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# Solving Recursive Equations

Posted on August 14, 2015
Tags: types

I wanted to write about something related to all the stuff I’ve been reading for research lately. I decided to talk about a super cool trick in a field called domain theory. It’s a method of generating a solution to a large class of recursive equations.

In order to go through this idea we’ve got some background to cover. I wanted to make this post readable even if you haven’t read too much domain theory (you do need to know what a functor/colimit is though, nothing crazy though). We’ll start with a whirlwind tutorial of the math behind domain theory. From there we’ll transform the problem of finding a solution to an equation into something categorically tractable. Finally, I’ll walk through the construction of a solution.

I decided not to show an example of applying this technique to model a language because that would warrant its own post, hopefully I’ll write about that soon :)

## Basic Domain Theory

The basic idea with domain theory comes from a simple problem. Suppose we want to model the lambda calculus. We want a collection of mathematical objects `D` so that we can treat element of `D` as a function `D -> D` and each function `D -> D` as an element of `D`. To see why this is natural, remember that we want to turn each program `E` into `d ∈ D`. If `E = λ x. E'` then we need to turn the function `e ↦ [e/x]E'` into a term. This means `D → D` needs to be embeddable in `D`. On the other hand, we might have `E = E' E''` in which case we need to turn `E'` into a function `D → D` so that we can apply it. This means we need to be able to embed `D` into `D → D`.

After this we can turn a lambda calculus program into a specific element of `D` and reason about its properties using the ambient mathematical tools for `D`. This is semantics, understanding programs by studying their meaning in some mathematical structure. In our specific case that structure is `D` with the isomorphism `D ≅ D → D`. However, there’s an issue! We know that `D` can’t just be a set because then there cannot be such an isomorphism! In the case where `D ≅ N`, then `D → D ≅ R` and there’s a nice proof by diagonalization that such an isomorphism cannot exist.

So what can we do? We know there are only countably many programs, but we’re trying to state that there exists an isomorphism between our programs (countable) and functions on them (uncountable). Well the issue is that we don’t really mean all functions on `D`, just the ones we can model as lambda terms. For example, the function which maps all divergent programs to `1` and all terminating ones to `0` need not be considered because there’s no lambda term for it! How do we consider “computable” functions though? It’s not obvious since we define computable functions using the lambda calculus, what we’re trying to model here. Let’s set aside this question for a moment.

Another question is how do we handle this program: `(λ x. x x) (λ x. x x)`? It doesn’t have a value after all! It doesn’t behave like a normal mathematical function because applying it to something doesn’t give us back a new term, it just runs forever! To handle this we do something really clever. We stop considering just a collection of terms and instead look at terms with an ordering relation `⊑`! The idea is that ⊑ represents definedness. A program which runs to a value is more defined than a program which just loops forever. Similarly, two functions behave the same on all inputs except for `0` where one loops we could say one is more defined than the other. What we’ll do is define ⊑ abstractly and then model programs into sets with such a relation defined upon them. In order to build up this theory we need a few definitions

A partially ordered set (poset) is a set `A` and a binary relation `⊑` where

1. `a ⊑ a`
2. `a ⊑ b` and `b ⊑ c` implies `a ⊑ c`
3. `a ⊑ b` and `b ⊑ a` implies `a = b`

We often just denote the pair `<A, ⊑>` as `A` when the ordering is clear. With a poset `A`, of particular interest are chains in it. A chain is collection of elements `aᵢ` so that `aᵢ ⊑ aⱼ` if `i ≤ j`. For example, in the partial order of natural numbers and `≤`, a chain is just a run of ascending numbers. Another fundamental concept is called a least upper bound (lub). A lub of a subset `P ⊆ A` is an element of `x ∈ A` so that `y ∈ P` implies `y ⊑ x` and if this property holds for some `z` also in `A`, then `x ⊑ z`. So a least upper bound is just the smallest thing bigger than the subset. This isn’t always guaranteed to exist, for example, in our poset of natural numbers `N`, the subset `N` has no upper bounds at all! When such a lub does exist, we denote it with `⊔P`. Some partial orders have an interesting property, all chains in them have least upper bounds. We call this posets complete partial orders or cpos.

For example while `N` isn’t a cpo, `ω` (the natural numbers + an element greater than all of them) is! As a quick puzzle, can you show that all finite partial orders are in fact CPOs?

We can define a number of basic constructions on cpos. The most common is the “lifting” operation which takes a cpo `D` and returns `D⊥`, a cpo with a least element `⊥`. A cpo with such a least element is called “pointed” and I’ll write that as cppo (complete pointed partial order). Another common example, given two cppos, `D` and `E`, we can construct `D ⊗ E`. An element of this cppo is either `⊥` or `<l, r>` where `l ∈ D - {⊥}` and `r ∈ E - {⊥}`. This is called the smash product because it “smashes” the ⊥s out of the components. Similarly, there’s smash sums `D ⊕ E`.

