Hand-wavy explanations of my lightning talk of the same title at Mercury’s October 2022 PDX Haskell meetup. Based on “Infinite sets that admit fast exhaustive search” by Martín Escardó.

## Searchable types

We start by defining a class, $$\mathtt{Searchable} \, a$$ for types $$a$$ that admit exhaustive search. Given $$\phi : a \to \mathtt{Bool}$$ and $$a_0 : a$$, we say $$a_0$$ satisfies $$\phi$$ if $$\phi \, a_0 = \mathtt{True}$$. We say $$\phi$$ is satisfiable if at least one non-bottom member of $$a$$ satisfies $$\phi$$.

A type $$a$$ is searchable if there is a total function $$\mathtt{search} : (a \to \mathtt{Bool}) \to a$$ where $$\mathtt{search} \, \phi$$ satisfies $$\phi$$ for every satisfiable total predicate $$\phi : a \to \mathtt{Bool}$$.

$\mathtt{Searchable} \, a \iff \left( \forall_{\phi : a \to \mathtt{Bool}} \, \left( \exists_{x : a} \, \phi \, x = \mathtt{True} \right) \Rightarrow \phi \left( \mathtt{search} \, \phi \right) = \mathtt{True} \right)$

In other words, $$\mathtt{search} \, \phi$$ will find a member of $$a$$ that satisfies $$\phi$$, if such a member exists. $$\mathtt{search} \, \phi$$ still produces a member of $$a$$ even if $$\phi$$ is not satisfiable, though, so it’s the caller’s responsibility to check and see whether or not the search result satisfies $$\phi$$. Let’s automate that check!

In typical Haskell fashion, we use our API to eliminate ambiguity. If $$\mathtt{query} \, \phi$$ is $$\mathtt{Just} \, a_0$$, then $$\phi$$ is satisfiable and $$a_0$$ satisfies $$\phi$$. If $$\mathtt{query} \, \phi$$ is $$\mathtt{Nothing}$$, then no member of $$a$$ will satisfy $$\phi$$, so $$\phi$$ is not satisfiable. Let’s automate that reasoning, too!

Given this framework, Our basic problem is to discover which types are searchable and to implement $$\mathtt{search}$$ for such types.

## Finite types are searchable

Duh…

A finite type $$a$$ is searchable, because we can implement $$\mathtt{search}$$ as follows. When given a predicate $$\phi : a \to \mathtt{Bool}$$ we can iterating over the members of $$a$$, applying $$\phi$$ to each. If we find a member of $$a$$ that satisfies $$\phi$$, we stop and return it. If we exhaust $$a$$, we return the last member we checked.

This is a brute-force approach, and many searchable types will have a more-elegant, more-efficient implementation of $$\mathtt{search}$$. Here’s an implementation for $$\mathtt{Bool}$$

To verify that our instance is lawful, we consider three cases.

1. Suppose $$\phi \, \mathtt{True}$$ is $$\mathtt{True}$$. Then $$\phi \left( \mathtt{search} \, \phi \right) = \phi \left( \phi \, \mathtt{True} \right) = \phi \, \mathtt{True} = \mathtt{True}$$, as required.

2. Suppose $$\phi \, \mathtt{True} = \mathtt{False}$$ and $$\phi \, \mathtt{False} = \mathrm{True}$$. Then $$\phi \left( \mathtt{search} \, \phi \right) = \phi \left( \phi \, \mathtt{True} \right) = \phi \, \mathtt{False} = \mathtt{True}$$, as required.

3. Suppose $$\phi \, \mathtt{True} = \phi \, \mathtt{False} = \mathrm{False}$$. Then $$\mathtt{search} \, \phi$$ is not required to satisfy $$\phi$$, so there’s nothing to check: we’re done.

These cases exhaust all possibilities, thus the claim is verified.

This is effectively the solution to the Knight/Knave Riddle. Two guards stand in front of two doors; one door leads to certain death. One guard, the knave, always lies, and the other, the knight, always tells the truth. You have no way of telling which is which. You may ask a single question to one of the guards. You need to craft your question so that it’s guaranteed to reveal which door is safe less lethal, no matter which guard you ask.

## Building searchable types inductively

If each of $$a$$ and $$b$$ is searchable, then their product is searchable.

