One of the missing pieces of lamba calculus as we have seen so far is the ability to express recursive functions, since the ability to define and use a function is absent in lambda calculus. But any reasonable functional programming relies on recursion to do anything non-trivial. Here we would develop some theory around how we implement recursion as a feature of the language.

## Definitions

So far in our programming language we never named any function, all functions that we defined and used were anonymous. We would now add the syntax for naming/defining a function, since any recursive function needs a way to refer to itself somehow. The syntax we introduce will be like

let x = P in Q

where $x$ is a term variable $P$ and $Q$ are lambda expressions. Though this form is equivalent to defining something like this

(\x. Q) P

But we will still introduce this let form, as it is much clearer, and also it allows us to define polymorphic functions. We would talk more about it when we get to type systems.

## Self recursion

Consider the program

let fact = \n . if n == 1 then 1 else (n * (fact (n - 1))) in (fact 5)

Though we dont yet have a notion of numbers and if ... then ... else ..., consider they are available for the moment (We would make them available soon). How do we transform that into a lambda expression?

Any self-recursive function will be of the form

let P = \x ... P ...

The primary problem is we need a way to use P inside the definition of P itself. We can rewrite this expression into

let P = (\q. (\x. .... q ...))P

And also if we were some how able to build a lambda expression, lets call it Y, which could do infinite self application of a function applied to itself like,

Yf = f(Yf) ie., Yf = f(f(f(f(f .....))))

Lets call E = \q.(\x. .... q ...), then the definition of P will be like

P = EP expanded to P = E(EP) to P = E(E(EP)) ….

Now we can use Y as

let P = Y(\q. (\x. .... q ...))

that removes the recursive usage of P in the definition of P

Our factorial example will be

let fact = Y(\fn. \n. if n == 1 then 1 else (n * (fn (n - 1)))) in (fact 5)

The execution of the same would be like this

• fact 5
• (\n. (if n ==1 then 1 else n * ((Y fn) (n - 1)))) 5
• 5 * ((\n. (if n ==1 then 1 else n * ((Y fn) (n - 1)))) 4)
• 5 * 4 * ((\n. (if n ==1 then 1 else n * ((Y fn) (n - 1)))) 3)
• 5 * 4 * 3 * ((\n. (if n ==1 then 1 else n * ((Y fn) (n - 1)))) 2)
• 5 * 4 * 3 * 2 * ((\n. (if n ==1 then 1 else n * ((Y fn) (n - 1)))) 1)
• 5 * 4 * 3 * 2 * 1

## Fixed point combinator

Now the question is more about is it possible to define such a Y in lambda calculus. We know of a lambda expression $(\lambda x. xxx)(\lambda x. xxx)$. which expands to $(\lambda x. xxx) (\lambda x.xxx)(\lambda x.xxx)$ and never terminates

Lets call $Z = \lambda x.xxx$ and $E = ZZ$. Its clear that $E \rightarrow ZE \rightarrow Z(ZE) ...$ which looks similar to our $Yf$ reduction, except if we could somehow keep generating $f$ instead of $Z$ so we can use that principle to build a lambda expression

$Y = \lambda f. (\lambda x. fxx)(\lambda x. fxx)$

It is called a fixed point combinator, since it can be used to find fixed point of lambda expression. Fixed point of a function $f$ is any $x$ such that $fx = x$. More on fixed points in some other post.

Thus we can get rid of recursion and compile all of let x = Q in P to (\x. P)(Y(\x. Q)). But we wont do that, since it gets really complicated when we have mutual recursion, though this is a nicer way to do execution incase we follow substitution. Execution with environments have better and easier ways to do recursion.