Sat 6 Jan 2018

Is `State`

a `Comonad`

?

Not `Costate`

or rather, `Store`

as we tend to call it today, but actually `State s`

itself?

Let's see!

Archived Posts from this Category

Sat 6 Jan 2018

Is `State`

a `Comonad`

?

Not `Costate`

or rather, `Store`

as we tend to call it today, but actually `State s`

itself?

Let's see!

Wed 13 Jan 2016

No Comments

A common occurrence in category theory is the adjoint triple. This is a pair of adjunctions relating three functors:

F ⊣ G ⊣ H F ⊣ G, G ⊣ H

Perhaps part of the reason they are so common is that (co)limits form one:

colim ⊣ Δ ⊣ lim

where `Δ : C -> C^J`

is the diagonal functor, which takes objects in `C`

to the constant functor returning that object. A version of this shows up in Haskell (with some extensions) and dependent type theories, as:

∃ ⊣ Const ⊣ ∀ Σ ⊣ Const ⊣ Π

where, if we only care about quantifying over a single variable, existential and sigma types can be seen as a left adjoint to a diagonal functor that maps types into constant type families (either over `*`

for the first triple in Haskell, or some other type for the second in a dependently typed language), while universal and pi types can be seen as a right adjoint to the same.

It's not uncommon to see the above information in type theory discussion forums. But, there are a few cute properties and examples of adjoint triples that I haven't really seen come up in such contexts.

To begin, we can compose the two adjunctions involved, since the common functor ensures things match up. By calculating on the hom definition, we can see:

Hom(FGA, B) Hom(GFA, B) ~= ~= Hom(GA, GB) Hom(FA, HB) ~= ~= Hom(A, HGB) Hom(A, GHB)

So there are two ways to compose the adjunctions, giving two induced adjunctions:

FG ⊣ HG, GF ⊣ GH

And there is something special about these adjunctions. Note that `FG`

is the comonad for the `F ⊣ G`

adjunction, while `HG`

is the monad for the `G ⊣ H`

adjunction. Similarly, `GF`

is the `F ⊣ G`

monad, and `GH`

is the `G ⊣ H`

comonad. So each adjoint triple gives rise to two adjunctions between monads and comonads.

The second of these has another interesting property. We often want to consider the algebras of a monad, and coalgebras of a comonad. The (co)algebra operations with carrier `A`

have type:

alg : GFA -> A coalg : A -> GHA

but these types are isomorphic according to the `GF ⊣ GH`

adjunction. Thus, one might guess that `GF`

monad algebras are also `GH`

comonad coalgebras, and that in such a situation, we actually have some structure that can be characterized both ways. In fact this is true for any monad left adjoint to a comonad; [0] but all adjoint triples give rise to these.

The first adjunction actually turns out to be more familiar for the triple examples above, though. (Edit: [2]) If we consider the `Σ ⊣ Const ⊣ Π`

adjunction, where:

Σ Π : (A -> Type) -> Type Const : Type -> (A -> Type)

we get:

ΣConst : Type -> Type ΣConst B = A × B ΠConst : Type -> Type ΠConst B = A -> B

So this is the familiar adjunction:

A × - ⊣ A -> -

But, there happens to be a triple that is a bit more interesting for both cases. It refers back to categories of functors vs. bare type constructors mentioned in previous posts. So, suppose we have a category called `Con`

whose objects are (partially applied) type constructors (f, g) with kind `* -> *`

, and arrows are polymorphic functions with types like:

```
forall x. f x -> g x
```

And let us further imagine that there is a similar category, called `Func`

, except its objects are the things with `Functor`

instances. Now, there is a functor:

U : Func -> Con

that 'forgets' the functor instance requirement. This functor is in the middle of an adjoint triple:

F ⊣ U ⊣ C F, C : Con -> Func

where `F`

creates the free functor over a type constructor, and `C`

creates the cofree functor over a type constructor. These can be written using the types:

data F f a = forall e. F (e -> a) (f e) newtype C f a = C (forall r. (a -> r) -> f r)

and these types will also serve as the types involved in the composite adjunctions:

FU ⊣ CU : Func -> Func UF ⊣ UC : Con -> Con

Now, `CU`

is a monad on functors, and the Yoneda lemma tells us that it is actually the identity monad. Similarly, `FU`

is a comonad, and the co-Yoneda lemma tells us that it is the identity comonad (which makes sense, because identity is self-adjoint; and the above is why `F`

and `C`

are often named `(Co)Yoneda`

in Haskell examples).

On the other hand, `UF`

is a *monad* on type constructors (note, `U`

isn't represented in the Haskell types; `F`

and `C`

just play triple duty, and the constraints on `f`

control what's going on):

eta :: f a -> F f a eta = F id transform :: (forall x. f x -> g x) -> F f a -> F g a transform tr (F g x) = F g (tr x) mu :: F (F f) a -> F f a mu (F g (F h x)) = F (g . h) x

and `UC`

is a *comonad*:

epsilon :: C f a -> f a epsilon (C e) = e id transform' :: (forall x. f x -> g x) -> C f a -> C g a transform' tr (C e) = C (tr . e) delta :: C f a -> C (C f) a delta (C e) = C $ \h -> C $ \g -> e (g . h)

These are not the identity (co)monad, but this is the case where we have algebras and coalgebras that are equivalent. So, what are the (co)algebras? If we consider `UF`

(and unpack the definitions somewhat):

alg :: forall e. (e -> a, f e) -> f a alg (id, x) = x alg (g . h, x) = alg (g, alg (h, x))

and for `UC`

:

coalg :: f a -> forall r. (a -> r) -> f r coalg x id = x coalg x (g . h) = coalg (coalg x h) g

in other words, (co)algebra actions of these (co)monads are (mangled) `fmap`

implementations, and the commutativity requirements are exactly what is required to be a law abiding instance. So the (co)algebras are exactly the `Functors`

. [1]

There are, of course, many other examples of adjoint triples. And further, there are even adjoint quadruples, which in turn give rise to adjoint triples of (co)monads. Hopefully this has sparked some folks' interest in finding and studying more interesting examples.

[0]: Another exmaple is `A × - ⊣ A -> -`

where the `A`

in question is a monoid. (Co)monad (co)algebras of these correspond to actions of the monoid on the carrier set.

[1]: This shouldn't be too surprising, because having a category of (co)algebraic structures that is equivalent to the category of (co)algebras of the (co)monad that comes from the (co)free-forgetful adjunction is the basis for doing algebra in category theory (with (co)monads, at least). However, it is somewhat unusual for a forgetful functor to have both a left and right adjoint. In many cases, something is either algebraic or coalgebraic, and not both.

[2]: Urs Schreiber informed me of an interesting interpretation of the `ConstΣ ⊣ ConstΠ`

adjunction. If you are familiar with modal logic and the possible worlds semantics thereof, you can probably imagine that we could model it using something like `P : W -> Type`

, where `W`

is the type of possible worlds, and propositions are types. Then values of type `Σ P`

demonstrate that `P`

holds in particular worlds, while values of type `Π P`

demonstrate that it holds in all worlds. `Const`

turns these types back into world-indexed 'propositions,' so `ConstΣ`

is the possibility modality and `ConstΠ`

is the necessity modality.

