{-# LANGUAGE CPP #-}
module TcFlatten(
FlattenMode(..),
flatten, flattenManyNom,
unflatten,
eqCanRewrite, eqCanRewriteFR, canRewriteOrSame,
CtFlavourRole, ctEvFlavourRole, ctFlavourRole
) where
#include "HsVersions.h"
import TcRnTypes
import TcType
import Type
import TcEvidence
import TyCon
import TypeRep
import Kind( isSubKind )
import Coercion ( tyConRolesX )
import Var
import VarEnv
import NameEnv
import Outputable
import TcSMonad as TcS
import DynFlags( DynFlags )
import Util
import Bag
import FastString
import Control.Monad
import MonadUtils ( zipWithAndUnzipM )
import GHC.Exts ( inline )
#if __GLASGOW_HASKELL__ < 709
import Control.Applicative ( Applicative(..), (<$>) )
#endif
{-
Note [The flattening story]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
* A CFunEqCan is either of form
[G] : F xis ~ fsk -- fsk is a FlatSkol
[W] x : F xis ~ fmv -- fmv is a unification variable,
-- but untouchable,
-- with MetaInfo = FlatMetaTv
where
x is the witness variable
fsk/fmv is a flatten skolem
xis are function-free
CFunEqCans are always [Wanted], or [Given], never [Derived]
fmv untouchable just means that in a CTyVarEq, say,
fmv ~ Int
we do NOT unify fmv.
* KEY INSIGHTS:
- A given flatten-skolem, fsk, is known a-priori to be equal to
F xis (the LHS), with evidence
- A unification flatten-skolem, fmv, stands for the as-yet-unknown
type to which (F xis) will eventually reduce
* Inert set invariant: if F xis1 ~ fsk1, F xis2 ~ fsk2
then xis1 /= xis2
i.e. at most one CFunEqCan with a particular LHS
* Each canonical CFunEqCan x : F xis ~ fsk/fmv has its own
distinct evidence variable x and flatten-skolem fsk/fmv.
Why? We make a fresh fsk/fmv when the constraint is born;
and we never rewrite the RHS of a CFunEqCan.
* Function applications can occur in the RHS of a CTyEqCan. No reason
not allow this, and it reduces the amount of flattening that must occur.
* Flattening a type (F xis):
- If we are flattening in a Wanted/Derived constraint
then create new [W] x : F xis ~ fmv
else create new [G] x : F xis ~ fsk
with fresh evidence variable x and flatten-skolem fsk/fmv
- Add it to the work list
- Replace (F xis) with fsk/fmv in the type you are flattening
- You can also add the CFunEqCan to the "flat cache", which
simply keeps track of all the function applications you
have flattened.
- If (F xis) is in the cache already, just
use its fsk/fmv and evidence x, and emit nothing.
- No need to substitute in the flat-cache. It's not the end
of the world if we start with, say (F alpha ~ fmv1) and
(F Int ~ fmv2) and then find alpha := Int. Athat will
simply give rise to fmv1 := fmv2 via [Interacting rule] below
* Canonicalising a CFunEqCan [G/W] x : F xis ~ fsk/fmv
- Flatten xis (to substitute any tyvars; there are already no functions)
cos :: xis ~ flat_xis
- New wanted x2 :: F flat_xis ~ fsk/fmv
- Add new wanted to flat cache
- Discharge x = F cos ; x2
* Unification flatten-skolems, fmv, ONLY get unified when either
a) The CFunEqCan takes a step, using an axiom
b) During un-flattening
They are never unified in any other form of equality.
For example [W] ffmv ~ Int is stuck; it does not unify with fmv.
* We *never* substitute in the RHS (i.e. the fsk/fmv) of a CFunEqCan.
That would destroy the invariant about the shape of a CFunEqCan,
and it would risk wanted/wanted interactions. The only way we
learn information about fsk is when the CFunEqCan takes a step.
However we *do* substitute in the LHS of a CFunEqCan (else it
would never get to fire!)
* [Interacting rule]
(inert) [W] x1 : F tys ~ fmv1
(work item) [W] x2 : F tys ~ fmv2
Just solve one from the other:
x2 := x1
fmv2 := fmv1
This just unites the two fsks into one.
Always solve given from wanted if poss.
* For top-level reductions, see Note [Top-level reductions for type functions]
in TcInteract
Why given-fsks, alone, doesn't work
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Could we get away with only flatten meta-tyvars, with no flatten-skolems? No.
[W] w : alpha ~ [F alpha Int]
---> flatten
w = ...w'...
[W] w' : alpha ~ [fsk]
[G] : F alpha Int ~ fsk
--> unify (no occurs check)
alpha := [fsk]
But since fsk = F alpha Int, this is really an occurs check error. If
that is all we know about alpha, we will succeed in constraint
solving, producing a program with an infinite type.
Even if we did finally get (g : fsk ~ Boo)l by solving (F alpha Int ~ fsk)
using axiom, zonking would not see it, so (x::alpha) sitting in the
tree will get zonked to an infinite type. (Zonking always only does
refl stuff.)
Why flatten-meta-vars, alone doesn't work
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Look at Simple13, with unification-fmvs only
[G] g : a ~ [F a]
---> Flatten given
g' = g;[x]
[G] g' : a ~ [fmv]
[W] x : F a ~ fmv
--> subst a in x
x = F g' ; x2
[W] x2 : F [fmv] ~ fmv
And now we have an evidence cycle between g' and x!
If we used a given instead (ie current story)
[G] g : a ~ [F a]
---> Flatten given
g' = g;[x]
[G] g' : a ~ [fsk]
[G] : F a ~ fsk
---> Substitute for a
[G] g' : a ~ [fsk]
[G] F (sym g'); : F [fsk] ~ fsk
Why is it right to treat fmv's differently to ordinary unification vars?
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
f :: forall a. a -> a -> Bool
g :: F Int -> F Int -> Bool
Consider
f (x:Int) (y:Bool)
This gives alpha~Int, alpha~Bool. There is an inconsistency,
but really only one error. SherLoc may tell you which location
is most likely, based on other occurrences of alpha.
Consider
g (x:Int) (y:Bool)
Here we get (F Int ~ Int, F Int ~ Bool), which flattens to
(fmv ~ Int, fmv ~ Bool)
But there are really TWO separate errors. We must not complain
about Int~Bool. Moreover these two errors could arise in entirely
unrelated parts of the code. (In the alpha case, there must be
*some* connection (eg v:alpha in common envt).)