We claim that the above instance is lawful. To that end, suppose $$\phi$$ is satisfiable. We need to show that $$(a_0, b_0)$$ satisfies $$\phi$$. Since $$\phi$$ is satisfiable, $$\phi_1$$ is satisfiable, because $$\mathtt{fst} : a \times b \to a$$ maps solutions of $$\phi$$ to solutions of $$\phi_1$$. Since $$\phi_1$$ is satisfiable, $$a_0 = \mathtt{search} \, \phi_1$$ satisfies $$\phi_1$$. Since $$a_0$$ satisfies $$\phi_1$$, $$\mathtt{exists} \left( b \mapsto \phi \, (a_0,b) \right)$$ is $$\mathtt{True}$$. But notice $$\mathtt{exists} \left( b \mapsto \phi \, (a0,b) \right)$$ is identically $$\mathtt{exists} \, \phi_2$$, thus $$\mathtt{exists} \, \phi_2$$ is $$\mathtt{True}$$. Since $$\mathtt{exists} \, \phi_2$$ is $$\mathtt{True}$$, $$\phi_2$$ is satisfiable. Since $$\phi_2$$ is satisfiable, $$b_0 = \mathtt{search} \, \phi_2$$ satisfies $$\phi_2$$. That is to say $$\phi \, (a_0, b_0)$$ is $$\mathtt{True}$$, completing the proof.

So we now know that the product of any two searchable types is searchable. This generalizes in a straight-forward manner, allowing us to conclude that the product of finitely-many searchable types is searchable. What about the product of infinitely-many searchable types? In particular, let’s consider infinite-length tuples of a searchable type $$a$$. Such infinite tuples are, of course, more appropriately thought of as infinite sequences of members of $$a$$. If $$a$$ is searchable, is the type comprised of infinite sequences of members of $$a$$ searchable? Surprisingly, it is!

$$\mathtt{Sequence} \, a$$ has infinitely-many non-bottom members as long as $$a$$ has at least two non-bottom members, so we now have examples of infinite types that are still searchable. $$\mathtt{tychonoff}$$ is doing all the heavy lifting here. $$\mathtt{tychonoff}$$ is a function that takes a sequence of searchers and uses that to build a searcher of sequences. We thus define $$\mathtt{search}$$ on $$\mathtt{Sequence} \, a$$ by applying $$\mathtt{tychonoff}$$ to the constant sequence that identically returns $$\mathtt{search}$$ on $$a$$.

I’ve read Escardó’s paper four times now, and I still have no idea how $$\mathtt{tychonoff}$$ works. Maybe on my fifth reading I’ll see the light. I will attempt to convey the gist of why this is even plausible in the first place, though. I defined searchable using the word total to describe predicates. The crucial fact that this gives us is that a total predicate $$\phi : a \to \mathtt{Bool}$$ must be non-bottom for all non-bottom members of $$a$$. Now, consider the case of a predicate defined on infinite sequences. In order for a predicate $$\phi : \mathtt{Sequence} \, a \to \mathtt{Bool}$$ to be non-bottom on all non-bottom members of $$\mathtt{Sequence} \, a$$, there must necessarily be a largest natural number $$n$$ beyond which $$\phi \, x$$ does not use any terms of $$x$$. This is because the expression $$\langle\phi \, x\rangle$$ reduces to an expression $$\langle\phi_1 \, [x \, 1] \, x\rangle$$ which reduces to an expression $$\langle\phi_2 \, [x \, 1] \, [x \, 2] \, x\rangle$$ and so on.

$\langle\phi \, x\rangle \Rightarrow \langle\phi_1 \, [x \, 1] \, x\rangle \Rightarrow \langle\phi_2 \, [x \, 1] \, [x \, 2] \, x\rangle \Rightarrow \dots \Rightarrow \langle\phi_i \, [x \, 1] \, ... \, [x \, i] \, x\rangle \Rightarrow \dots$

(Where $$\langle\cdot\rangle$$ is $$\mathtt{apply}$$ and $$[\cdot]$$ is $$\mathtt{eval}$$.)

This reduction process generates an infinite sequence of expressions. Because we require $$\phi \, x$$ be non-bottom, the tail of the sequence eventually stabilizes at either the expression $$\langle\mathtt{True}\rangle$$ or the expression $$\langle\mathtt{False}\rangle$$. So given any non-bottom sequence $$x$$, we know that $$\phi \, x$$ will terminate.