Wed 22 Jul 2015

One area where I'm at odds with the prevailing winds in Haskell is lazy I/O. It's often said that lazy I/O is evil, scary and confusing, and it breaks things like referential transparency. Having a soft spot for it, and not liking most of the alternatives, I end up on the opposite side when the topic comes up, if I choose to pick the fight. I usually don't feel like I come away from such arguments having done much good at giving lazy I/O its justice. So, I thought perhaps it would be good to spell out my whole position, so that I can give the best defense I can give, and people can continue to ignore it, without costing me as much time in the future. :)

So, what's the argument that lazy I/O, or `unsafeInterleaveIO`

on which it's based, breaks referential transparency? It usually looks something like this:

swap (x, y) = (y, x) setup = do r1 < - newIORef True r2 <- newIORef True v1 <- unsafeInterleaveIO $ do writeIORef r2 False ; readIORef r1 v2 <- unsafeInterleaveIO $ do writeIORef r1 False ; readIORef r2 return (v1, v2) main = do p1 <- setup p2 <- setup print p1 print . swap $ p2

I ran this, and got:

(True, False) (True, False)

So this is supposed to demonstrate that the pure values depend on evaluation order, and we have broken a desirable property of Haskell.

First a digression. Personally I distinguish the terms, "referential transparency," and, "purity," and use them to identify two desirable properties of Haskell. The first I use for the property that allows you to factor your program by introducing (or eliminating) named subexpressions. So, instead of:

f e e

we are free to write:

let x = e in f x x

or some variation. I have no argument for this meaning, other than it's what I thought it meant when I first heard the term used with respect to Haskell, it's a useful property, and it's the best name I can think of for the property. I also (of course) think it's better than some of the other explanations you'll find for what people mean when they say Haskell has referential transparency, since it doesn't mention functions or "values". It's just about equivalence of expressions.

Anyhow, for me, the above example is in no danger of violating referential transparency. There is no factoring operation that will change the meaning of the program. I can even factor out `setup`

(or inline it, since it's already named):

main = let m = setup in do p1 < - m p2 <- m print p1 print . swap $ p2

This is the way in which `IO`

preserves referential transparency, unlike side effects, in my view (note: the embedded language represented by `IO`

does not have this property, since otherwise `p1`

could be used in lieu of `p2`

; this is why you shouldn't spend much time writing `IO`

stuff, because it's a bad language embedded in a good one).

The other property, "purity," I pull from Amr Sabry's paper, What is a Purely Functional Language? There he argues that a functional language should be considered "pure" if it is an extension of the lambda calculus in which there are no contexts which observe differences in evaluation order. Effectively, evaluation order must only determine whether or not you get an answer, not change the answer you get.

This is slightly different from my definition of referential transparency earlier, but it's also a useful property to have. Referential transparency tells us that we can freely refactor, and purity tells us that we can change the order things are evaluated, both without changing the meaning of our programs.

Now, it would seem that the original interleaving example violates purity. Depending on the order that the values are evaluated, opponents of lazy I/O say, the values change. However, this argument doesn't impress me, because I think the proper way to think about `unsafeInterleaveIO`

is as concurrency, and in that case, it isn't very strange that the results of running it would be non-deterministic. And in that case, there's not much you can do to prove that the evaluation order is affecting results, and that you aren't simply very unlucky and always observing results that happen to correspond to evaluation order.

In fact, there's something I didn't tell you. I didn't use the `unsafeInterleaveIO`

from base. I wrote my own. It looks like this:

unsafeInterleaveIO :: IO a -> IO a unsafeInterleaveIO action = do iv < - new forkIO $ randomRIO (1,5) >>= threadDelay . (*1000) >> action >>= write iv return . read $ iv

`iv`

is an `IVar`

(I used ivar-simple). The pertinent operations on them are:

new :: IO (IVar a) write :: IVar a -> a -> IO () read :: IVar a -> a

`new`

creates an empty `IVar`

, and we can `write`

to one only once; trying to write a second time will throw an exception. But this is no problem for me, because I obviously only attempt to write once. `read`

will block until its argument is actually is set, and since that can only happen once, it is considered safe for `read`

to not require `IO`

. [1]

Using this and `forkIO`

, one can easily write something like `unsafeInterleaveIO`

, which accepts an `IO a`

argument and yields an `IO a`

whose result is guaranteed to be the result of running the argument at some time in the future. The only difference is that the real `unsafeInterleaveIO`

schedules things just in time, whereas mine schedules them in a relatively random order (I'll admit I had to try a few times before I got the 'expected' lazy IO answer).

But, we could even take this to be the specification of interleaving. It runs `IO`

actions concurrently, and you will be fine as long as you aren't attempting to depend on the exact scheduling order (or whether things get scheduled at all in some cases).

In fact, thinking of lazy I/O as concurrency turns most spooky examples into threading problems that I would expect most people to consider rather basic. For instance:

- Don't pass a handle to another thread and close it in the original.
- Don't fork more file-reading threads than you have file descriptors.
- Don't fork threads to handle files if you're concerned about the files being closed deterministically.
- Don't read from the same handle in multiple threads (unless you don't care about each thread seeing a random subsequence of the stream).

And of course, the original example in this article is just non-determinism introduced by concurrency, but not of a sort that requires fundamentally different explanation than fork. The main pitfall, in my biased opinion, is that the scheduling for interleaving is explained in a way that encourages people to try to guess exactly what it will do. But the presumption of purity (and the reordering GHC actually does based on it) actually means that you cannot assume that much more about the scheduling than you can about my scheduler, at least in general.

This isn't to suggest that lazy I/O is appropriate for every situation. Sometimes the above advice means that it is not appropriate to use concurrency. However, in my opinion, people are over eager to ban lazy I/O even for simple uses where it is the nicest solution, and justify it based on the 'evil' and 'confusing' ascriptions. But, personally, I don't think this is justified, unless one does the same for pretty much all concurrency.

I suppose the only (leading) question left to ask is which should be declared unsafe, fork or ivars, since together they allow you to construct a(n even less deterministic) `unsafeInterleaveIO`

?

[1] Note that there are other implementations of `IVar`

. I'd expect the most popular to be in monad-par by Simon Marlow. That allows one to construct an operation like `read`

, but it is actually *less* deterministic in my construction, because it seems that it will not block unless perhaps you write and read within a single 'transaction,' so to speak.

In fact, this actually breaks referential transparency in conjunction with `forkIO`

:

deref = runPar . get randomDelay = randomRIO (1,10) >>= threadDelay . (1000*) myHandle m = m `catch` \(_ :: SomeExpression) -> putStrLn "Bombed" mySpawn :: IO a -> IO (IVar a) mySpawn action = do iv < - runParIO new forkIO $ randomDelay >> action >>= runParIO . put_ iv return iv main = do iv < - mySpawn (return True) myHandle . print $ deref iv randomDelay myHandle . print $ deref iv

Sometimes this will print "Bombed" twice, and sometimes it will print "Bombed" followed by "True". The latter will never happen if we factor out the `deref iv`

however. The blocking behavior is essential to `deref`

maintaining referential transparency, and it seems like monad-par only blocks within a single `runPar`

, not across multiples. Using ivar-simple in this example always results in "True" being printed twice.

It is also actually possible for `unsafeInterleaveIO`

to break referential transparency if it is implemented incorrectly (or if the optimizer mucks with the internals in some bad way). But I haven't seen an example that couldn't be considered a bug in the implementation rather than some fundamental misbehavior. And my reference implementation here (with a suboptimal scheduler) suggests that there is no break that isn't just a bug.