Note [Orient equalities with flatten-meta-vars on the left]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
This example comes from IndTypesPerfMerge
From the ambiguity check for
f :: (F a ~ a) => a
we get:
[G] F a ~ a
[W] F alpha ~ alpha, alpha ~ a
From Givens we get
[G] F a ~ fsk, fsk ~ a
Now if we flatten we get
[W] alpha ~ fmv, F alpha ~ fmv, alpha ~ a
Now, processing the first one first, choosing alpha := fmv
[W] F fmv ~ fmv, fmv ~ a
And now we are stuck. We must either *unify* fmv := a, or
use the fmv ~ a to rewrite F fmv ~ fmv, so we can make it
meet up with the given F a ~ blah.
Solution: always put fmvs on the left, so we get
[W] fmv ~ alpha, F alpha ~ fmv, alpha ~ a
The point is that fmvs are very uninformative, so doing alpha := fmv
is a bad idea. We want to use other constraints on alpha first.
************************************************************************
* *
* Other notes (Oct 14)
I have not revisted these, but I didn't want to discard them
* *
************************************************************************
Try: rewrite wanted with wanted only for fmvs (not all meta-tyvars)
But: fmv ~ alpha[0]
alpha[0] ~ fmv’
Now we don’t see that fmv ~ fmv’, which is a problem for injectivity detection.
Conclusion: rewrite wanteds with wanted for all untouchables.
skol ~ untch, must re-orieint to untch ~ skol, so that we can use it to rewrite.
************************************************************************
* *
* Examples
Here is a long series of examples I had to work through
* *
************************************************************************
Simple20
~~~~~~~~
axiom F [a] = [F a]
[G] F [a] ~ a
-->
[G] fsk ~ a
[G] [F a] ~ fsk (nc)
-->
[G] F a ~ fsk2
[G] fsk ~ [fsk2]
[G] fsk ~ a
-->
[G] F a ~ fsk2
[G] a ~ [fsk2]
[G] fsk ~ a
-----------------------------------
----------------------------------------
indexed-types/should_compile/T44984
[W] H (F Bool) ~ H alpha
[W] alpha ~ F Bool
-->
F Bool ~ fmv0
H fmv0 ~ fmv1
H alpha ~ fmv2
fmv1 ~ fmv2
fmv0 ~ alpha
flatten
~~~~~~~
fmv0 := F Bool
fmv1 := H (F Bool)
fmv2 := H alpha
alpha := F Bool
plus
fmv1 ~ fmv2
But these two are equal under the above assumptions.
Solve by Refl.
--- under plan B, namely solve fmv1:=fmv2 eagerly ---
[W] H (F Bool) ~ H alpha
[W] alpha ~ F Bool
-->
F Bool ~ fmv0
H fmv0 ~ fmv1
H alpha ~ fmv2
fmv1 ~ fmv2
fmv0 ~ alpha
-->
F Bool ~ fmv0
H fmv0 ~ fmv1
H alpha ~ fmv2 fmv2 := fmv1
fmv0 ~ alpha
flatten
fmv0 := F Bool
fmv1 := H fmv0 = H (F Bool)
retain H alpha ~ fmv2
because fmv2 has been filled
alpha := F Bool
----------------------------
indexed-types/should_failt/T4179
after solving
[W] fmv_1 ~ fmv_2
[W] A3 (FCon x) ~ fmv_1 (CFunEqCan)
[W] A3 (x (aoa -> fmv_2)) ~ fmv_2 (CFunEqCan)
----------------------------------------
indexed-types/should_fail/T7729a
a) [W] BasePrimMonad (Rand m) ~ m1
b) [W] tt m1 ~ BasePrimMonad (Rand m)
---> process (b) first
BasePrimMonad (Ramd m) ~ fmv_atH
fmv_atH ~ tt m1
---> now process (a)
m1 ~ s_atH ~ tt m1 -- An obscure occurs check
----------------------------------------
typecheck/TcTypeNatSimple
Original constraint
[W] x + y ~ x + alpha (non-canonical)
==>
[W] x + y ~ fmv1 (CFunEqCan)
[W] x + alpha ~ fmv2 (CFuneqCan)
[W] fmv1 ~ fmv2 (CTyEqCan)
(sigh)
----------------------------------------
indexed-types/should_fail/GADTwrong1
[G] Const a ~ ()
==> flatten
[G] fsk ~ ()
work item: Const a ~ fsk
==> fire top rule
[G] fsk ~ ()
work item fsk ~ ()
Surely the work item should rewrite to () ~ ()? Well, maybe not;
it'a very special case. More generally, our givens look like
F a ~ Int, where (F a) is not reducible.
----------------------------------------
indexed_types/should_fail/T8227:
Why using a different can-rewrite rule in CFunEqCan heads
does not work.
Assuming NOT rewriting wanteds with wanteds
Inert: [W] fsk_aBh ~ fmv_aBk -> fmv_aBk
[W] fmv_aBk ~ fsk_aBh
[G] Scalar fsk_aBg ~ fsk_aBh
[G] V a ~ f_aBg
Worklist includes [W] Scalar fmv_aBi ~ fmv_aBk
fmv_aBi, fmv_aBk are flatten unificaiton variables
Work item: [W] V fsk_aBh ~ fmv_aBi
Note that the inert wanteds are cyclic, because we do not rewrite
wanteds with wanteds.
Then we go into a loop when normalise the work-item, because we
use rewriteOrSame on the argument of V.
Conclusion: Don't make canRewrite context specific; instead use
[W] a ~ ty to rewrite a wanted iff 'a' is a unification variable.
----------------------------------------
Here is a somewhat similar case:
type family G a :: *
blah :: (G a ~ Bool, Eq (G a)) => a -> a
blah = error "urk"
foo x = blah x
For foo we get
[W] Eq (G a), G a ~ Bool
Flattening
[W] G a ~ fmv, Eq fmv, fmv ~ Bool
We can't simplify away the Eq Bool unless we substitute for fmv.
Maybe that doesn't matter: we would still be left with unsolved
G a ~ Bool.
