Programming Languages in Haskell

Support for EECS 662 at KU


Typed Recursion

Typing Omega

Having established that the type of a lambda is of the form D:->:R, typing an increment function is quite simple:

typeof cont (lambda (x:TNum) in x + 1)
== TNum -> TNum

It follows directly that a lambda taking another lambda and applying it to a value is similarly typed:

typeof cont (lambda (f:Num -> Num) in (lambda (x:Num) in (app f x)))
== (TNum -> TNum) -> TNum

The first argument to the expression is the lambda to be applied while the second is the number applied. It is thus quite natural to have lambda expressions passed to other lambda expressions as arguments. The app in the body of the lambda applies the function argument to a value. We have seen this before. Specifically, when we looked at untyped recursion and the $\Omega$ and Y.

Remember $\Omega$, a simple recursive function that does not terminate:

(app (lambda x in (app x x))
	 (lambda x in (app x x)))

Let’s add a type placeholder, call it T->T, to the lambda and determine what a typed version of the $\Omega$ would look like:

(app (lambda (x:T->T) in (app x x))
	 (lambda (x:T->T) in (app x x)))

Will this work? Remember that to determine the type of an app f a we fine the type of f which must be a function type. Then if a has the same type as the domain of f, app f a has the same type as the range of the type of f. However, we will never find a type for app x x because the type of x must be the domain type of x. If x has type T->T, then for app x x to have type T, x must also have type T. This will never work as x would need to have both type T and T->T. The only way this will happen is when T = T->T, something that is not possible in our current type system.

$\Omega$ cannot have a type. Using the same argument, Y cannot have a type. Thus, neither can be written in our new language that includes function types.


Normalization is a term used to talk about termination. We say that a normal form is a term in a language that cannot be reduced by evaluation. In our languages terms like 1, true and lambda x in x + 1 are normal forms because eval simply returns their values without reduction. We also define these normal forms to be values representing acceptable evaluation results.

In all the languages we have written so far, the set of values and normal forms are the same. What if we removed the evaluation case for + implying that 1+1 would not evaluate further. Now terms involving + join values as normal forms. Unlike values, + terms should reduce. Normal forms that are not values are referred to as stuck values and represent errors in a language definition. It is always desirable to show that the only normal forms in a language are values.

Normalization is the property that all evaluating any term in a language always terminates resulting in a value. No non-termination and no stuck terms. Evaluation always halts with a value.

As it turns out, normalization and typing are strongly related. In our latest language, function types ensure termination. We saw this when failing to find a type for $\Omega$, but did not generalize to all function applications. Look at the types involved in this simple evaluation:

(app (lambda x in x+1) 3)

The type resulting from app is smaller than the type of lambda. This is always true because of the way app types are defined. Specifically, given f:D->R and (app f d) where d:D, the result will always be R. The base types for numbers and booleans are the smallest possible types. Like number and Boolean values, there is no way to reduce them - they are the base cases for types. Thus, app aways makes a type smaller and eventually will get to a base type. Just like subtracting 1 repeatedly from a number will eventually result in 0. Given that’s the case, if we do repeatedly execute we will always get to that smallest type which is always associated with a value.

There are ins and outs to this normalization property. It is great to know that programs termindate in many cases. But it is not great to have to know how many times any iterative construct executes when writing programs. Furthermore, there are many programs we do not want to terminate. Think operating systems or controllers as two examples. Clearly languages with function types allow this, we just need to figure out how.

Manipulating Closures

Lets revisit our original problem by going back and revisit the most famous of all recursive function, factorial:

bind fact = (lambda x in (if (isZero x) then 1 else x * (app fact x-1)))

If this is dynamically scoped it executes recursively just fine. But when static scoping and types come into play, things go downhill fast. Here is a basic definition of fact that will not execute if statically scoped:

bind fact = (lambda (x:TNum) in (if (isZero x) then 1 else x * (app fact x-1)))

Because fact is not in scope, technically there is no type for this expression either. If we can get fact in scope, maybe this will work. Is there a way to do recursion in a typed, statically scoped language? Our techniques for untyped recursion do not work. What can we do?

We need fact to be in its closure’s environment. That’s a fancy way of saying fact needs to know about itself. Let’s look at the closure resulting from evaluating the lambda defining fact:

(ClosureV "x" TNum (if (isZero x) then 1 else x * (app fact x-1)) [])

Let’s see if it works by applying the closure to 0:

(app (ClosureV "x" TNum (if (isZero x) then 1 else x * (app fact x-1)) []) 0)   []
== (if (isZero 0) then 1 else x * (app fact x-1))                               []
== 1

Remember that when we execute an app, the environment from the closure replaces the local environment. That enviornment here is empty, thus the recursive reference to fact is not in the closure’s environment. app fact 0 works only because the recursive call to fact never occurs. fact 0 is the base case and its value can be calculated directly. Still can’t fined a type for fact, but we’ll worry about that later.

