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# Roles

The idea of *roles* comes from the paper Generative Type Abstraction and Type-level Computation, published at POPL 2011. The implementation of roles in GHC, however, is somewhat different than stated in that paper. This page focuses on the user-visible features of roles. RolesImplementation talks about the implementation in GHC.

## The problem we wish to solve

GHC has had a hole in its type system for several years, as documented in #1496, #4846, #5498, and #7148. The common cause behind all of this is the magic behind `-XGeneralizedNewtypeDeriving`. Here is an example:

newtype Age = MkAge { unAge :: Int } type family Inspect x type instance Inspect Age = Int type instance Inspect Int = Bool class BadIdea a where bad :: a -> Inspect a instance BadIdea Int where bad = (> 0) deriving instance BadIdea Age

This code is accepted by GHC 7.6.3. Yet, it goes wrong when you say `bad (MkAge 5)` -- we see the internal encoding of `Bool`! Let's trace what is happening here.

A newtype is a new algebraic datatype that wraps up exactly one field (in our example, of type `Int`). Yet, the semantics of Haskell makes a guarantee that wrapping and unwrapping a value (with `MkAge` or `unAge`) has no runtime cost. Thus, internally, we must consider `Age` to be wholly equivalent to `Int`.

The problem with this idea comes with type families. (There are other ways to tickle the bug, but one example is enough here.) A type family can branch on *Haskell* type, and of course, in Haskell (unlike in the internals of a compiler), `Age` is most certainly *not* `Int`. (If it were, newtypes would be useless for controlling instance selection, a very popular use case.) So, in our example, we see that `Inspect Age` is `Int`, but `Inspect Int` is `Bool`. Now, note the type of `bad`, the method in class `BadIdea`. When passed an `Int`, `bad` will return a `Bool`. When passed an `Age`, `bad` will return an `Int`. What happens on the last line above, when we use GeneralizedNewtypeDeriving? Internally, we take the existing instance for `Int` and just transmogrify it into an instance for `Age`. But, this transformation is very dumb -- because `Age` and `Int` are the same, internally, the code for the `Age` instance and the code for the `Int` instance are the same. This means that when we call `bad (MkAge 5)`, we run `5` through the existing implementation for `bad`, which produces a `Bool`. But, of course, the type of `bad (MkAge 5)` is `Int`, and so we have effectively converted a `Bool` to an `Int`. Yuck.

## The solution

What to do? It turns out we need a subtler definition of type equality than what we have had. Specifically, we must differentiate between *nominal* equality and *representational* equality (abbreviated to N and R, respectively). N equality (called C in the paper cited above) is the Haskell equality we all know and love. If two types have the same name, they are N-equal. If they don't have the same name (expanding type synonyms), they are not N-equal. R equality, on the other hand, shows that two types have the same *representation*. This is the equality that newtypes produce -- `Age` is R-equal to `Int`, but they are not N-equal.

Datatypes, classes, and type synonyms can be parametric in their type arguments or not. By "parametric", I mean that they do not *inspect* the type argument. A non-parametric type variable is inspect. Here are some examples:

data List a = Nil | Cons a (List a) -- parametric data GADT a where -- non-parametric GAge :: GADT Age GInt :: GADT Int class C1 a where -- parametric foo :: a -> List a class C2 a where -- non-parametric bar :: a -> GADT a class BadIdea a where -- non-parametric bad :: a -> Inspect a

In the terminology here, non-parametric types and classes care, in some fundamental way, what type parameter they are given. Parametric ones don't. We can generalize this idea a bit further to label each type variable as either parametric or not. For example,

data Mixed a b where MInt :: a -> Mixed a Int MAge :: a -> Mixed a Age

is parametric in its first parameter but not its second. For reasons that will soon become clear, we say that a parametric type variable has role R and a non-parametric one has role N.

With this notion in place, we simply need the following rule: **GeneralizedNewtypeDeriving works only with classes whose last type parameter is at role R.**

## Phantom parameters

It turns out that a third role is also useful (though unnecessary for type soundness): the phantom role, abbreviated P. It is often the case that programmers use type variables simply to constrain the type checker, not to make any statement about the runtime representation of a type. For example `data Phant a = MkPhant Int`. Because `a` doesn't appear on the right-hand side, we say that `a` is at role P. Why is this nice? Because it allows us to say that, say, `Phant Int` and `Phant Bool` are R-equal, because they really do have the same representation.