--------------------------
Trac #9318 has a very simple program leading to
[W] F Int ~ Int
[W] F Int ~ Bool
We don't want to get "Error Int~Bool". But if fmv's can rewrite
wanteds, we will
[W] fmv ~ Int
[W] fmv ~ Bool
--->
[W] Int ~ Bool
************************************************************************
* *
* FlattenEnv & FlatM
* The flattening environment & monad
* *
************************************************************************
-}
type FlatWorkListRef = TcRef [Ct] -- See Note [The flattening work list]
data FlattenEnv
= FE { fe_mode :: FlattenMode
, fe_loc :: CtLoc -- See Note [Flattener CtLoc]
, fe_flavour :: CtFlavour
, fe_eq_rel :: EqRel -- See Note [Flattener EqRels]
, fe_work :: FlatWorkListRef } -- See Note [The flattening work list]
data FlattenMode -- Postcondition for all three: inert wrt the type substitution
= FM_FlattenAll -- Postcondition: function-free
| FM_SubstOnly -- See Note [Flattening under a forall]
-- | FM_Avoid TcTyVar Bool -- See Note [Lazy flattening]
-- -- Postcondition:
-- -- * tyvar is only mentioned in result under a rigid path
-- -- e.g. [a] is ok, but F a won't happen
-- -- * If flat_top is True, top level is not a function application
-- -- (but under type constructors is ok e.g. [F a])
mkFlattenEnv :: FlattenMode -> CtEvidence -> FlatWorkListRef -> FlattenEnv
mkFlattenEnv fm ctev ref = FE { fe_mode = fm
, fe_loc = ctEvLoc ctev
, fe_flavour = ctEvFlavour ctev
, fe_eq_rel = ctEvEqRel ctev
, fe_work = ref }
-- | The 'FlatM' monad is a wrapper around 'TcS' with the following
-- extra capabilities: (1) it offers access to a 'FlattenEnv';
-- and (2) it maintains the flattening worklist.
-- See Note [The flattening work list].
newtype FlatM a
= FlatM { runFlatM :: FlattenEnv -> TcS a }
instance Monad FlatM where
return x = FlatM $ const (return x)
m >>= k = FlatM $ \env ->
do { a <- runFlatM m env
; runFlatM (k a) env }
instance Functor FlatM where
fmap = liftM
instance Applicative FlatM where
pure = return
(<*>) = ap
liftTcS :: TcS a -> FlatM a
liftTcS thing_inside
= FlatM $ const thing_inside
emitFlatWork :: Ct -> FlatM ()
-- See Note [The flattening work list]
emitFlatWork ct = FlatM $ \env -> updTcRef (fe_work env) (ct :)
runFlatten :: FlattenMode -> CtEvidence -> FlatM a -> TcS a
-- Run thing_inside (which does flattening), and put all
-- the work it generates onto the main work list
-- See Note [The flattening work list]
-- NB: The returned evidence is always the same as the original, but with
-- perhaps a new CtLoc
runFlatten mode ev thing_inside
= do { flat_ref <- newTcRef []
; let fmode = mkFlattenEnv mode ev flat_ref
; res <- runFlatM thing_inside fmode
; new_flats <- readTcRef flat_ref
; updWorkListTcS (add_flats new_flats)
; return res }
where
add_flats new_flats wl
= wl { wl_funeqs = add_funeqs new_flats (wl_funeqs wl) }
add_funeqs [] wl = wl
add_funeqs (f:fs) wl = add_funeqs fs (f:wl)
-- add_funeqs fs ws = reverse fs ++ ws
-- e.g. add_funeqs [f1,f2,f3] [w1,w2,w3,w4]
-- = [f3,f2,f1,w1,w2,w3,w4]
traceFlat :: String -> SDoc -> FlatM ()
traceFlat herald doc = liftTcS $ traceTcS herald doc
getFlatEnvField :: (FlattenEnv -> a) -> FlatM a
getFlatEnvField accessor
= FlatM $ \env -> return (accessor env)
getEqRel :: FlatM EqRel
getEqRel = getFlatEnvField fe_eq_rel
getRole :: FlatM Role
getRole = eqRelRole <$> getEqRel
getFlavour :: FlatM CtFlavour
getFlavour = getFlatEnvField fe_flavour
getFlavourRole :: FlatM CtFlavourRole
getFlavourRole
= do { flavour <- getFlavour
; eq_rel <- getEqRel
; return (flavour, eq_rel) }
getMode :: FlatM FlattenMode
getMode = getFlatEnvField fe_mode
getLoc :: FlatM CtLoc
getLoc = getFlatEnvField fe_loc
checkStackDepth :: Type -> FlatM ()
checkStackDepth ty
= do { loc <- getLoc
; liftTcS $ checkReductionDepth loc ty }
-- | Change the 'EqRel' in a 'FlatM'.
setEqRel :: EqRel -> FlatM a -> FlatM a
setEqRel new_eq_rel thing_inside
= FlatM $ \env ->
if new_eq_rel == fe_eq_rel env
then runFlatM thing_inside env
else runFlatM thing_inside (env { fe_eq_rel = new_eq_rel })
-- | Change the 'FlattenMode' in a 'FlattenEnv'.
setMode :: FlattenMode -> FlatM a -> FlatM a
setMode new_mode thing_inside
= FlatM $ \env ->
if new_mode `eq` fe_mode env
then runFlatM thing_inside env
else runFlatM thing_inside (env { fe_mode = new_mode })
where
FM_FlattenAll `eq` FM_FlattenAll = True
FM_SubstOnly `eq` FM_SubstOnly = True
-- FM_Avoid tv1 b1 `eq` FM_Avoid tv2 b2 = tv1 == tv2 && b1 == b2
_ `eq` _ = False
bumpDepth :: FlatM a -> FlatM a
bumpDepth (FlatM thing_inside)
= FlatM $ \env -> do { let env' = env { fe_loc = bumpCtLocDepth (fe_loc env) }
; thing_inside env' }
-- Flatten skolems
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
newFlattenSkolemFlatM :: TcType -- F xis
-> FlatM (CtEvidence, TcTyVar) -- [W] x:: F xis ~ fsk
newFlattenSkolemFlatM ty
= do { flavour <- getFlavour
; loc <- getLoc
; liftTcS $ newFlattenSkolem flavour loc ty }
{-
Note [The flattening work list]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The "flattening work list", held in the fe_work field of FlattenEnv,
is a list of CFunEqCans generated during flattening. The key idea
is this. Consider flattening (Eq (F (G Int) (H Bool)):
* The flattener recursively calls itself on sub-terms before building
the main term, so it will encounter the terms in order
G Int
H Bool
F (G Int) (H Bool)
flattening to sub-goals
w1: G Int ~ fuv0
w2: H Bool ~ fuv1
w3: F fuv0 fuv1 ~ fuv2
* Processing w3 first is BAD, because we can't reduce i t,so it'll
get put into the inert set, and later kicked out when w1, w2 are
solved. In Trac #9872 this led to inert sets containing hundreds
of suspended calls.