Mon 25 May 2015

[4] Comments

In the last couple posts I've used some 'free' constructions, and not remarked too much on how they arise. In this post, I'd like to explore them more. This is going to be something of a departure from the previous posts, though, since I'm not going to worry about thinking precisely about bottom/domains. This is more an exercise in applying some category theory to Haskell, "fast and loose".

(Advance note: for some continuous code to look at see this file.)

First, it'll help to talk about how some categories can work in Haskell. For any kind `k`

made of `*`

and `(->)`

, [0] we can define a category of type constructors. Objects of the category will be first-class [1] types of that kind, and arrows will be defined by the following type family:

newtype Transformer f g = Transform { ($$) :: forall i. f i ~> g i } type family (~>) :: k -> k -> * where (~>) = (->) (~>) = Transformer type a < -> b = (a -> b, b -> a) type a < ~> b = (a ~> b, b ~> a)

So, for a base case, * has monomorphic functions as arrows, and categories for higher kinds have polymorphic functions that saturate the constructor:

Int ~> Char = Int -> Char Maybe ~> [] = forall a. Maybe a -> [a] Either ~> (,) = forall a b. Either a b -> (a, b) StateT ~> ReaderT = forall s m a. StateT s m a -> ReaderT s m a

We can of course define identity and composition for these, and it will be handy to do so:

class Morph (p :: k -> k -> *) where id :: p a a (.) :: p b c -> p a b -> p a c instance Morph (->) where id x = x (g . f) x = g (f x) instance Morph ((~>) :: k -> k -> *) => Morph (Transformer :: (i -> k) -> (i -> k) -> *) where id = Transform id Transform f . Transform g = Transform $ f . g

These categories can be looked upon as the most basic substrates in Haskell. For instance, every type of kind `* -> *`

is an object of the relevant category, even if it's a GADT or has other structure that prevents it from being nicely functorial.

The category for * is of course just the normal category of types and functions we usually call Hask, and it is fairly analogous to the category of sets. One common activity in category theory is to study categories of sets equipped with extra structure, and it turns out we can do this in Haskell, as well. And it even makes some sense to study categories of structures over any of these type categories.

When we equip our types with structure, we often use type classes, so that's how I'll do things here. Classes have a special status socially in that we expect people to only define instances that adhere to certain equational rules. This will take the place of equations that we are not able to state in the Haskell type system, because it doesn't have dependent types. So using classes will allow us to define more structures that we normally would, if only by convention.

So, if we have a kind `k`

, then a corresponding structure will be `σ :: k -> Constraint`

. We can then define the category `(k,σ)`

as having objects `t :: k`

such that there is an instance `σ t`

. Arrows are then taken to be `f :: t ~> u`

such that `f`

"respects" the operations of `σ`

.

As a simple example, we have:

k = * σ = Monoid :: * -> Constraint Sum Integer, Product Integer, [Integer] :: (*, Monoid) f :: (Monoid m, Monoid n) => m -> n if f mempty = mempty f (m <> n) = f m <> f n

This is just the category of monoids in Haskell.

As a side note, we will sometimes be wanting to quantify over these "categories of structures". There isn't really a good way to package together a kind and a structure such that they work as a unit, but we can just add a constraint to the quantification. So, to quantify over all `Monoid`

s, we'll use '`forall m. Monoid m => ...`

'.

Now, once we have these categories of structures, there is an obvious forgetful functor back into the unadorned category. We can then look for free and cofree functors as adjoints to this. More symbolically:

Forget σ :: (k,σ) -> k Free σ :: k -> (k,σ) Cofree σ :: k -> (k,σ) Free σ ⊣ Forget σ ⊣ Cofree σ

However, what would be nicer (for some purposes) than having to look for these is being able to construct them all systematically, without having to think much about the structure `σ`

.

Category theory gives a hint at this, too, in the form of Kan extensions. In category terms they look like:

p : C -> C' f : C -> D Ran p f : C' -> D Lan p f : C' -> D Ran p f c' = end (c : C). Hom_C'(c', p c) ⇒ f c Lan p f c' = coend (c : c). Hom_C'(p c, c') ⊗ f c

where `⇒`

is a "power" and `⊗`

is a copower, which are like being able to take exponentials and products by sets (or whatever the objects of the hom category are), instead of other objects within the category. Ends and coends are like universal and existential quantifiers (as are limits and colimits, but ends and coends involve mixed-variance).

Some handy theorems relate Kan extensions and adjoint functors:

if L ⊣ R then L = Ran R Id and R = Lan L Id if Ran R Id exists and is absolute then Ran R Id ⊣ R if Lan L Id exists and is absolute then L ⊣ Lan L Id Kan P F is absolute iff forall G. (G . Kan P F) ~= Kan P (G . F)

It turns out we can write down Kan extensions fairly generally in Haskell. Our restricted case is:

p = Forget σ :: (k,σ) -> k f = Id :: (k,σ) -> (k,σ) Free σ = Ran (Forget σ) Id :: k -> (k,σ) Cofree σ = Lan (Forget σ) Id :: k -> (k,σ) g :: (k,σ) -> j g . Free σ = Ran (Forget σ) g g . Cofree σ = Lan (Forget σ) g

As long as the final category is like one of our type constructor categories, ends are universal quantifiers, powers are function types, coends are existential quantifiers and copowers are product spaces. This only breaks down for our purposes when `g`

is contravariant, in which case they are flipped. For higher kinds, these constructions occur point-wise. So, we can break things down into four general cases, each with cases for each arity:

newtype Ran0 σ p (f :: k -> *) a = Ran0 { ran0 :: forall r. σ r => (a ~> p r) -> f r } newtype Ran1 σ p (f :: k -> j -> *) a b = Ran1 { ran1 :: forall r. σ r => (a ~> p r) -> f r b } -- ... data RanOp0 σ p (f :: k -> *) a = forall e. σ e => RanOp0 (a ~> p e) (f e) -- ... data Lan0 σ p (f :: k -> *) a = forall e. σ e => Lan0 (p e ~> a) (f e) data Lan1 σ p (f :: k -> j -> *) a b = forall e. σ e => Lan1 (p e ~> a) (f e b) -- ... data LanOp0 σ p (f :: k -> *) a = LanOp0 { lan0 :: forall r. σ r => (p r -> a) -> f r } -- ...

The more specific proposed (co)free definitions are:

type family Free :: (k -> Constraint) -> k -> k type family Cofree :: (k -> Constraint) -> k -> k newtype Free0 σ a = Free0 { gratis0 :: forall r. σ r => (a ~> r) -> r } type instance Free = Free0 newtype Free1 σ f a = Free1 { gratis1 :: forall g. σ g => (f ~> g) -> g a } type instance Free = Free1 -- ... data Cofree0 σ a = forall e. σ e => Cofree0 (e ~> a) e type instance Cofree = Cofree0 data Cofree1 σ f a = forall g. σ g => Cofree1 (g ~> f) (g a) type instance Cofree = Cofree1 -- ...

We can define some handly classes and instances for working with these types, several of which generalize existing Haskell concepts:

class Covariant (f :: i -> j) where comap :: (a ~> b) -> (f a ~> f b) class Contravariant f where contramap :: (b ~> a) -> (f a ~> f b) class Covariant m => Monad (m :: i -> i) where pure :: a ~> m a join :: m (m a) ~> m a class Covariant w => Comonad (w :: i -> i) where extract :: w a ~> a split :: w a ~> w (w a) class Couniversal σ f | f -> σ where couniversal :: σ r => (a ~> r) -> (f a ~> r) class Universal σ f | f -> σ where universal :: σ e => (e ~> a) -> (e ~> f a) instance Covariant (Free0 σ) where comap f (Free0 e) = Free0 (e . (.f)) instance Monad (Free0 σ) where pure x = Free0 $ \k -> k x join (Free0 e) = Free0 $ \k -> e $ \(Free0 e) -> e k instance Couniversal σ (Free0 σ) where couniversal h (Free0 e) = e h -- ...