## Role inference

How do we know what role a type parameter should have? We use role inference! We start with a few base facts: `(->)` has two R parameters; `(~)` has two N parameters; and all type families' parameters are N. Then, we just propagate the information. By defaulting parameters to role P, any parameters unused in the right-hand side (or used only in other types in P positions) will be P. Whenever a parameter is used in an R position (that is, used as a type argument to a constructor whose corresponding variable is at role R), we raise its role from P to R. Similarly, when a parameter is used in an N position, its role is upgraded to N. We never downgrade a role from N to P or R, or from R to P. In this way, we infer the most-general role for each parameter.

## Role annotations

As we have learned with type and kind inference, sometimes the programmer wants to constrain the inference process. For example, the base library contains the following definition:

data Ptr a = Ptr Addr#

The idea is that `a` should really be an R parameter, but role inference assigns it to P. This makes some level of sense: a pointer to an `Int` really *is* representationally the same as a pointer to a `Bool`. But, that's not at all how we want to use `Ptr`s! So, we want to be able to say

data Ptr a@R = Ptr Addr#

The `@R` annotation forces the parameter `a` to be at role R, not role P. We, then, of course, need to *check* the user-supplied roles to make sure they don't break any promises. It would be bad if the user could make `BadIdea`'s role be R!

The other place where role annotations may be necessary are in .hs-boot files, where the right-hand sides of definitions can be omitted. As usual, the types/classes declared in an .hs-boot file must match up with the definitions in the .hs file, including down to the roles. The default role will be R in hs-boot files, corresponding to the common use case. Note that this **will break code**. But, the change is necessary to close the type-safety hole discussed above.

Role annotations will be allowed on type variables in `data`, `newtype`, `class`, and `type` declarations. They will not be allowed on type/data family declarations or in explicit `forall`s in function type signatures.

## Role *and* kind annotations

What if the user wants both a role and a kind annotation on a type variable? There are two possibilities:

data Foo (a :: k)@R

data Bar (a@R :: k)

I (Richard E) propose the syntax for `Foo`, for no reason I can articulate. Note that the parentheses enclosing the kind annotation are required whether or not there is a role annotation.

## Roles and `Traversable`

Though a minor issue in the overall scheme, work on Roles had led to an interesting interaction with `Traversable`, excerpted here:

class Traversable t where traverse :: Applicative f => (a -> f b) -> t a -> f (t b)

According to the rules for roles, the parameter `t` must be at role N, as it is used as a parameter to the type variable `f`. We must account for the possibility that `f` will be instantiated with a type whose last parameter is at role N, so we force `t` to be at role N as well.

This means that GeneralizedNewtypeDeriving (hereafter "GND") no longer works with Traversable. But, DeriveTraversable *does* still work. However, GHC previously preferred using GND over DeriveTraversable when both were available, which assumedly produced the same code. How is this all possible? If GND doesn't work anymore, is there something wrong with DeriveTraversable? The answer is that GND and DeriveTraversable make *different* instances, contrary to expectations. The reason is that DeriveTraversable has to use `fmap` to cast the result of `traverse` from the representation type back to the newtype. According to the Functor laws, `fmap`ping this cast should be a no-op (the law of interest is `fmap id == id`). But, if that law is not obeyed, `fmap`ping the cast may change the result of the `traverse`. Contrast this with a GND instance, which magically casts the result, without using `fmap`. If the Functor law is not obeyed, these two instances have different behavior.

Despite this, I believe that using GND with `Traversable` is indeed type-safe. Why? Because of the parametricity guaranteed in `Functor` and `Applicative`. The reason GND is prohibited with `Traversable` is that we are worried `f`'s last parameter will be at role N. While it is possible to write `Functor` and `Applicative` instances for such a type, the methods of those classes can't really use the any constructors that force the role to be N. For example, consider this:

data G a where GInt :: a -> G Int Ga :: a -> G a instance Functor G where fmap f (GInt _) = error "urk" -- no way out here fmap f (Ga a) = Ga (f a) instance Applicative G where pure a = Ga a (Ga f) <*> (Ga a) = Ga (f a) _ <*> _ = error "urk" -- no way out here, either

There's no way to usefully interact with the `GInt` constructor and get the code to type-check. Thus, I believe (but haven't yet proved), that using GND with `Traversable` is safe, because the `f` in `traverse` can't ever do bad things with its argument. If you, the reader, have more insight into this (or a counterexample!), please write!