* So we want to process w1, w2 first.
* So you might think that we should just use a FIFO deque for the work-list,
so that putting adding goals in order w1,w2,w3 would mean we processed
w1 first.
* BUT suppose we have 'type instance G Int = H Char'. Then processing
w1 leads to a new goal
w4: H Char ~ fuv0
We do NOT want to put that on the far end of a deque! Instead we want
to put it at the *front* of the work-list so that we continue to work
on it.
So the work-list structure is this:
* The wl_funeqs (in TcS) is a LIFO stack; we push new goals (such as w4) on
top (extendWorkListFunEq), and take new work from the top
(selectWorkItem).
* When flattening, emitFlatWork pushes new flattening goals (like
w1,w2,w3) onto the flattening work list, fe_work, another
push-down stack.
* When we finish flattening, we *reverse* the fe_work stack
onto the wl_funeqs stack (which brings w1 to the top).
The function runFlatten initialises the fe_work stack, and reverses
it onto wl_fun_eqs at the end.
Note [Flattener EqRels]
~~~~~~~~~~~~~~~~~~~~~~~
When flattening, we need to know which equality relation -- nominal
or representation -- we should be respecting. The only difference is
that we rewrite variables by representational equalities when fe_eq_rel
is ReprEq, and that we unwrap newtypes when flattening w.r.t.
representational equality.
Note [Flattener CtLoc]
~~~~~~~~~~~~~~~~~~~~~~
The flattener does eager type-family reduction.
Type families might loop, and we
don't want GHC to do so. A natural solution is to have a bounded depth
to these processes. A central difficulty is that such a solution isn't
quite compositional. For example, say it takes F Int 10 steps to get to Bool.
How many steps does it take to get from F Int -> F Int to Bool -> Bool?
10? 20? What about getting from Const Char (F Int) to Char? 11? 1? Hard to
know and hard to track. So, we punt, essentially. We store a CtLoc in
the FlattenEnv and just update the environment when recurring. In the
TyConApp case, where there may be multiple type families to flatten,
we just copy the current CtLoc into each branch. If any branch hits the
stack limit, then the whole thing fails.
A consequence of this is that setting the stack limits appropriately
will be essentially impossible. So, the official recommendation if a
stack limit is hit is to disable the check entirely. Otherwise, there
will be baffling, unpredictable errors.
Note [Lazy flattening]
~~~~~~~~~~~~~~~~~~~~~~
The idea of FM_Avoid mode is to flatten less aggressively. If we have
a ~ [F Int]
there seems to be no great merit in lifting out (F Int). But if it was
a ~ [G a Int]
then we *do* want to lift it out, in case (G a Int) reduces to Bool, say,
which gets rid of the occurs-check problem. (For the flat_top Bool, see
comments above and at call sites.)
HOWEVER, the lazy flattening actually seems to make type inference go
*slower*, not faster. perf/compiler/T3064 is a case in point; it gets
*dramatically* worse with FM_Avoid. I think it may be because
floating the types out means we normalise them, and that often makes
them smaller and perhaps allows more re-use of previously solved
goals. But to be honest I'm not absolutely certain, so I am leaving
FM_Avoid in the code base. What I'm removing is the unique place
where it is *used*, namely in TcCanonical.canEqTyVar.
See also Note [Conservative unification check] in TcUnify, which gives
other examples where lazy flattening caused problems.
Bottom line: FM_Avoid is unused for now (Nov 14).
Note: T5321Fun got faster when I disabled FM_Avoid
T5837 did too, but it's pathalogical anyway
Note [Phantoms in the flattener]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have
data Proxy p = Proxy
and we're flattening (Proxy ty) w.r.t. ReprEq. Then, we know that `ty`
is really irrelevant -- it will be ignored when solving for representational
equality later on. So, we omit flattening `ty` entirely. This may
violate the expectation of "xi"s for a bit, but the canonicaliser will
soon throw out the phantoms when decomposing a TyConApp. (Or, the
canonicaliser will emit an insoluble, in which case the unflattened version
yields a better error message anyway.)
-}
{- *********************************************************************
* *
* Externally callable flattening functions *
* *
* They are all wrapped in runFlatten, so their *
* flattening work gets put into the work list *
* *
********************************************************************* -}
flatten :: FlattenMode -> CtEvidence -> TcType
-> TcS (Xi, TcCoercion)
flatten mode ev ty
= runFlatten mode ev (flatten_one ty)
flattenManyNom :: CtEvidence -> [TcType] -> TcS ([Xi], [TcCoercion])
-- Externally-callable, hence runFlatten
-- Flatten a bunch of types all at once; in fact they are
-- always the arguments of a saturated type-family, so
-- ctEvFlavour ev = Nominal
-- and we want to flatten all at nominal role
flattenManyNom ev tys
= runFlatten FM_FlattenAll ev (flatten_many_nom tys)
{- *********************************************************************
* *
* The main flattening functions
* *
********************************************************************* -}
{- Note [Flattening]
~~~~~~~~~~~~~~~~~~~~
flatten ty ==> (xi, cc)
where
xi has no type functions, unless they appear under ForAlls
cc = Auxiliary given (equality) constraints constraining
the fresh type variables in xi. Evidence for these
is always the identity coercion, because internally the
fresh flattening skolem variables are actually identified
with the types they have been generated to stand in for.
Note that it is flatten's job to flatten *every type function it sees*.
flatten is only called on *arguments* to type functions, by canEqGiven.
Recall that in comments we use alpha[flat = ty] to represent a
flattening skolem variable alpha which has been generated to stand in
for ty.
----- Example of flattening a constraint: ------
flatten (List (F (G Int))) ==> (xi, cc)
where
xi = List alpha
cc = { G Int ~ beta[flat = G Int],
F beta ~ alpha[flat = F beta] }
Here
* alpha and beta are 'flattening skolem variables'.
* All the constraints in cc are 'given', and all their coercion terms
are the identity.
NB: Flattening Skolems only occur in canonical constraints, which
are never zonked, so we don't need to worry about zonking doing
accidental unflattening.
Note that we prefer to leave type synonyms unexpanded when possible,
so when the flattener encounters one, it first asks whether its
transitive expansion contains any type function applications. If so,
it expands the synonym and proceeds; if not, it simply returns the
unexpanded synonym.