The only unfamiliar classes here should be `(Co)Universal`

. They are for witnessing the adjunctions that make `Free σ`

the initial `σ`

and `Cofree σ`

the final `σ`

in the relevant way. Only one direction is given, since the opposite is very easy to construct with the (co)monad structure.

`Free σ`

is a monad and couniversal, `Cofree σ`

is a comonad and universal.

We can now try to convince ourselves that `Free σ`

and `Cofree σ`

are absolute Here are some examples:

free0Absolute0 :: forall g σ a. (Covariant g, σ (Free σ a)) => g (Free0 σ a) < -> Ran σ Forget g a free0Absolute0 = (l, r) where l :: g (Free σ a) -> Ran σ Forget g a l g = Ran0 $ \k -> comap (couniversal $ remember0 . k) g r :: Ran σ Forget g a -> g (Free σ a) r (Ran0 e) = e $ Forget0 . pure free0Absolute1 :: forall (g :: * -> * -> *) σ a x. (Covariant g, σ (Free σ a)) => g (Free0 σ a) x < -> Ran σ Forget g a x free0Absolute1 = (l, r) where l :: g (Free σ a) x -> Ran σ Forget g a x l g = Ran1 $ \k -> comap (couniversal $ remember0 . k) $$ g r :: Ran σ Forget g a x -> g (Free σ a) x r (Ran1 e) = e $ Forget0 . pure free0Absolute0Op :: forall g σ a. (Contravariant g, σ (Free σ a)) => g (Free0 σ a) < -> RanOp σ Forget g a free0Absolute0Op = (l, r) where l :: g (Free σ a) -> RanOp σ Forget g a l = RanOp0 $ Forget0 . pure r :: RanOp σ Forget g a -> g (Free σ a) r (RanOp0 h g) = contramap (couniversal $ remember0 . h) g -- ...

As can be seen, the definitions share a lot of structure. I'm quite confident that with the right building blocks these could be defined once for each of the four types of Kan extensions, with types like:

freeAbsolute :: forall g σ a. (Covariant g, σ (Free σ a)) => g (Free σ a) < ~> Ran σ Forget g a cofreeAbsolute :: forall g σ a. (Covariant g, σ (Cofree σ a)) => g (Cofree σ a) < ~> Lan σ Forget g a freeAbsoluteOp :: forall g σ a. (Contravariant g, σ (Free σ a)) => g (Free σ a) < ~> RanOp σ Forget g a cofreeAbsoluteOp :: forall g σ a. (Contravariant g, σ (Cofree σ a)) => g (Cofree σ a) < ~> LanOp σ Forget g a

However, it seems quite difficult to structure things in a way such that GHC will accept the definitions. I've successfully written `freeAbsolute`

using some axioms, but turning those axioms into class definitions and the like seems impossible.

Anyhow, the punchline is that we can prove absoluteness using only the premise that there is a valid `σ`

instance for `Free σ`

and `Cofree σ`

. This tends to be quite easy; we just borrow the structure of the type we are quantifying over. This means that in all these cases, we are justified in saying that `Free σ ⊣ Forget σ ⊣ Cofree σ`

, and we have a very generic presentations of (co)free structures in Haskell. So let's look at some.

We've already seen `Free Monoid`

, and last time we talked about `Free Applicative`

, and its relation to traversals. But, `Applicative`

is to traversal as `Functor`

is to lens, so it may be interesting to consider constructions on that. Both `Free Functor`

and `Cofree Functor`

make `Functor`

s:

instance Functor (Free1 Functor f) where fmap f (Free1 e) = Free1 $ fmap f . e instance Functor (Cofree1 Functor f) where fmap f (Cofree1 h e) = Cofree1 h (fmap f e)

And of course, they are (co)monads, covariant functors and (co)universal among `Functor`

s. But, it happens that I know some other types with these properties:

data CoYo f a = forall e. CoYo (e -> a) (f e) instance Covariant CoYo where comap f = Transform $ \(CoYo h e) -> CoYo h (f $$ e) instance Monad CoYo where pure = Transform $ CoYo id join = Transform $ \(CoYo h (CoYo h' e)) -> CoYo (h . h') e instance Functor (CoYo f) where fmap f (CoYo h e) = CoYo (f . h) e instance Couniversal Functor CoYo where couniversal tr = Transform $ \(CoYo h e) -> fmap h (tr $$ e) newtype Yo f a = Yo { oy :: forall r. (a -> r) -> f r } instance Covariant Yo where comap f = Transform $ \(Yo e) -> Yo $ (f $$) . e instance Comonad Yo where extract = Transform $ \(Yo e) -> e id split = Transform $ \(Yo e) -> Yo $ \k -> Yo $ \k' -> e $ k' . k instance Functor (Yo f) where fmap f (Yo e) = Yo $ \k -> e (k . f) instance Universal Functor Yo where universal tr = Transform $ \e -> Yo $ \k -> tr $$ fmap k e

These are the types involved in the (co-)Yoneda lemma. `CoYo`

is a monad, couniversal among functors, and `CoYo f`

is a `Functor`

. `Yo`

is a comonad, universal among functors, and is always a `Functor`

. So, are these equivalent types?

coyoIso :: CoYo < ~> Free Functor coyoIso = (Transform $ couniversal pure, Transform $ couniversal pure) yoIso :: Yo < ~> Cofree Functor yoIso = (Transform $ universal extract, Transform $ universal extract)

Indeed they are. And similar identities hold for the contravariant versions of these constructions.

I don't have much of a use for this last example. I suppose to be perfectly precise, I should point out that these uses of `(Co)Yo`

are not actually part of the (co-)Yoneda lemma. They are two different constructions. The (co-)Yoneda lemma can be given in terms of Kan extensions as:

yoneda :: Ran Id f < ~> f coyoneda :: Lan Id f < ~> f

But, the use of `(Co)Yo`

to make `Functor`

s out of things that aren't necessarily is properly thought of in other terms. In short, we have some kind of category of Haskell types with only identity arrows---it is discrete. Then any type constructor, even non-functorial ones, is certainly a functor from said category (call it Haskrete) into the normal one (Hask). And there is an inclusion functor from Haskrete into Hask:

F Haskrete -----> Hask | /| | / | / Incl | / | / Ran/Lan Incl F | / | / v / Hask

So, `(Co)Free Functor`

can also be thought of in terms of these Kan extensions involving the discrete category.

To see more fleshed out, loadable versions of the code in this post, see this file. I may also try a similar Agda development at a later date, as it may admit the more general absoluteness constructions easier.

[0]: The reason for restricting ourselves to kinds involving only `*`

and `(->)`

is that they work much more simply than data kinds. Haskell values can't depend on type-level entities without using type classes. For *, this is natural, but for something like `Bool -> *`

, it is more natural for transformations to be able to inspect the booleans, and so should be something more like `forall b. InspectBool b => f b -> g b`

.

[1]: First-class types are what you get by removing type families and synonyms from consideration. The reason for doing so is that these can't be used properly as parameters and the like, except in cases where they reduce to some other type that is first-class. For example, if we define:

```
type I a = a
```

even though GHC will report `I :: * -> *`

, it is not legal to write `Transform I I`

.