This explains how $$\mathtt{tychonoff} \, (\mathtt{const} \, \mathtt{search}) \, \phi$$ terminates in the case where $$\phi$$ is satisfiable. If we assume for the moment that we are clairvoyant and we know that $$\phi$$ is satisfiable before the fact, then we further know that iterating over all non-bottom sequences will eventually yield a sequence that satisfies $$\phi$$.

What’s still not clear to me is how $$\mathtt{tychonoff} \, (\mathtt{const} \, \mathtt{search}) \, \phi$$ terminates in the case where $$\phi$$ is not satisfiable. Here’s my best guess, though. Suppose a predicate $$\phi$$ is not satisfiable. For each non-bottom sequence $$x$$, let $$N_\phi(x)$$ be the smallest natural number $$n$$ where

$\langle\phi_n \, [x \, 1] \, ... \, [x \, n] \, x\rangle \in \{\langle\mathtt{True}\rangle, \langle\mathtt{False}\rangle\} \text{.}$

(I.e., $$N_\phi(x)$$ is the point at which the sequence of reductions of the expression $$\langle \phi \, x \rangle$$ stabilizes.) We know that $$N_\phi(x) < \infty$$ for each non-bottom $$x$$. What we’d like is for $$N_\phi$$ to have a maximum, and we’d like to be able to compute that maximum. If $$N_\phi$$ had a computable maximum, we’d then know when to make $$\mathtt{tychonoff} \, (\mathtt{const} \, \mathtt{search}) \, \phi$$ stop looking. Now, this is just one approach, one very hopeful approach. It’s totally conceivable that there’s some predicate $$\phi$$ out there where $$N_\phi$$ doesn’t have a maximum. So maybe the stopping condition for $$\mathtt{tychonoff} \, (\mathtt{const} \, \mathtt{search}) \, \phi$$ is found some other way.

Tangentially, we want our search to be fast, so we memoize our sequence. A common memoization strategy in Haskell is passing a function through a concrete data structure; here we use a binary tree. This step is optional: the search algorithm takes longer but still works without it.

Here’s where we stand. Pick your favorite searchable type $$a$$. Thanks to $$\mathtt{tychonoff}$$, we know that the type comprised of infinite sequences of members of $$a$$—clearly an infinite type if $$a$$ has at least two non-bottom members—is also searchable. Exhaustively. In finite time. Mind blowing.

## Equality of functions

Functions from $$a$$ to $$b$$ may be tested for equality—extensional equality, not reference equality—deterministically and in finite time if a is searchable and b instances Eq.

Real numbers between zero and one (inclusive) may be represented by infinite sequences of bits. This is the binary encoding of the number’s decimal expansion, as decoded by approx. The Show instance for Real makes use of only eight bits of information, but this will prove to be fairly expensive. In our demo, showing Real numbers will end up taking longer than testing Real-variable functions for equality.

## Demonstration

Here’s the code I used to verify the searchability of Real, a file called demo.hs.

We define several Integer-valued functions of a Real variable to test on. We avoid returning a Real or a Double because Eq (a -> b) requires Eq b, and while Double has an Eq instance, equality for Double values is somewhat flaky. I didn’t want risk having that flakiness change the outcome of any of the tests.

Our demo will test these functions for equality, finding inputs on which they disagree if applicable. We’ll annotate our test cases with various stringy labels and also the time they take to run. In order to accurately measure this time, we force evaluation by branching on length . show of the output before grabbing the end time. Otherwise, laziness could cause misleading timings.

### Interpreted

Ran the main action of demo.hs in GHCi, which precludes any kind of optimizations and is also slower than native binaries generally. It took about 2.5 seconds and 2.6 seconds to determine that f and g disagree and that g and h disagree, respectively. Interestingly, it only took about 1.1 seconds to determine that f and h agree at every input. It took about 2.6 seconds each to find inputs where f and g disagree and where g and h disagree, unsurprisingly. To find and also print the input on which they disagree took a whopping 10.4 and 10.7 seconds, respectively. Printing the input takes longer than finding it.

Times in the output are in milliseconds.

### Compiled

Native code generated without optimizations fared much better. Finding inputs now takes less than half a second in all cases. Finding and printing takes nearly two seconds.

### Compiled with optimizations

Optimizations further bring these numbers down, to less than a quarter of a second to find inputs and less than a half of a second extra if you want to print them.