Note [flatten_many performance]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In programs with lots of type-level evaluation, flatten_many becomes
part of a tight loop. For example, see test perf/compiler/T9872a, which
calls flatten_many a whopping 7,106,808 times. It is thus important
that flatten_many be efficient.
Performance testing showed that the current implementation is indeed
efficient. It's critically important that zipWithAndUnzipM be
specialized to TcS, and it's also quite helpful to actually `inline`
it. On test T9872a, here are the allocation stats (Dec 16, 2014):
* Unspecialized, uninlined: 8,472,613,440 bytes allocated in the heap
* Specialized, uninlined: 6,639,253,488 bytes allocated in the heap
* Specialized, inlined: 6,281,539,792 bytes allocated in the heap
To improve performance even further, flatten_many_nom is split off
from flatten_many, as nominal equality is the common case. This would
be natural to write using mapAndUnzipM, but even inlined, that function
is not as performant as a hand-written loop.
* mapAndUnzipM, inlined: 7,463,047,432 bytes allocated in the heap
* hand-written recursion: 5,848,602,848 bytes allocated in the heap
If you make any change here, pay close attention to the T9872{a,b,c} tests
and T5321Fun.
If we need to make this yet more performant, a possible way forward is to
duplicate the flattener code for the nominal case, and make that case
faster. This doesn't seem quite worth it, yet.
-}
flatten_many :: [Role] -> [Type] -> FlatM ([Xi], [TcCoercion])
-- Coercions :: Xi ~ Type, at roles given
-- Returns True iff (no flattening happened)
-- NB: The EvVar inside the 'fe_ev :: CtEvidence' is unused,
-- we merely want (a) Given/Solved/Derived/Wanted info
-- (b) the GivenLoc/WantedLoc for when we create new evidence
flatten_many roles tys
-- See Note [flatten_many performance]
= inline zipWithAndUnzipM go roles tys
where
go Nominal ty = setEqRel NomEq $ flatten_one ty
go Representational ty = setEqRel ReprEq $ flatten_one ty
go Phantom ty = -- See Note [Phantoms in the flattener]
return (ty, mkTcPhantomCo ty ty)
-- | Like 'flatten_many', but assumes that every role is nominal.
flatten_many_nom :: [Type] -> FlatM ([Xi], [TcCoercion])
flatten_many_nom [] = return ([], [])
-- See Note [flatten_many performance]
flatten_many_nom (ty:tys)
= do { (xi, co) <- flatten_one ty
; (xis, cos) <- flatten_many_nom tys
; return (xi:xis, co:cos) }
------------------
flatten_one :: TcType -> FlatM (Xi, TcCoercion)
-- Flatten a type to get rid of type function applications, returning
-- the new type-function-free type, and a collection of new equality
-- constraints. See Note [Flattening] for more detail.
--
-- Postcondition: Coercion :: Xi ~ TcType
-- The role on the result coercion matches the EqRel in the FlattenEnv
flatten_one xi@(LitTy {})
= do { role <- getRole
; return (xi, mkTcReflCo role xi) }
flatten_one (TyVarTy tv)
= do { mb_yes <- flatten_tyvar tv
; role <- getRole
; case mb_yes of
Left tv' -> -- Done
do { traceFlat "flattenTyVar1" (ppr tv $$ ppr (tyVarKind tv'))
; return (ty', mkTcReflCo role ty') }
where
ty' = mkTyVarTy tv'
Right (ty1, co1) -- Recurse
-> do { (ty2, co2) <- flatten_one ty1
; traceFlat "flattenTyVar2" (ppr tv $$ ppr ty2)
; return (ty2, co2 `mkTcTransCo` co1) } }
flatten_one (AppTy ty1 ty2)
= do { (xi1,co1) <- flatten_one ty1
; eq_rel <- getEqRel
; case (eq_rel, nextRole xi1) of
(NomEq, _) -> flatten_rhs xi1 co1 NomEq
(ReprEq, Nominal) -> flatten_rhs xi1 co1 NomEq
(ReprEq, Representational) -> flatten_rhs xi1 co1 ReprEq
(ReprEq, Phantom) ->
return (mkAppTy xi1 ty2, co1 `mkTcAppCo` mkTcNomReflCo ty2) }
where
flatten_rhs xi1 co1 eq_rel2
= do { (xi2,co2) <- setEqRel eq_rel2 $ flatten_one ty2
; traceFlat "flatten/appty"
(ppr ty1 $$ ppr ty2 $$ ppr xi1 $$
ppr co1 $$ ppr xi2)
; role1 <- getRole
; let role2 = eqRelRole eq_rel2
; return ( mkAppTy xi1 xi2
, mkTcTransAppCo role1 co1 xi1 ty1
role2 co2 xi2 ty2
role1 ) } -- output should match fmode
flatten_one (FunTy ty1 ty2)
= do { (xi1,co1) <- flatten_one ty1
; (xi2,co2) <- flatten_one ty2
; role <- getRole
; return (mkFunTy xi1 xi2, mkTcFunCo role co1 co2) }
flatten_one (TyConApp tc tys)
-- Expand type synonyms that mention type families
-- on the RHS; see Note [Flattening synonyms]
| Just (tenv, rhs, tys') <- expandSynTyCon_maybe tc tys
, let expanded_ty = mkAppTys (substTy (mkTopTvSubst tenv) rhs) tys'
= do { mode <- getMode
; let used_tcs = tyConsOfType rhs
; case mode of
FM_FlattenAll | anyNameEnv isTypeFamilyTyCon used_tcs
-> flatten_one expanded_ty
_ -> flatten_ty_con_app tc tys }
-- Otherwise, it's a type function application, and we have to
-- flatten it away as well, and generate a new given equality constraint
-- between the application and a newly generated flattening skolem variable.
| isTypeFamilyTyCon tc
= flatten_fam_app tc tys
-- For * a normal data type application
-- * data family application
-- we just recursively flatten the arguments.
| otherwise
-- FM_Avoid stuff commented out; see Note [Lazy flattening]
-- , let fmode' = case fmode of -- Switch off the flat_top bit in FM_Avoid
-- FE { fe_mode = FM_Avoid tv _ }
-- -> fmode { fe_mode = FM_Avoid tv False }
-- _ -> fmode
= flatten_ty_con_app tc tys
flatten_one ty@(ForAllTy {})
-- We allow for-alls when, but only when, no type function
-- applications inside the forall involve the bound type variables.