The next question is the classic algebraic question to ask about a structure: what are the interesting functions on it? We’ll in particular be interested in functions which preserve the ⊑ relation and the taking of lub’s on chains. For this we have two more definitions:

1. A function is monotone if `x ⊑ y` implies `f(x) ⊑ f(y)`
2. A function is continuous if it is monotone and for all chains `C`, `⊔ f(P) = f(⊔ P)`.

Notably, the collection of cppos and continuous functions form a category! This is because clearly `x ↦ x` is continuous and the composition of two continuous functions is continuous. This category is called `Cpo`. It’s here that we’re going to do most of our interesting constructions.

Finally, we have to discuss one important construction on `Cpo`: `D → E`. This is the set of continuous functions from `D` to `E`. The ordering on this is pointwise, meaning that `f ⊑ g` if for all `x ∈ D`, `f(x) ⊑ g(x)`. This is a cppo where `⊥` is `x ↦ ⊥` and all the lubs are determined pointwise.

This gives us most of the mathematics we need to do the constructions we’re going to want, to demonstrate something cool here’s a fun theorem which turns out to be incredibly useful: Any continuous function `f : D → D` on a cppo `D` has a least fixed point.

To construct this least point we need to find an `x` so that `x = f(x)`. To do this, note first that `x ⊑ f(x)` by definition and by the monotonicity of `f`: `f(x) ⊑ f(y)` if `x ⊆ y`. This means that the collection of elements `fⁱ(⊥)` forms a chain with the ith element being the ith iteration of `f`! Since `D` is a cppo, this chain has an upper bound: `⊔ fⁱ(⊥)`. Moreover, `f(⊔ fⁱ(⊥)) = ⊔ f(fⁱ(⊥))` by the continuity of `f`, but `⊔ fⁱ(⊥) = ⊥ ⊔ (⊔ f(fⁱ(⊥))) = ⊔ f(fⁱ(⊥))` so this is a fixed point! The proof that it’s a least fixed point is elided because typesetting in markdown is a bit of a bother.

So there you have it, very, very basic domain theory. I can now answer the question we weren’t sure about before, the slogan is “computable functions are continuous functions”.

## Solving Recursive Equations in `Cpo`

So now we can get to the result showing domain theory incredibly useful. Remember our problem before? We wanted to find a collection `D` so that

``D ≅ D → D``

However it wasn’t clear how to do this due to size issues. In `Cpo` however, we can absolutely solve this. This huge result was due to Dana Scott. First, we make a small transformation to the problem that’s very common in these scenarios. Instead of trying to solve this equation (something we don’t have very many tools for) we’re going to instead look for the fixpoint of this functor

``F(X) = X → X``

The idea here is that we’re going to prove that all well behaved endofunctors on Cpo have fixpoints. By using this viewpoint we get all the powerful tools we normally have for reasoning about functors in category theory. However, there’s a problem: the above isn’t a functor! It has both positive and negative occurrences of `X` so it’s neither a co nor contravariant functor. To handle this we apply another clever trick. Let’s not look at endofunctors, but rather functors `Cpoᵒ × Cpo → Cpo` (I believe this should be attributed to Freyd). This is a binary functor which is covariant in the second argument and contravariant in the first. We’ll use the first argument everywhere there’s a negative occurrence of `X` and the second for every positive occurrence. Take note: we need things to be contravariant in the first argument because we’re using that first argument negatively: if we didn’t do that we wouldn’t have a functor.

Now we have

``F(X⁻, X⁺) = X⁻ → X⁺``

This is functorial. We can also always recover the original map simply by diagonalizing: `F(X) = F(X, X)`. We’ll now look for an object `D` so that `F(D, D) ≅ D`. Not quite a fixed point, but still equivalent to the equation we were looking at earlier.

Furthermore, we need one last critical property, we want `F` to be locally continuous. This means that the maps on morphisms determined by `F` should be continuous so `F(⊔ P, g) = ⊔ F(P, g)` and vice-versa (here `P` is a set of functions). Note that such morphisms have an ordering because they belong to the pointwise ordered cppo we talked about earlier.

We have one final thing to set up before this proof: what about if there’s multiple non-isomorphic solutions to `F`? We want a further coherence condition that’s going to provide us with 2 things

1. An ability to uniquely determine a solution
2. A powerful proof technique that isolates us from the particulars of the construction

What we want is called minimal invariance. Suppose we have a `D` and an `i : D ≅ F(D, D)`. This is the minimal invariant solution if and only if the least fixed point of `f(e) = i⁻ ∘ F(e, e) ∘ i` is `id`. In other words, we want it to be the case that

``d = ⊔ₓ fˣ(⊥)(d) (d ∈ D)``

I mentally picture this as saying that the isomorphism is set up so that for any particular `d` we choose, if we apply `i`, `fmap` over it, apply `i` again, repeat and repeat, eventually this process will halt and we’ll run out of things to `fmap` over. It’s a sort of a statement that each `d ∈ D` is “finite” in a very, very handwavy sense. Don’t worry if that didn’t make much sense, it’s helpful to me but it’s just my intuition. This property has some interesting effects though: it means that if we find such a `D` then `(D, D)` is going to be both the initial algebra and final coalgebra of `F`.

