
 <math>\equiv (\lambda x y . x y\ \mathrm{FALSE})\ \mathrm{TRUE}\ \mathrm{FALSE} \to_{\beta} \mathrm{TRUE}\ \mathrm{FALSE}\ \mathrm{FALSE}<math>
 <math>\equiv (\lambda x y . x)\ \mathrm{FALSE}\ \mathrm{FALSE} \to_{\beta} \mathrm{FALSE}<math>
and we see that TRUE AND FALSE is equivalent to FALSE.
A predicate is a function which returns a boolean value. The most fundamental predicate is ISZERO which returns true if and only if its argument is zero:
 ISZERO := λ n. n (λ x. FALSE) TRUE
The availability of predicates and the above definition of TRUE and FALSE make it convenient to write "ifthenelse" statements in lambda calculus.
Recursion
Recursion is the definition of a function using the function itself; on the face of it, lambda calculus does not allow this. However, this impression is misleading. Consider for instance the factorial function f(n) recursively defined by
 f(n) = 1, if n = 0; and n·f(n1), if n>0.
In lambda calculus, one cannot define a function which includes itself. To get around this, one may start by defining a function, here called g, which takes a function f as an argument and returns another function that takes n as an argument:
 g := λ f. λ n. (1, if n = 0; and n·f(n1), if n>0).
The function that g returns either returns the constant 1, or returns n times the application of the function f to "n1". Using the ISZERO predicate, and boolean and algebraic definitions described above, the function g can be defined in lambda calculus.
However, g by itself is still not recursive; in order to use g to create the recursive factorial function, the function passed to g as f must have specific properties. Namely, the function passed as f must expand to the function g called with one argument  and that argument must be the function that was passed as f again!
In other words, f must expand to g(f). This call to g will then expand to the above factorial function and calculate down to another level of recursion. In that expansion the function f will appear again, and will again expand to g(f) and continue the recursion. This kind of function, where f = g(f), is called a fixedpoint of g, and it turns out that it can be implemented in the lambda calculus using what is known as the paradoxical operator or fixedpoint operator and is represented as Y  the Y combinator:
 Y = λ g. (λ x. g (x x)) (λ x. g (x x))
In the lambda calculus, Y g is a fixedpoint of g, as it expands to g (Y g). Now, to complete our recursive call to the factorial function, we would simply call g (Y g) n, where n is the number we are calculating the factorial of.
Given n = 5, for example, this expands to:
 λ n. (1, if n = 0; and n·(Y g)(n1), if n>0) 5
 1, if 5 = 0; and 5·(g(Y g))(51), if 5>0
 5·g(Y g) 4
 5·(λ n. (1, if n = 0; and n·(Y g)(n1), if n>0) 4)
 5·(1, if 4 = 0; and 4·(g(Y g))(41), if 4>0)
 5·(4·g(Y g) 3)
 5·(4·(λ n. (1, if n = 0; and n·(Y g)(n1), if n>0) 3))
 5·(4·(1, if 3 = 0; and 3·(g(Y g))(31), if 3>0))
 5·(4·(3·g(Y g) 2)
 ...
And so on, evaluating the structure of the algorithm recursively. Every recursively defined function can be seen as a fixed point of some other suitable function, and therefore, using Y, every recursively defined function can be expressed as a lambda expression. In particular, we can now cleanly define the subtraction, multiplication and comparison predicate of natural numbers recursively.
Computable functions and lambda calculus
A function F : N → N of natural numbers is a computable function if and only if there exists a lambda expression f such that for every pair of x, y in N, F(x) = y if and only if the expressions f x and y are equivalent. This is one of the many ways to define computability; see the ChurchTuring thesis for a discussion of other approaches and their equivalence.
Undecidability of equivalence
There is no algorithm which takes as input two lambda expressions and outputs TRUE or FALSE depending on whether or not the two expressions are equivalent. This was historically the first problem for which the unsolvability could be proven. Of course, in order to do so, the notion of algorithm has to be cleanly defined; Church used a definition via recursive functions, which is now known to be equivalent to all other reasonable definitions of the notion.
Church's proof first reduces the problem to determining whether a given lambda expression has a normal form. A normal form is an equivalent expression which cannot be reduced any further. Then he assumes that this predicate is computable, and can hence be expressed in lambda calculus. Building on earlier work by Kleene and constructing a Gödel numbering for lambda expressions, he constructs a lambda expression e which closely follows the proof of Gödel's first incompleteness theorem. If e is applied to its own Gödel number, a contradiction results.
Lambda calculus and programming languages
Most functional programming languages are equivalent to lambda calculus extended with constants and datatypes. Lisp uses a variant of lambda notation for defining functions, but only its purely functional subset is really equivalent to lambda calculus. Strictly speaking, this holds only for modern dialects of Lisp, such as Common Lisp and Scheme programming language. More archaic Lisps, such as Emacs Lisp, still use dynamic binding, and so are not based on the lambda calculus. Rather, they are based on the syntax of the lambda calculus, together with a misunderstanding of the notion of binding and substitution in the lambda calculus. In Lisp terminology, this is known as the Funarg problem.
The lambda calculus is a fundamental in theoretical computer science and in programming language theory.
See also
 SKI combinator calculus
 H.P. Barendregt's Lambda cube
 JeanYves Girard's System F
 Thierry Coquand's calculus of constructions
 Typed lambda calculus
 CurryHoward isomorphism
References
 Abelson, Harold; & Gerald Jay Sussman. Structure and Interpretation of Computer Programs. The MIT Press. ISBN 0262510871.
 Barendregt, Henk, The lambda calculus, its syntax and semantics (http://www.elsevier.com/wps/find/bookdescription.cws_home/501727/description#description), NorthHolland (1984), is the comprehensive reference on the (untyped) lambda calculus.
 Church, Alonzo, An unsolvable problem of elementary number theory, American Journal of Mathematics, 58 (1936), pp 345  363. This paper contains the proof that the equivalence of lambda expressions is in general not decidable.
 Gupta, Amit; & Ashutosh Agte, Untyped lambdacalculus, alpha, beta and eta reductions and recursion
 Henz, Martin, The Lambda Calculus. Formally correct development of the Lambda calculus.
 Kleene, Stephen, A theory of positive integers in formal logic, American Journal of Mathematics, 57 (1935), pp 153  173 and 219  244. Contains the lambda calculus definitions of several familiar functions.
 Larson, Jim, An Introduction to Lambda Calculus and Scheme (http://www.jetcafe.org/~jim/lambda.html). A gentle introduction for programmers.
Some parts of this article are based on material from FOLDOC, used with permission.
External links
 L. Allison, Some executable λcalculus examples (http://www.csse.monash.edu.au/~lloyd/tildeFP/Lambda/Examples/)
 Georg P. Loczewski, The Lambda Calculus and A++ (http://www.lambdabound.com/book/lambdacalc/lcalconl.html)
 To Dissect a Mockingbird: A Graphical Notation for the Lambda Calculus with Animated Reduction (http://uq.net.au/~zzdkeena/Lambda/index.htm)
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