= do { let (tvs, rho) = splitForAllTys ty
; (rho', co) <- setMode FM_SubstOnly $ flatten_one rho
-- Substitute only under a forall
-- See Note [Flattening under a forall]
; return (mkForAllTys tvs rho', foldr mkTcForAllCo co tvs) }
flatten_ty_con_app :: TyCon -> [TcType] -> FlatM (Xi, TcCoercion)
flatten_ty_con_app tc tys
= do { eq_rel <- getEqRel
; let role = eqRelRole eq_rel
; (xis, cos) <- case eq_rel of
NomEq -> flatten_many_nom tys
ReprEq -> flatten_many (tyConRolesX role tc) tys
; return (mkTyConApp tc xis, mkTcTyConAppCo role tc cos) }
{-
Note [Flattening synonyms]
~~~~~~~~~~~~~~~~~~~~~~~~~~
Not expanding synonyms aggressively improves error messages, and
keeps types smaller. But we need to take care.
Suppose
type T a = a -> a
and we want to flatten the type (T (F a)). Then we can safely flatten
the (F a) to a skolem, and return (T fsk). We don't need to expand the
synonym. This works because TcTyConAppCo can deal with synonyms
(unlike TyConAppCo), see Note [TcCoercions] in TcEvidence.
But (Trac #8979) for
type T a = (F a, a) where F is a type function
we must expand the synonym in (say) T Int, to expose the type function
to the flattener.
Note [Flattening under a forall]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Under a forall, we
(a) MUST apply the inert substitution
(b) MUST NOT flatten type family applications
Hence FMSubstOnly.
For (a) consider c ~ a, a ~ T (forall b. (b, [c]))
If we don't apply the c~a substitution to the second constraint
we won't see the occurs-check error.
For (b) consider (a ~ forall b. F a b), we don't want to flatten
to (a ~ forall b.fsk, F a b ~ fsk)
because now the 'b' has escaped its scope. We'd have to flatten to
(a ~ forall b. fsk b, forall b. F a b ~ fsk b)
and we have not begun to think about how to make that work!
************************************************************************
* *
Flattening a type-family application
* *
************************************************************************
-}
flatten_fam_app, flatten_exact_fam_app, flatten_exact_fam_app_fully
:: TyCon -> [TcType] -> FlatM (Xi, TcCoercion)
-- flatten_fam_app can be over-saturated
-- flatten_exact_fam_app is exactly saturated
-- flatten_exact_fam_app_fully lifts out the application to top level
-- Postcondition: Coercion :: Xi ~ F tys
flatten_fam_app tc tys -- Can be over-saturated
= ASSERT( tyConArity tc <= length tys ) -- Type functions are saturated
-- The type function might be *over* saturated
-- in which case the remaining arguments should
-- be dealt with by AppTys
do { let (tys1, tys_rest) = splitAt (tyConArity tc) tys
; (xi1, co1) <- flatten_exact_fam_app tc tys1
-- co1 :: xi1 ~ F tys1
-- all Nominal roles b/c the tycon is oversaturated
; (xis_rest, cos_rest) <- flatten_many (repeat Nominal) tys_rest
-- cos_res :: xis_rest ~ tys_rest
; return ( mkAppTys xi1 xis_rest -- NB mkAppTys: rhs_xi might not be a type variable
-- cf Trac #5655
, mkTcAppCos co1 cos_rest -- (rhs_xi :: F xis) ; (F cos :: F xis ~ F tys)
) }
flatten_exact_fam_app tc tys
= do { mode <- getMode
; role <- getRole
; case mode of
FM_FlattenAll -> flatten_exact_fam_app_fully tc tys
FM_SubstOnly -> do { (xis, cos) <- flatten_many roles tys
; return ( mkTyConApp tc xis
, mkTcTyConAppCo role tc cos ) }
where
-- These are always going to be Nominal for now,
-- but not if #8177 is implemented
roles = tyConRolesX role tc }
-- FM_Avoid tv flat_top ->
-- do { (xis, cos) <- flatten_many fmode roles tys
-- ; if flat_top || tv `elemVarSet` tyVarsOfTypes xis
-- then flatten_exact_fam_app_fully fmode tc tys
-- else return ( mkTyConApp tc xis
-- , mkTcTyConAppCo (feRole fmode) tc cos ) }
flatten_exact_fam_app_fully tc tys
-- See Note [Reduce type family applications eagerly]
= try_to_reduce tc tys False id $
do { -- First, flatten the arguments
(xis, cos) <- setEqRel NomEq $ flatten_many_nom tys
; eq_rel <- getEqRel
; let role = eqRelRole eq_rel
ret_co = mkTcTyConAppCo role tc cos
-- ret_co :: F xis ~ F tys
-- Now, look in the cache
; mb_ct <- liftTcS $ lookupFlatCache tc xis
; flavour_role <- getFlavourRole
; case mb_ct of
Just (co, rhs_ty, flav) -- co :: F xis ~ fsk
| (flav, NomEq) `canRewriteOrSameFR` flavour_role
-> -- Usable hit in the flat-cache
-- We certainly *can* use a Wanted for a Wanted
do { traceFlat "flatten/flat-cache hit" $ (ppr tc <+> ppr xis $$ ppr rhs_ty)
; (fsk_xi, fsk_co) <- flatten_one rhs_ty
-- The fsk may already have been unified, so flatten it
-- fsk_co :: fsk_xi ~ fsk
; return (fsk_xi, fsk_co `mkTcTransCo`
maybeTcSubCo eq_rel
(mkTcSymCo co) `mkTcTransCo`
ret_co) }
-- :: fsk_xi ~ F xis
-- Try to reduce the family application right now
-- See Note [Reduce type family applications eagerly]
_ -> try_to_reduce tc xis True (`mkTcTransCo` ret_co) $
do { let fam_ty = mkTyConApp tc xis
; (ev, fsk) <- newFlattenSkolemFlatM fam_ty
; let fsk_ty = mkTyVarTy fsk
co = ctEvCoercion ev
; liftTcS $ extendFlatCache tc xis (co, fsk_ty, ctEvFlavour ev)
-- The new constraint (F xis ~ fsk) is not necessarily inert
-- (e.g. the LHS may be a redex) so we must put it in the work list
; let ct = CFunEqCan { cc_ev = ev
, cc_fun = tc
, cc_tyargs = xis
, cc_fsk = fsk }
; emitFlatWork ct
; traceFlat "flatten/flat-cache miss" $ (ppr fam_ty $$ ppr fsk)
; return (fsk_ty, maybeTcSubCo eq_rel
(mkTcSymCo co)
`mkTcTransCo` ret_co) }
}
where
try_to_reduce :: TyCon -- F, family tycon
-> [Type] -- args, not necessarily flattened
-> Bool -- add to the flat cache?