Wed 29 Apr 2015

No Comments

Last time I looked at free monoids, and noticed that in Haskell lists don't really cut it. This is a consequence of laziness and general recursion. To model a language with those properties, one needs to use domains and monotone, continuous maps, rather than sets and total functions (a call-by-value language with general recursion would use domains and strict maps instead).

This time I'd like to talk about some other examples of this, and point out how doing so can (perhaps) resolve some disagreements that people have about the specific cases.

The first example is not one that I came up with: induction. It's sometimes said that Haskell does not have inductive types at all, or that we cannot reason about functions on its data types by induction. However, I think this is (techincally) inaccurate. What's true is that we cannot simply pretend that that our types are sets and use the induction principles for sets to reason about Haskell programs. Instead, one has to figure out what inductive domains would be, and what their proof principles are.

Fortunately, there are some papers about doing this. The most recent (that I'm aware of) is Generic Fibrational Induction. I won't get too into the details, but it shows how one can talk about induction in a general setting, where one has a category that roughly corresponds to the type theory/programming language, and a second category of proofs that is 'indexed' by the first category's objects. Importantly, it is not required that the second category is somehow 'part of' the type theory being reasoned about, as is often the case with dependent types, although that is also a special case of their construction.

One of the results of the paper is that this framework can be used to talk about induction principles for types that don't make sense as sets. Specifically:

newtype Hyp = Hyp ((Hyp -> Int) -> Int)

the type of "hyperfunctions". Instead of interpreting this type as a set, where it would effectively require a set that is isomorphic to the power set of its power set, they interpret it in the category of domains and strict functions mentioned earlier. They then construct the proof category in a similar way as one would for sets, except instead of talking about predicates as sub*sets*, we talk about sub-*domains* instead. Once this is done, their framework gives a notion of induction for this type.

This example is suitable for ML (and suchlike), due to the strict functions, and sort of breaks the idea that we can really get away with only thinking about sets, even there. Sets are good enough for some simple examples (like flat domains where we don't care about ⊥), but in general we have to generalize induction itself to apply to all types in the 'good' language.

While I haven't worked out how the generic induction would work out for Haskell, I have little doubt that it would, because ML actually contains all of Haskell's data types (and vice versa). So the fact that the framework gives meaning to induction for ML implies that it does so for Haskell. If one wants to know what induction for Haskell's 'lazy naturals' looks like, they can study the ML analogue of:

data LNat = Zero | Succ (() -> LNat)

because function spaces lift their codomain, and make things 'lazy'.

----

The other example I'd like to talk about hearkens back to the previous article. I explained how `foldMap`

is the proper fundamental method of the `Foldable`

class, because it can be massaged to look like:

foldMap :: Foldable f => f a -> FreeMonoid a

and lists are not the free monoid, because they do not work properly for various infinite cases.

I also mentioned that `foldMap`

looks a lot like `traverse`

:

foldMap :: (Foldable t , Monoid m) => (a -> m) -> t a -> m traverse :: (Traversable t, Applicative f) => (a -> f b) -> t a -> f (t b)

And of course, we have `Monoid m => Applicative (Const m)`

, and the functions are expected to agree in this way when applicable.

Now, people like to get in arguments about whether traversals are allowed to be infinite. I know Ed Kmett likes to argue that they can be, because he has lots of examples. But, not everyone agrees, and especially people who have papers proving things about traversals tend to side with the finite-only side. I've heard this includes one of the inventors of `Traversable`

, Conor McBride.

In my opinion, the above disagreement is just another example of a situation where we have a generic notion instantiated in two different ways, and intuition about one does not quite transfer to the other. If you are working in a language like Agda or Coq (for proving), you will be thinking about traversals in the context of sets and total functions. And there, traversals are finite. But in Haskell, there are infinitary cases to consider, and they should work out all right when thinking about domains instead of sets. But I should probably put forward some argument for this position (and even if I don't need to, it leads somewhere else interesting).

One example that people like to give about finitary traversals is that they can be done via lists. Given a finite traversal, we can traverse to get the elements (using `Const [a]`

), traverse the list, then put them back where we got them by traversing again (using `State [a]`

). Usually when you see this, though, there's some subtle cheating in relying on the list to be exactly the right length for the second traversal. It will be, because we got it from a traversal of the same structure, but I would expect that proving the function is actually total to be a lot of work. Thus, I'll use this as an excuse to do my own cheating later.

Now, the above uses lists, but why are we using lists when we're in Haskell? We know they're deficient in certain ways. It turns out that we can give a lot of the same relevant structure to the better free monoid type:

newtype FM a = FM (forall m. Monoid m => (a -> m) -> m) deriving (Functor) instance Applicative FM where pure x = FM ($ x) FM ef < *> FM ex = FM $ \k -> ef $ \f -> ex $ \x -> k (f x) instance Monoid (FM a) where mempty = FM $ \_ -> mempty mappend (FM l) (FM r) = FM $ \k -> l k <> r k instance Foldable FM where foldMap f (FM e) = e f newtype Ap f b = Ap { unAp :: f b } instance (Applicative f, Monoid b) => Monoid (Ap f b) where mempty = Ap $ pure mempty mappend (Ap l) (Ap r) = Ap $ (<>) < $> l < *> r instance Traversable FM where traverse f (FM e) = unAp . e $ Ap . fmap pure . f

So, free monoids are `Monoids`

(of course), `Foldable`

, and even `Traversable`

. At least, we can define something with the right type that wouldn't bother anyone if it were written in a total language with the right features, but in Haskell it happens to allow various infinite things that people don't like.

Now it's time to cheat. First, let's define a function that can take any `Traversable`

to our free monoid:

toFreeMonoid :: Traversable t => t a -> FM a toFreeMonoid f = FM $ \k -> getConst $ traverse (Const . k) f

Now let's define a `Monoid`

that's not a monoid:

data Cheat a = Empty | Single a | Append (Cheat a) (Cheat a) instance Monoid (Cheat a) where mempty = Empty mappend = Append

You may recognize this as the data version of the free monoid from the previous article, where we get the real free monoid by taking a quotient. using this, we can define an `Applicative`

that's not valid:

newtype Cheating b a = Cheating { prosper :: Cheat b -> a } deriving (Functor) instance Applicative (Cheating b) where pure x = Cheating $ \_ -> x Cheating f < *> Cheating x = Cheating $ \c -> case c of Append l r -> f l (x r)

Given these building blocks, we can define a function to relabel a traversable using a free monoid:

relabel :: Traversable t => t a -> FM b -> t b relabel t (FM m) = propser (traverse (const hope) t) (m Single) where hope = Cheating $ \c -> case c of Single x -> x

And we can implement any traversal by taking a trip through the free monoid:

slowTraverse :: (Applicative f, Traversable t) => (a -> f b) -> t a -> f (t b) slowTraverse f t = fmap (relabel t) . traverse f . toFreeMonoid $ t

And since we got our free monoid via traversing, all the partiality I hid in the above won't blow up in practice, rather like the case with lists and finite traversals.

Arguably, this is worse cheating. It relies on the exact association structure to work out, rather than just number of elements. The reason is that for infinitary cases, you cannot flatten things out, and there's really no way to detect when you have something infinitary. The finitary traversals have the luxury of being able to reassociate everything to a canonical form, while the infinite cases force us to not do any reassociating at all. So this might be somewhat unsatisfying.