Unfortunately, any value other that 0 triggers the recursive call where fact must be found in the environment. Looking at fact 1 demonstrates the problem immediately:

(app (ClosureV "x" TNum (if (isZero x) then 1 else x * (app fact x-1)) []) 1)  env
== eval (if (isZero 1) then 1 else 1 * (app fact 0))                           []
== 1 * (app fact 0)                                                            []
== error on lookup of fact in []

The actual value of env is immaterial here because we’re using static scoping. The empty environment in the closure is where we look for the definition of fact.

The easiest fix is to simply add fact to its closure’s environment. We know the definition of fact when it is defined, so we can simply add it to it’s own closure. That should do the trick because the lookup of fact will find it. Let’s give it a try:

(ClosureV "x" TNum (if x=0 then 1 else x * (app fact x-1))
    [(fact,(ClosureV "x" TNum (if ...) []))])

There are two closures here. The outer closure is what we will evaluate with app and the inner closure defines the value of fact in the outer closure’s environment. Now fact will work for 0 as it did before and should work for other values as well. Let’s give it a shot for 1:

(app (ClosureV "x" TNum (if (isZero x) then 1 else x * (app fact x-1)) []) 1) []
== (if (isZero x) then 1 else x * (app fact x-1)         [(x,1),(fact,(ClosureV "x" TNum (if ...) []))]
== (if (isZero 1) then 1 else 1 * (app fact 0))          [(x,1),(fact,(ClosureV "x" TNum (if ...) []))]
== eval 1 * (app fact 0)                                 [(x,1),(fact,(ClosureV "x" TNum (if ...) []))]
== 1 * (app (lambda x in (if (isZero x) then 1 else x * (app fact x-1)) [])  0 [(x,1),(fact,(ClosureV "x" TNum (if ...) []))]
== 1 * (if (isZero x) then 1 else x * (app fact x-1)     [(x,0)]
== 1 * 1

Good for now, but look at the last app where fact is no longer in the environment. What happened? The environment of the inner closure becomes the new environment when it is applied. This means app fact 2 will fail when the if evaluates and tries to apply fact. We have a fix for that! Let’s just add the closure again - add the closure to the environment of the closure in the closures’ environment: (Say that fast 5 times)

(ClosureV "x" (if x=0 then 1 else x * (app fact x-1))
    [(fact,(ClosureV "x" (if ...)
       [(fact,(ClosureV "x" (if ...) [(x,??)])),(x,??)])

Bingo. Now the closure in the environment for the closure knows about the closure. Now we can call fact on 0, 1, and 2. But not 3. Do you see why? The innermost closure can never have fact in its environment because it is, in effect, the base case. Any number of nested closures you choose can always be exceeded by 1. Build 10 and fact will fail for 11, build 100 and it fails for 101 and so forth. No matter how deep the nesting, eventually the recurse call to fact fails.

What can we do to solve this? In the immortal words of Dr. Seuss, no matter how many turtles we add there is always one at the bottom we can try to jump under. We cannot write $\Omega$ thus we cannot write Y. We cannot use closure magic. The only thing we are left with is adding a new construct to our language with a different execution behavior.

The Fix

In this case, the fix is adding a fixed point, concrete syntax fix t, to our statically scoped language. Instead of using the language to write a fixed point construct like Y, we will build the fixed point into the language directly and take advantage of controlling evaluation more precisely.

Fixed points are common structures in mathematics, but we only need to understand what a basic fixed point structure looks like to solve our problem.

The rule for the general recursive structure is:

Evaluating fix uses substitution to replace the called function with fix over the called function. Note that eval appears on both sides of the definition.

The let evaluates t to get a closure. the body of the closure is evaluated in e replacing i with lambda i b. What the heck?

eval env (Fix f) = let (ClosureV i b e) = (eval env f) in
                     eval e (subst i (Fix (Lambda i b)) b)

To better understand how the fix operation works, let’s evaluate factorial of 3 using our new operation. First, let’s define f, the function we will use to implement factorial:

bind fact = (lambda g in (lambda x in (if x=0 then 1 else x * (app g x-1)))) in ...

fact is not recursive. It takes a function that it will call in the recursive case and return something that looks a great deal like factorial. Let’s do a quick thought experiment to see what fact would look like if it were called on itself:

((lambda g in (lambda x in (if x=0 then 1 else x * (app g x-1)))) fact)
== (lambda x in (if x=0 then 1 else x * (app fact x-1))))

That looks exactly like what we want, but it won’t work until we use fix to perform the instantiation of f.