-> ( TcCoercion -- :: xi ~ F args
-> TcCoercion ) -- what to return from outer function
-> FlatM (Xi, TcCoercion) -- continuation upon failure
-> FlatM (Xi, TcCoercion)
try_to_reduce tc tys cache update_co k
= do { checkStackDepth (mkTyConApp tc tys)
; mb_match <- liftTcS $ matchFam tc tys
; case mb_match of
Just (norm_co, norm_ty)
-> do { traceFlat "Eager T.F. reduction success" $
vcat [ppr tc, ppr tys, ppr norm_ty, ppr cache]
; (xi, final_co) <- bumpDepth $ flatten_one norm_ty
; let co = norm_co `mkTcTransCo` mkTcSymCo final_co
; flavour <- getFlavour
; when cache $
liftTcS $
extendFlatCache tc tys (co, xi, flavour)
; return (xi, update_co $ mkTcSymCo co) }
Nothing -> k }
{- Note [Reduce type family applications eagerly]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
If we come across a type-family application like (Append (Cons x Nil) t),
then, rather than flattening to a skolem etc, we may as well just reduce
it on the spot to (Cons x t). This saves a lot of intermediate steps.
Examples that are helped are tests T9872, and T5321Fun.
Performance testing indicates that it's best to try this *twice*, once
before flattening arguments and once after flattening arguments.
Adding the extra reduction attempt before flattening arguments cut
the allocation amounts for the T9872{a,b,c} tests by half.
An example of where the early reduction appears helpful:
type family Last x where
Last '[x] = x
Last (h ': t) = Last t
workitem: (x ~ Last '[1,2,3,4,5,6])
Flattening the argument never gets us anywhere, but trying to flatten
it at every step is quadratic in the length of the list. Reducing more
eagerly makes simplifying the right-hand type linear in its length.
Testing also indicated that the early reduction should *not* use the
flat-cache, but that the later reduction *should*. (Although the
effect was not large.) Hence the Bool argument to try_to_reduce. To
me (SLPJ) this seems odd; I get that eager reduction usually succeeds;
and if don't use the cache for eager reduction, we will miss most of
the opportunities for using it at all. More exploration would be good
here.
At the end, once we've got a flat rhs, we extend the flatten-cache to record
the result. Doing so can save lots of work when the same redex shows up more
than once. Note that we record the link from the redex all the way to its
*final* value, not just the single step reduction. Interestingly, using the
flat-cache for the first reduction resulted in an increase in allocations
of about 3% for the four T9872x tests. However, using the flat-cache in
the later reduction is a similar gain. I (Richard E) don't currently (Dec '14)
have any knowledge as to *why* these facts are true.
************************************************************************
* *
Flattening a type variable
* *
************************************************************************
Note [The inert equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Definition [Can-rewrite relation]
A "can-rewrite" relation between flavours, written f1 >= f2, is a
binary relation with the following properties
R1. >= is transitive
R2. If f1 >= f, and f2 >= f,
then either f1 >= f2 or f2 >= f1
Lemma. If f1 >= f then f1 >= f1
Proof. By property (R2), with f1=f2
Definition [Generalised substitution]
A "generalised substitution" S is a set of triples (a -f-> t), where
a is a type variable
t is a type
f is a flavour
such that
(WF1) if (a -f1-> t1) in S
(a -f2-> t2) in S
then neither (f1 >= f2) nor (f2 >= f1) hold
(WF2) if (a -f-> t) is in S, then t /= a
Definition [Applying a generalised substitution]
If S is a generalised substitution
S(f,a) = t, if (a -fs-> t) in S, and fs >= f
= a, otherwise
Application extends naturally to types S(f,t), modulo roles.
See Note [Flavours with roles].
Theorem: S(f,a) is well defined as a function.
Proof: Suppose (a -f1-> t1) and (a -f2-> t2) are both in S,
and f1 >= f and f2 >= f
Then by (R2) f1 >= f2 or f2 >= f1, which contradicts (WF)
Notation: repeated application.
S^0(f,t) = t
S^(n+1)(f,t) = S(f, S^n(t))
Definition: inert generalised substitution
A generalised substitution S is "inert" iff
(IG1) there is an n such that
for every f,t, S^n(f,t) = S^(n+1)(f,t)
(IG2) if (b -f-> t) in S, and f >= f, then S(f,t) = t
that is, each individual binding is "self-stable"
----------------------------------------------------------------
Our main invariant:
the inert CTyEqCans should be an inert generalised substitution
----------------------------------------------------------------
Note that inertness is not the same as idempotence. To apply S to a
type, you may have to apply it recursive. But inertness does
guarantee that this recursive use will terminate.
---------- The main theorem --------------
Suppose we have a "work item"
a -fw-> t
and an inert generalised substitution S,
such that
(T1) S(fw,a) = a -- LHS of work-item is a fixpoint of S(fw,_)
(T2) S(fw,t) = t -- RHS of work-item is a fixpoint of S(fw,_)
(T3) a not in t -- No occurs check in the work item
(K1) if (a -fs-> s) is in S then not (fw >= fs)
(K2) if (b -fs-> s) is in S, where b /= a, then
(K2a) not (fs >= fs)
or (K2b) not (fw >= fs)
or (K2c) a not in s
(K3) If (b -fs-> s) is in S with (fw >= fs), then
(K3a) If the role of fs is nominal: s /= a
(K3b) If the role of fs is representational: EITHER
a not in s, OR
the path from the top of s to a includes at least one non-newtype
then the extended substition T = S+(a -fw-> t)
is an inert generalised substitution.
The idea is that
* (T1-2) are guaranteed by exhaustively rewriting the work-item
with S(fw,_).
* T3 is guaranteed by a simple occurs-check on the work item.
* (K1-3) are the "kick-out" criteria. (As stated, they are really the
"keep" criteria.) If the current inert S contains a triple that does
not satisfy (K1-3), then we remove it from S by "kicking it out",
and re-processing it.
* Note that kicking out is a Bad Thing, because it means we have to
re-process a constraint. The less we kick out, the better.
TODO: Make sure that kicking out really *is* a Bad Thing. We've assumed
this but haven't done the empirical study to check.