But, what if we didn't have to cheat at all? We can get the free monoid by tweaking `foldMap`

, and it looks like `traverse`

, so what happens if we do the same manipulation to the latter?

It turns out that lens has a type for this purpose, a slight specialization of which is:

newtype Bazaar a b t = Bazaar { runBazaar :: forall f. Applicative f => (a -> f b) -> f t }

Using this type, we can reorder `traverse`

to get:

howBizarre :: Traversable t => t a -> Bazaar a b (t b) howBizarre t = Bazaar $ \k -> traverse k t

But now, what do we do with this? And what even is it? [1]

If we continue drawing on intuition from `Foldable`

, we know that `foldMap`

is related to the free monoid. `Traversable`

has more indexing, and instead of `Monoid`

uses `Applicative`

. But the latter are actually related to the former; `Applicative`

s are monoidal (closed) functors. And it turns out, `Bazaar`

has to do with free `Applicative`

s.

If we want to construct free `Applicative`

s, we can use our universal property encoding trick:

newtype Free p f a = Free { gratis :: forall g. p g => (forall x. f x -> g x) -> g a }

This is a higher-order version of the free `p`

, where we parameterize over the constraint we want to use to represent structures. So `Free Applicative f`

is the free `Applicative`

over a type constructor `f`

. I'll leave the instances as an exercise.

Since free monoid is a monad, we'd expect `Free p`

to be a monad, too. In this case, it is a McBride style indexed monad, as seen in The Kleisli Arrows of Outrageous Fortune.

type f ~> g = forall x. f x -> g x embed :: f ~> Free p f embed fx = Free $ \k -> k fx translate :: (f ~> g) -> Free p f ~> Free p g translate tr (Free e) = Free $ \k -> e (k . tr) collapse :: Free p (Free p f) ~> Free p f collapse (Free e) = Free $ \k -> e $ \(Free e') -> e' k

That paper explains how these are related to Atkey style indexed monads:

data At key i j where At :: key -> At key i i type Atkey m i j a = m (At a j) i ireturn :: IMonad m => a -> Atkey m i i a ireturn = ... ibind :: IMonad m => Atkey m i j a -> (a -> Atkey m j k b) -> Atkey m i k b ibind = ...

It turns out, `Bazaar`

is exactly the Atkey indexed monad derived from the `Free Applicative`

indexed monad (with some arguments shuffled) [2]:

hence :: Bazaar a b t -> Atkey (Free Applicative) t b a hence bz = Free $ \tr -> runBazaar bz $ tr . At forth :: Atkey (Free Applicative) t b a -> Bazaar a b t forth fa = Bazaar $ \g -> gratis fa $ \(At a) -> g a imap :: (a -> b) -> Bazaar a i j -> Bazaar b i j imap f (Bazaar e) = Bazaar $ \k -> e (k . f) ipure :: a -> Bazaar a i i ipure x = Bazaar ($ x) (>>>=) :: Bazaar a j i -> (a -> Bazaar b k j) -> Bazaar b k i Bazaar e >>>= f = Bazaar $ \k -> e $ \x -> runBazaar (f x) k (>==>) :: (s -> Bazaar i o t) -> (i -> Bazaar a b o) -> s -> Bazaar a b t (f >==> g) x = f x >>>= g

As an aside, `Bazaar`

is also an (Atkey) indexed comonad, and the one that characterizes traversals, similar to how indexed store characterizes lenses. A `Lens s t a b`

is equivalent to a coalgebra `s -> Store a b t`

. A traversal is a similar `Bazaar`

coalgebra:

s -> Bazaar a b t ~ s -> forall f. Applicative f => (a -> f b) -> f t ~ forall f. Applicative f => (a -> f b) -> s -> f t

It so happens that Kleisli composition of the Atkey indexed monad above `(>==>)`

is traversal composition.

Anyhow, `Bazaar`

also inherits `Applicative`

structure from `Free Applicative`

:

instance Functor (Bazaar a b) where fmap f (Bazaar e) = Bazaar $ \k -> fmap f (e k) instance Applicative (Bazaar a b) where pure x = Bazaar $ \_ -> pure x Bazaar ef < *> Bazaar ex = Bazaar $ \k -> ef k < *> ex k

This is actually analogous to the `Monoid`

instance for the free monoid; we just delegate to the underlying structure.

The more exciting thing is that we can fold and traverse over the first argument of `Bazaar`

, just like we can with the free monoid:

bfoldMap :: Monoid m => (a -> m) -> Bazaar a b t -> m bfoldMap f (Bazaar e) = getConst $ e (Const . f) newtype Comp g f a = Comp { getComp :: g (f a) } deriving (Functor) instance (Applicative f, Applicative g) => Applicative (Comp g f) where pure = Comp . pure . pure Comp f < *> Comp x = Comp $ liftA2 (< *>) f x btraverse :: (Applicative f) => (a -> f a') -> Bazaar a b t -> Bazaar a' b t btraverse f (Bazaar e) = getComp $ e (c . fmap ipure . f)

This is again analogous to the free monoid code. `Comp`

is the analogue of `Ap`

, and we use `ipure`

in `traverse`

. I mentioned that `Bazaar`

is a comonad:

extract :: Bazaar b b t -> t extract (Bazaar e) = runIdentity $ e Identity

And now we are finally prepared to not cheat:

honestTraverse :: (Applicative f, Traversable t) => (a -> f b) -> t a -> f (t b) honestTraverse f = fmap extract . btraverse f . howBizarre

So, we can traverse by first turning out `Traversable`

into some structure that's kind of like the free monoid, except having to do with `Applicative`

, traverse that, and then pull a result back out. `Bazaar`

retains the information that we're eventually building back the same type of structure, so we don't need any cheating.

To pull this back around to domains, there's nothing about this code to object to if done in a total language. But, if we think about our free `Applicative`

-ish structure, in Haskell, it will naturally allow infinitary expressions composed of the `Applicative`

operations, just like the free monoid will allow infinitary monoid expressions. And this is okay, because *some* `Applicative`

s can make sense of those, so throwing them away would make the type not free, in the same way that even finite lists are not the free monoid in Haskell. And this, I think, is compelling enough to say that infinite traversals are right for Haskell, just as they are wrong for Agda.

For those who wish to see executable code for all this, I've put a files here and here. The latter also contains some extra goodies at the end that I may talk about in further installments.

[1] Truth be told, I'm not exactly sure.

[2] It turns out, you can generalize `Bazaar`

to have a correspondence for every choice of `p`

newtype Bizarre p a b t = Bizarre { bizarre :: forall f. p f => (a -> f b) -> f t }

`hence`

and `forth`

above go through with the more general types. This can be seen here.

Sat 21 Feb 2015

[7] Comments

It is often stated that `Foldable`

is effectively the `toList`

class. However, this turns out to be wrong. The real fundamental member of `Foldable`

is `foldMap`

(which should look suspiciously like `traverse`

, incidentally). To understand exactly why this is, it helps to understand another surprising fact: lists are not free monoids in Haskell.