Without further ado, let’s prove that every locally continuous functor `F`. We start by defining the following

``````D₀ = {⊥}
Dᵢ  = F(Dᵢ₋₁, Dᵢ₋₁)``````

This gives us a chain of cppos that gradually get larger. How do we show that they’re getting larger? By defining an section from `Dᵢ` to `Dⱼ` where `j = i + 1`. A section is a function `f` which is paired with a (unique) function `f⁰` so that `f⁰f = id` and `ff⁰ ⊑ id`. In other words, `f` embeds its domain into the codomain and `f⁰` tells us how to get it out. Putting something in and taking it out is a round trip. Since the codomain may be bigger though taking something out and putting it back only approximates a round trip. Our sections are defined thusly

``````s₀ = x ↦ ⊥         r₀ = x ↦ ⊥
sᵢ  = F(rᵢ₋₁, sᵢ₋₁)   rᵢ = F(rᵢ₋₁, sᵢ₋₁)``````

It would be very instructive to work out that these definitions are actually sections and retractions. Since type-setting this subscripts is a little rough, if it’s clear from context I’ll just write `r` and `s`. Now we’ve got this increasing chain, we define an interesting object

`` D = {x ∈ Πᵢ Dᵢ | x.(i-1) = r(x.i)}``

In other words, `D` is the collection of infinitely large pairs. Each component if from one of those `Dᵢ`s above and they cohere with each other so using `s` and `r` to step up the chain takes you from one component to the next. Next we define a way to go from a single `Dᵢ` to a `D`: `upᵢ : Dᵢ → D` where

``````upᵢ(x).j =  x    if i = j
| rᵈ(x) if i - j = d > 0
| sᵈ(x) if j - i = d > 0``````

Interestingly, note that `πᵢ ∘ upᵢ = id` (easy proof) and that `upᵢ ∘ πᵢ ⊑ id` (slightly harder proof). This means that we’ve got more sections lying around: every `Dᵢ` can be fed into `D`. Consider the following diagram

``````    s      s      s
D0 ——> D1 ——> D2 ——> ...``````

I claim that `D` is the colimit to this diagram where the collection of arrows mapping into it are given with `upᵢ`. Seeing this is a colimit follows from the fact that `πᵢ ∘ upᵢ` is just `id`. Specifically, suppose we have some object `C` and a family of morphisms `cᵢ : Dᵢ → C` which commute properly with `s`. We need to find a unique morphism `h` so that `cᵢ = h ∘ upᵢ`. Define `h` as `⊔ᵢ cᵢπᵢ`. Then

``h ∘ upⱼ = (⊔j<i cᵢsʲrʲ) ⊔ cᵢ ⊔ (⊔j>i cᵢrʲsʲ) = (⊔j<i cᵢsʲrʲ) ⊔ cᵢ``

The last step follows from the fact that `rʲsʲ = id`. Furthermore, `sʲrʲ ⊑ id` so `cᵢsʲrʲ ⊑ cᵢ` so that whole massive term just evaluates to `cᵢ` as required. So we have a colimit. Notice that if we apply `F` to each `Dᵢ` in the diagram we end up with a new diagram.

``````    s      s      s
D1 ——> D2 ——> D3 ——> ...``````

`D` is still the colimit (all we’ve done is shift the diagram over by one) but by identical reasoning to `D` being a colimit, so is `F(D, D)`. This means we have a unique isomorphism `i : D ≅ F(D, D)`. The fact that `i` is the minimal invariant follows from the properties we get from the fact that `i` comes from a colimit.

With this construction we can construct our model of the lambda calculus simply by finding the minimal invariant of the locally continuous functor `F(D⁻, D⁺) = D⁻ → D⁺` (it’s worth proving it’s locally continuous). Our denotation is defined as `[e]ρ ∈ D` where `e` is a lambda term and `ρ` is a map of the free variables of `e` to other elements of `D`. This is inductively defined as

``````[λx. e]ρ = i⁻(d ↦ [e]ρ[x ↦ d])
[e e']ρ = i([e]ρ)([e']ρ)
[x]ρ = ρ(x)``````

Notice here that for the two main constructions we just use `i` and `i⁻` to fold and unfold the denotations to treat them as functions. We could go on to prove that this denotation is sound and complete but that’s something for another post.

## Wrap Up

That’s the main result I wanted to demonstrate. With this single proof we can actually model a very large class of programming languages into `Cpo`. Hopefully I’ll get around to showing how we can pull a similar trick with a relational structure on `Cpo` in order to prove full abstraction. This is nicely explained in Andrew Pitt’s “Relational Properties of Domains”.

If you’re interested in domain theory I learned from Gunter’s “Semantics of Programming Languages” book and recommend it.