After evaluating f and pulling the resulting closure apart, we have the following bindings that will get used in the substitution:

i = g
b = (lambda x in (if x=0 then 1 else x * (app g x-1))
e = []

The parameter defined by the fact lambda expression is g. Thus, the argument name in the closure is g. The body of the fact lambda is what we think of as factorial with the recursive call replaced by a call to g. The environment is empty because there is nothing defined when we defined the fact lambda. Let’s start the evaluation by applying the fixed point of fact to 3:

app (fix fact) 3

Note that we are not applying fact to 3, but instead applying the fixed point of fact to 3 to build a recursive function. To evaluate the app we evaluate fact and apply the resulting value to 3. Let’s evaluate (fix fact) using the definition from above by replacing g with `(fix (lambda g in b)):

== app [g->(fix (lambda g in b))]b 3
== app (lambda x in (if x=0 then 1 else x * (app (fix (lambda g in b)) x-1) 3

Now we have something we understand. Specifically, application of a lambda to the term, 3. Substituting 3 for x results in:

== if 3 = 0 then 1 else 3 * (app (fix (lambda g in b)) 3-1)
== 3 * (app (fix (lambda g in b)) 2)

Well look at that. We got exactly what we want! We started by applying the fixed point of the lambda to a value and we just got the same thing here. Exactly the same thing with the argument decremented by 1. Let’s keep going by applying the same steps again:

== 3 * app [g->(fix (lambda g b))] b 2
== 3 * app (lambda x in (if x=0 then 1 else x * (app (fix (lambda g in b)) x-1) 2
== 3 * if 2 = 0 then 1 else 2 * (app (fix (lambda g in b)) 2-1)
== 3 * 2 * (app (fix (lambda g in b)) 1)

We are recursively executing just like we hoped we would. Now we just need to worrying about termination. Same steps again:

== 3 * 2 * (app [g->(fix (lambda g b))] b 1)
== 3 * 2 * (app (lambda x in (if x=0 then 1 else x * (app (fix (lambda g in b)) x-1) 1)
== 3 * 2 * (if 1=0 then 1 else 1 * (app (fix (lambda g in b)) 1-1))))

… and again:

== 3 * 2 * 1 * (app (fix (lambda g in b)) 0)
== 3 * 2 * 1 * (app [g->(fix (lambda g b))] b 0)
== 3 * 2 * 1 * (app (lambda x in (if x=0 then 1 else x * (app (fix (lambda g in b)) x-1) 0)
== 3 * 2 * 1 * if 0=0 then 1 else 0 * (app (fix (lambda g in b)) -1)

This time the if condition is true, so the function returns 1 rather than evaluating the fixed point again. The result is exactly what we would expect:

== 3 * 2 * 1 * 1
== 6

Finally evaluating the resulting product gives us 6 as anticipated. Our newly added fix operation takes a properly formed function and creates a recursive function. Let’s look back at the fact function given to fix:

bind fact = (lambda g in (lambda x in (if x=0 then 1 else x * (app g x-1)))) in
   (app (fix fact) 3)

This looks exactly like our original fact definition with a “hole” for the recursive call. Where fact appears in the initial definition, the function g from the outer lambda appears. This is the general form for any recursive construction we would like to create. Specifically, create the recursive function and replace the recursive instance with a variable created by an outer lambda.

Typing fix

This entire discussion started with an attempt to create a statically scoped, well-typed recursive construction. We could not find a type for $\Omega$ or Y, we couldn’t hack closures, and finally resorted to extending the core language to include a fix operator. We now have a statically scoped fix expression that will create recursive constructs for us.

One task remains. What is the type of fix? Looking at how we created factorial from fact gives is a great clue:

bind fact = (lambda g in (lambda x in (if x=0 then 1 else x * (app g x-1)))) in
   (app (fix fact) 3)

fix fact gives us the factorial fucntion. The previous definition could be rewritten as:

bind fact = (lambda g in (lambda x in (if x=0 then 1 else x * (app g x-1)))) in
   bind factorial = (fix fact) in
     (app factorial 3)

Looking at factorial this way, it should be clear the type of factorial must be TNat->TNat. Given a number, factorial will return another number. What then is the type of fact? It takes a value g and returns a function that calls g. So, the argument to fact must be a function. The result must also be a function because it is applied to a value. fact takes a function and returns a function. If we call typeof on just fact we learn:

typeof [] (lambda (ie:Nat->Nat) in
             (lambda (x:Nat) in
                (if (isZero x) then x else x + app ie x - 1)))
 == (Nat->Nat) -> (Nat->Nat)

fix takes fact and creates factorial. Instead of applying fix to an argument like app, fix skips the argument and environment going straight to the substitution. Given a function like fact, fix creates a recursive function from the body of fact using the function itself. Just like an app, the type of fix is the range of the input function:

typeof cont (Fix t) = let d:->:r = typeof cont t
                      in r