This latter fact can be seen relatively easily by considering another list-like type:

data SL a = Empty | SL a :> a instance Monoid (SL a) where mempty = Empty mappend ys Empty = ys mappend ys (xs :> x) = (mappend ys xs) :> x single :: a -> SL a single x = Empty :> x

So, we have a type `SL a`

of snoc lists, which are a monoid, and a function that embeds `a`

into `SL a`

. If (ordinary) lists were the free monoid, there would be a unique monoid homomorphism from lists to snoc lists. Such a homomorphism (call it `h`

) would have the following properties:

h [] = Empty h (xs <> ys) = h xs <> h ys h [x] = single x

And in fact, this (together with some general facts about Haskell functions) should be enough to define `h`

for our purposes (or any purposes, really). So, let's consider its behavior on two values:

h [1] = single 1 h [1,1..] = h ([1] <> [1,1..]) -- [1,1..] is an infinite list of 1s = h [1] <> h [1,1..]

This second equation can tell us what the value of `h`

is at this infinite value, since we can consider it the definition of a possibly infinite value:

x = h [1] <> x = fix (single 1 <>) h [1,1..] = x

`(single 1 <>)`

is a strict function, so the fixed point theorem tells us that `x = ⊥`

.

This is a problem, though. Considering some additional equations:

[1,1..] <> [n] = [1,1..] -- true for all n h [1,1..] = ⊥ h ([1,1..] <> [1]) = h [1,1..] <> h [1] = ⊥ <> single 1 = ⊥ :> 1 ≠ ⊥

So, our requirements for `h`

are contradictory, and no such homomorphism can exist.

The issue is that Haskell types are domains. They contain these extra partially defined values and infinite values. The monoid structure on (cons) lists has infinite lists absorbing all right-hand sides, while the snoc lists are just the opposite.

This also means that finite lists (or any method of implementing finite sequences) are not free monoids in Haskell. They, as domains, still contain the additional bottom element, and it absorbs all other elements, which is incorrect behavior for the free monoid:

pure x <> ⊥ = ⊥ h ⊥ = ⊥ h (pure x <> ⊥) = [x] <> h ⊥ = [x] ++ ⊥ = x:⊥ ≠ ⊥

So, what is the free monoid? In a sense, it can't be written down at all in Haskell, because we cannot enforce value-level equations, and because we don't have quotients. But, if conventions are good enough, there is a way. First, suppose we have a free monoid type `FM a`

. Then for any other monoid `m`

and embedding `a -> m`

, there must be a monoid homomorphism from `FM a`

to `m`

. We can model this as a Haskell type:

forall a m. Monoid m => (a -> m) -> FM a -> m

Where we consider the `Monoid m`

constraint to be enforcing that `m`

actually has valid monoid structure. Now, a trick is to recognize that this sort of universal property can be used to *define* types in Haskell (or, GHC at least), due to polymorphic types being first class; we just rearrange the arguments and quantifiers, and take `FM a`

to be the polymorphic type:

newtype FM a = FM { unFM :: forall m. Monoid m => (a -> m) -> m }

Types defined like this are automatically universal in the right sense. [1] The only thing we have to check is that `FM a`

is actually a monoid over `a`

. But that turns out to be easily witnessed:

embed :: a -> FM a embed x = FM $ \k -> k x instance Monoid (FM a) where mempty = FM $ \_ -> mempty mappend (FM e1) (FM e2) = FM $ \k -> e1 k <> e2 k

Demonstrating that the above is a proper monoid delegates to instances of `Monoid`

being proper monoids. So as long as we trust that convention, we have a free monoid.

However, one might wonder what a free monoid would look like as something closer to a traditional data type. To construct that, first ignore the required equations, and consider only the generators; we get:

```
data FMG a = None | Single a | FMG a :<> FMG a
```

Now, the proper `FM a`

is the quotient of this by the equations:

None :<> x = x = x :<> None x :<> (y :<> z) = (x :<> y) :<> z

One way of mimicking this in Haskell is to hide the implementation in a module, and only allow elimination into `Monoid`

s (again, using the convention that `Monoid`

ensures actual monoid structure) using the function:

unFMG :: forall a m. Monoid m => FMG a -> (a -> m) -> m unFMG None _ = mempty unFMG (Single x) k = k x unFMG (x :<> y) k = unFMG x k <> unFMG y k

This is actually how quotients can be thought of in richer languages; the quotient does not eliminate any of the generated structure internally, it just restricts the way in which the values can be consumed. Those richer languages just allow us to prove equations, and enforce properties by proof obligations, rather than conventions and structure hiding. Also, one should note that the above should look pretty similar to our encoding of `FM a`

using universal quantification earlier.

Now, one might look at the above and have some objections. For one, we'd normally think that the quotient of the above type is just `[a]`

. Second, it seems like the type is revealing something about the associativity of the operations, because defining recursive values via left nesting is different from right nesting, and this difference is observable by extracting into different monoids. But aren't monoids supposed to remove associativity as a concern? For instance:

ones1 = embed 1 <> ones1 ones2 = ones2 <> embed 1

Shouldn't we be able to prove these are the same, becuase of an argument like:

ones1 = embed 1 <> (embed 1 <> ...) ... reassociate ... = (... <> embed 1) <> embed 1 = ones2

The answer is that the equation we have only specifies the behavior of associating three values:

x <> (y <> z) = (x <> y) <> z

And while this is sufficient to nail down the behavior of *finite* values, and *finitary* reassociating, it does not tell us that *infinitary* reassociating yields the same value back. And the "... reassociate ..." step in the argument above was decidedly infinitary. And while the rules tell us that we can peel any finite number of copies of `embed 1`

to the front of `ones1`

or the end of `ones2`

, it does not tell us that `ones1 = ones2`

. And in fact it is vital for `FM a`

to have distinct values for these two things; it is what makes it the free monoid when we're dealing with domains of lazy values.

Finally, we can come back to `Foldable`

. If we look at `foldMap`

:

foldMap :: (Foldable f, Monoid m) => (a -> m) -> f a -> m

we can rearrange things a bit, and get the type:

Foldable f => f a -> (forall m. Monoid m => (a -> m) -> m)

And thus, the most fundamental operation of `Foldable`

is not `toList`

, but `toFreeMonoid`

, and lists are not free monoids in Haskell.

[1]: What we are doing here is noting that (co)limits are objects that internalize natural transformations, but the natural transformations expressible by quantification in GHC are already automatically internalized using quantifiers. However, one has to be careful that the quantifiers are actually enforcing the relevant naturality conditions. In many simple cases they are.

Tue 30 Dec 2014

[3] Comments

Emil Axelsson and Koen Claessen wrote a functional pearl last year about Using Circular Programs for Higher-Order Syntax.

About 6 months ago I had an opportunity to play with this approach in earnest, and realized we can speed it up a great deal. This has kept coming up in conversation ever since, so I've decided to write up an article here.

In my bound library I exploit the fact that monads are about substitution to make a monad transformer that manages substitution for me.

Here I'm going to take a more coupled approach.

To have a type system with enough complexity to be worth examining, I'll adapt Dan Doel's UPTS, which is a pure type system with universe polymorphism. I won't finish the implementation here, but from where we get it should be obvious how to finish the job.

Mon 1 Apr 2013

No Comments

A couple of weeks back one of my coworkers brought to my attention a several hour long workshop in Japan to go over and describe a number of my libraries, hosted by TANAKA Hideyuki — not the voice actor, I checked!

I was incredibly honored and I figured that if that many people (they had 30 or so registered attendees and 10 presentations) were going to spend that much time going over software that I had written, I should at least offer to show up!

I'd like to apologize for any errors in the romanization of people's names or misunderstandings I may have in the following text. My grasp of Japanese is very poor! Please feel free to send me corrections or additions!

Sadly, my boss's immediate reaction to hearing that there was a workshop in Japan about my work was to quip that "You're saying you're huge in Japan?" With him conspicuously not offering to fly me out here, I had to settle for surprising the organizers and attending via Google Hangout.

