# Flexibly-Phased Constraints in Haskell.

Unfortunately, Haskell’s Prelude abounds with partial functions, like head, tail, read, etc. There have been some attempts to resolve this by replacing them with safe variants that return Maybe a rather than a. But what about when we can make static guarantees about data?

Because Haskell doesn’t have real $$\prod$$-types, we can’t bring those invariants into our program specifications in the same way a dependently-typed language could. But with GHC 7.4’s new -XDataKinds extension, many data types are automatically promoted into the kind-level; this means that we can maintain a static-dynamic phase distinction at the same time as having data structures mirrored into the kind level. Combined with phantom-types, this is enough to start encoding interesting invariants into the type system. But, reusing statically verified code for dynamic values can be problematic.

Let’s look back at our length-indexed vectors to see the problem first-hand:

> {-# LANGUAGE GADTs #-}
> {-# LANGUAGE DataKinds #-}
> {-# LANGUAGE KindSignatures #-}
> {-# LANGUAGE ExistentialQuantification #-}
> {-# LANGUAGE StandaloneDeriving #-}
> {-# LANGUAGE TypeFamilies #-}
> module PhasedConstraints where
> import Control.Applicative ((<$>)) > data Nat = Z | S Nat > infixr :> > data Vect :: * -> Nat -> * where > VNil :: Vect a Z > (:>) :: a -> Vect a n -> Vect a (S n) > deriving instance Show a => Show (Vect a n) It’s trivial to write safe head and tail functions now: vhead :: Vect a (S n) -> a vhead (x :> xs) = x vtail :: Vect a (S n) -> Vect a n vtail (x :> xs) = xs But what if we don’t statically know the size of our vector? For instance, if we are converting a plain old list to a vector, or if we are parsing some file. We’d presumably box up the vector and existentially quantify its length: > data EVect :: * -> * where > EVect :: Vect a n -> EVect a > deriving instance Show a => Show (EVect a) Now we can marshall $$[\alpha]$$ into $$\forall (\alpha:Set). \exists (n:\mathbb{N}). a^n$$: > fromList :: [a] -> EVect a > fromList [] = EVect VNil > fromList (x:xs) = > case fromList xs of > EVect v -> EVect$ x :> v

But now, we’re in a bit of a pickle, since we cannot reuse our vhead and vtail. So, we can just write new ones that target Maybe:

evhead :: EVect a -> Maybe a
evhead (EVect VNil)     = Nothing
evhead (EVect (a :> b)) = Just a
evtail :: EVect a -> Maybe (EVect a)
evtail (EVect VNil)     = Nothing
evtail (EVect (a :> b)) = Just (EVect b)

But this is certainly less-than-desirable. We haven’t reused any of our code. Imagine how unfortunate this would be for a function that is more complicated than head or tail… What we actually need is a way to marshall a function with static guarantees to a function with dynamic guarantees without rewriting everything, and without case analysis.

### Types as Propositions

What if instead of restricting the type of the input of vhead, we just required a proof that it was a nonempty vector? Let’s try that. Under the Curry-Howard correspondance, types are propositions, and values are proofs. So, we can create a type that represents the proposition greater-than-zero for natural numbers:

> data NotZero :: Nat -> * where
>   NotZero :: NotZero (S n)
> deriving instance Show (NotZero n)

Basically, excluding bottom, NotZero n does not have any inhabitants of type NotZero Z. So, forall natural numbers n, any value NotZero n is a proof that n is not zero. And we can just take this as a parameter in our new version of vhead:

> vhead :: Vect a n -> NotZero n -> a
> vhead (x :> xs) NotZero = x

To use vhead, just include the proof in the parameters:

exampleVector = "hello" :> "world" :> VNil
hello = vhead exampleVector NotZero
-- typeError = vhead VNil NotZero

It’s a bit more complicated for vtail, since the length of its input must be universally quantified for it to be able to be applied to a vector with existentially quantified length; that is, behavior must not depend on the value of the index, since the index will just be a skolem type variable in the unpacked existential. So, we need to do subtraction in the right-hand side of the arrow, rather than structural recursion on the left side. To do this, we can use a type family:

> type family Prev (n :: Nat) :: Nat
> type instance Prev (S n) = n
> vtail :: Vect a n -> NotZero n -> Vect a (Prev n)
> vtail (x :> xs) NotZero = xs

Now, in order to make this work for vectors of unknown length, we need a function that can generate a NotZero proof for all vectors if applicable:

> notEmpty :: Vect a n -> Maybe (NotZero n)
> notEmpty VNil = Nothing
> notEmpty (x :> xs) = Just NotZero

And now, implementing evhead and evtail in terms of the originals is trivial, thanks to currying:

> evhead :: EVect a -> Maybe a
> evhead (EVect v) = vhead v <$> notEmpty v > evtail :: EVect a -> Maybe (EVect a) > evtail (EVect v) = EVect . vtail v <$> notEmpty v

Gloriously, all the pattern matching and so forth is factored out into the proof generator. This is especially helpful if you have a whole family of functions that you want to be available to both your statically indexed values, and those whose indices are only known dynamically.

### Taking a breath.

The reason we even bothered to do this is that there are more instances of statically verifiable computation in our programs than we like to think. For instance, any time we use a literal form for a data structure, rather than building it up inductively from IO, we are doing something that might benefit from static verification.

But in the popular literature on dependent types (and faking-dependent-types-with-GADTs), there’s quite a bit of focus on building machinery to facilitate this static verification, but not a lot of details on how to make that machinery useful for dynamic data. By factoring constraints into explicit proofs, we can allow our constraints to be checked in different program phases (compilation or execution), according to the requirements of our data (static or dynamic). In a sense, we get to have our cake, and eat it too.