@nushio was very helpful in getting me connected, and while the speakers gave their talks I sat on the irc.freenode.net #haskell-lens channel and Google Hangout and answered questions and provided a running commentary with more details and references. Per freenode policy the fact that we were logging the channel was announced -- well, at least before things got too far underway.

Here is the IRC session log as a gist. IKEGAMI Daisuke @ikegami__ (`ikeg`

in the IRC log) tried to keep up a high-level running commentary about what was happening in the video to the log, which may be helpful if you are trying to follow along through each retroactively.

Other background chatter and material is strewn across twitter under the #ekmett_conf hash tag and on a japanese twitter aggregator named togetter

Sat 22 Sep 2012

[3] Comments

Recently, a fellow in category land discovered a fact that we in Haskell land have actually known for a while (in addition to things most of us probably don't). Specifically, given two categories and , a functor , and provided some conditions in hold, there exists a monad , the codensity monad of .

In category theory, the codensity monad is given by the rather frightening expression:

Wed 29 Aug 2012

Luite Stegeman has a mirror of the packages from Hackage.

He uses it to power his incredibly useful hdiff website.

During a Hackage outage, you can set up your local cabal configuration to point to it instead by (temporarily) replacing the remote-repo in your `~/.cabal/config`

file with:

remote-repo:

hdiff.luite.com:http://hdiff.luite.com/packages/archive

and then running `cabal update`

.

I have a `~/.cabal/config`

that I use whenever hackage goes down in my lens package.

If you use travis-ci, you can avoid build failures during hackage outages by first copying that config to ~/.cabal/config during before_install. -- You'll still be stuck waiting while it first tries to refresh from the real hackage server, but it only adds a few minutes to buildbot times.

Sun 24 Jun 2012

[9] Comments

Lenses are a great way to deal with functional references, but there are two common issues that arise from their use.

- There is a long-standing folklore position that lenses do not support polymorphic updates. This has actually caused a fair bit of embarrassment for the folks who'd like to incorporate lenses in any Haskell record system improvement.
- Access control. It'd be nice to have read-only or write-only properties -- "one-way" or "mirrored" lenses, as it were. Moreover, lenses are commonly viewed as an all or nothing proposition, in that it is hard to mix them with arbitrary user functions.
- Finally there is a bit of a cult around trying to generalize lenses by smashing a monad in the middle of them somewhere, it would be nice to be able to get into a list and work with each individual element in it without worrying about someone mucking up our lens laws, and perhaps avoid the whole generalized lens issue entirely.

We'll take a whack at each of these concerns in turn today.

(more...)

Sun 10 Jun 2012

No, I don't mean like this, but rather, If you spent any time trying to figure out xkcd's Umwelt April Fool comic this year, you may be interested in the Haskell source code. They used all sorts of information about you, the browser you were using, the resolution of your screen, to the geocoding of the network address you came from, etc. to serve up a custom web comic.

Today, davean posted to github the code for waldo, the engine he wrote to drive that comic.

Alas, he was not kind enough to actually supply the code for the umwelt comic strip itself, so you'll still be left wondering if the internet managed to find all of the Easter eggs. (Are they still Easter eggs when you release something a week before Easter?) You may find the list of links below useful if you want to get a feel for the different responses it gave people.

[ Article | xkcd's Forum | Hacker News | /r/haskell ]

**[Update: Jun 10, 9:09pm]** davean just posted a rather insightful post mortem of the development of waldo that talks a bit about why xkcd uses Haskell internally.

Mon 2 Apr 2012

In light of the burgeoning length of the ongoing record discussion sparked off by Simon Peyton-Jones in October, I would like to propose that we recognize an extension to Wadler's law (supplied in bold), which I'll refer to as the "Weak Record Conjecture" below.

In any language design, the total time spent discussing a feature in this list is proportional to two raised to the power of its position.

- 0. Semantics
- 1. Syntax
- 2. Lexical syntax
- 3. Lexical syntax of comments
4. Semantics of records

Sun 25 Dec 2011

[13] Comments

Andrej Bauer recently gave a really nice talk on how you can exploit side-effects to make a faster version of Martin Escardo's pseudo-paradoxical combinators.

A video of his talk is available over on his blog, and his presentation is remarkably clear, and would serve as a good preamble to the code I'm going to present below.

Andrej gave a related invited talk back at MSFP 2008 in Iceland, and afterwards over lunch I cornered him (with Dan Piponi) and explained how you could use parametricity to close over the side-effects of monads (or arrows, etc) but I think that trick was lost in the chaos of the weekend, so I've chosen to resurrect it here, and improve it to handle some of his more recent performance enhancements, and show that you don't need side-effects to speed up the search after all!

Thu 3 Nov 2011

[13] Comments

Thu 3 Nov 2011

[7] Comments

Max Bolingbroke has done a wonderful job on adding Constraint kinds to GHC.

Constraint Kinds adds a new kind `Constraint`

, such that `Eq :: * -> Constraint`

, `Monad :: (* -> *) -> Constraint`

, but since it is a kind, we can make type families for constraints, and even parameterize constraints *on* constraints.

So, let's play with them and see what we can come up with!

Thu 27 Oct 2011

No Comments

As requested, here are the slides from Dan Doel's excellent presentation on Homotopy and Directed Type Theory from this past Monday's Boston Haskell.

Tue 25 Oct 2011

I am very pleased to officially announce Hac Boston, a Haskell hackathon to be held January 20-22, 2012 at MIT in Cambridge, MA. The hackathon will officially kick off at 2:30 Friday afternoon, and go until 5pm on Sunday with the occasional break for sleep.

Everyone is welcome -- you do not have to be a Haskell guru to attend! Helping hack on someone else's project could be a great way to increase your Haskell skills.

If you plan on coming, please officially register, even if you already put your name on the wiki. Registration, travel, some information about lodging and many other details can now be found on the Hac Boston wiki.

We have confirmed space for about 30 people, so please register early! Beyond that we'll have to either seek additional space or close registration.

We're also looking for a few people interested in giving short (15-20 min.) talks, probably on Saturday afternoon. Anything of interest to the Haskell community is fair game---a project you've been working on, a paper, a quick tutorial. If you'd like to give a talk, add it on the wiki.

We look forward to seeing you at MIT!

Fri 23 Sep 2011

No Comments

Last time, I showed that we can build a small parsec clone with packrat support.

This time I intend to implement packrat directly on top of Parsec 3.

One of the main topics of discussion when it comes to packrat parsing since Bryan Ford's initial release of Pappy has been the fact that in general you shouldn't use packrat to memoize every rule, and that instead you should apply Amdahl's law to look for the cases where the lookup time is paid back in terms of repetitive evaluation, computation time and the hit rate. This is great news for us, since, we only want to memoize a handful of expensive combinators.

Fri 23 Sep 2011

[2] Comments

You never heard of the Millenium Falcon? It's the ship that made the Kessel Run in 12 parsecs.

I've been working on a parser combinator library called trifecta, and so I decided I'd share some thoughts on parsing.

Packrat parsing (as provided by frisby, pappy, rats! and the Scala parsing combinators) and more traditional recursive descent parsers (like Parsec) are often held up as somehow different.

Today I'll show that you can add monadic parsing to a packrat parser, sacrificing asymptotic guarantees in exchange for the convenient context sensitivity, and conversely how you can easily add packrat parsing to a traditional monadic parser combinator library.