* Assume we have G>=G, G>=W, D>=D, and that's all. Then, when performing
a unification we add a new given a -G-> ty. But doing so does NOT require
us to kick out an inert wanted that mentions a, because of (K2a). This
is a common case, hence good not to kick out.
* Lemma (L1): The conditions of the Main Theorem imply that there is no
(a fs-> t) in S, s.t. (fs >= fw).
Proof. Suppose the contrary (fs >= fw). Then because of (T1),
S(fw,a)=a. But since fs>=fw, S(fw,a) = s, hence s=a. But now we
have (a -fs-> a) in S, which contradicts (WF2).
* The extended substitution satisfies (WF1) and (WF2)
- (K1) plus (L1) guarantee that the extended substiution satisfies (WF1).
- (T3) guarantees (WF2).
* (K2) is about inertness. Intuitively, any infinite chain T^0(f,t),
T^1(f,t), T^2(f,T).... must pass through the new work item infnitely
often, since the substution without the work item is inert; and must
pass through at least one of the triples in S infnitely often.
- (K2a): if not(fs>=fs) then there is no f that fs can rewrite (fs>=f),
and hence this triple never plays a role in application S(f,a).
It is always safe to extend S with such a triple.
(NB: we could strengten K1) in this way too, but see K3.
- (K2b): If this holds, we can't pass through this triple infinitely
often, because if we did then fs>=f, fw>=f, hence fs>=fw,
contradicting (L1), or fw>=fs contradicting K2b.
- (K2c): if a not in s, we hae no further opportunity to apply the
work item.
NB: this reasoning isn't water tight.
Key lemma to make it watertight.
Under the conditions of the Main Theorem,
forall f st fw >= f, a is not in S^k(f,t), for any k
Also, consider roles more carefully. See Note [Flavours with roles].
Completeness
~~~~~~~~~~~~~
K3: completeness. (K3) is not necessary for the extended substitution
to be inert. In fact K1 could be made stronger by saying
... then (not (fw >= fs) or not (fs >= fs))
But it's not enough for S to be inert; we also want completeness.
That is, we want to be able to solve all soluble wanted equalities.
Suppose we have
work-item b -G-> a
inert-item a -W-> b
Assuming (G >= W) but not (W >= W), this fulfills all the conditions,
so we could extend the inerts, thus:
inert-items b -G-> a
a -W-> b
But if we kicked-out the inert item, we'd get
work-item a -W-> b
inert-item b -G-> a
Then rewrite the work-item gives us (a -W-> a), which is soluble via Refl.
So we add one more clause to the kick-out criteria
Another way to understand (K3) is that we treat an inert item
a -f-> b
in the same way as
b -f-> a
So if we kick out one, we should kick out the other. The orientation
is somewhat accidental.
When considering roles, we also need the second clause (K3b). Consider
inert-item a -W/R-> b c
work-item c -G/N-> a
The work-item doesn't get rewritten by the inert, because (>=) doesn't hold.
We've satisfied conditions (T1)-(T3) and (K1) and (K2). If all we had were
condition (K3a), then we would keep the inert around and add the work item.
But then, consider if we hit the following:
work-item2 b -G/N-> Id
where
newtype Id x = Id x
For similar reasons, if we only had (K3a), we wouldn't kick the
representational inert out. And then, we'd miss solving the inert, which
now reduced to reflexivity. The solution here is to kick out representational
inerts whenever the tyvar of a work item is "exposed", where exposed means
not under some proper data-type constructor, like [] or Maybe. See
isTyVarExposed in TcType. This is encoded in (K3b).
Note [Flavours with roles]
~~~~~~~~~~~~~~~~~~~~~~~~~~
The system described in Note [The inert equalities] discusses an abstract
set of flavours. In GHC, flavours have two components: the flavour proper,
taken from {Wanted, Derived, Given}; and the equality relation (often called
role), taken from {NomEq, ReprEq}. When substituting w.r.t. the inert set,
as described in Note [The inert equalities], we must be careful to respect
roles. For example, if we have
inert set: a -G/R-> Int
b -G/R-> Bool
type role T nominal representational
and we wish to compute S(W/R, T a b), the correct answer is T a Bool, NOT
T Int Bool. The reason is that T's first parameter has a nominal role, and
thus rewriting a to Int in T a b is wrong. Indeed, this non-congruence of
subsitution means that the proof in Note [The inert equalities] may need
to be revisited, but we don't think that the end conclusion is wrong.
-}
flatten_tyvar :: TcTyVar
-> FlatM (Either TyVar (TcType, TcCoercion))
-- "Flattening" a type variable means to apply the substitution to it
-- Specifically, look up the tyvar in
-- * the internal MetaTyVar box
-- * the inerts
-- Return (Left tv') if it is not found, tv' has a properly zonked kind
-- (Right (ty, co) if found, with co :: ty ~ tv;
flatten_tyvar tv
| not (isTcTyVar tv) -- Happens when flatten under a (forall a. ty)
= Left `liftM` flattenTyVarFinal tv
-- So ty contains refernces to the non-TcTyVar a
| otherwise
= do { mb_ty <- liftTcS $ isFilledMetaTyVar_maybe tv
; role <- getRole
; case mb_ty of {
Just ty -> do { traceFlat "Following filled tyvar" (ppr tv <+> equals <+> ppr ty)
; return (Right (ty, mkTcReflCo role ty)) } ;
Nothing ->
-- Try in the inert equalities
-- See Definition [Applying a generalised substitution]
do { ieqs <- liftTcS $ getInertEqs
; flavour_role <- getFlavourRole
; eq_rel <- getEqRel
; case lookupVarEnv ieqs tv of
Just (ct:_) -- If the first doesn't work,
-- the subsequent ones won't either
| CTyEqCan { cc_ev = ctev, cc_tyvar = tv, cc_rhs = rhs_ty } <- ct
, ctEvFlavourRole ctev `eqCanRewriteFR` flavour_role
-> do { traceFlat "Following inert tyvar" (ppr tv <+> equals <+> ppr rhs_ty $$ ppr ctev)
; let rewrite_co1 = mkTcSymCo (ctEvCoercion ctev)
rewrite_co = case (ctEvEqRel ctev, eq_rel) of
(ReprEq, _rel) -> ASSERT( _rel == ReprEq )
-- if this ASSERT fails, then
-- eqCanRewriteFR answered incorrectly
rewrite_co1
(NomEq, NomEq) -> rewrite_co1
(NomEq, ReprEq) -> mkTcSubCo rewrite_co1
; return (Right (rhs_ty, rewrite_co)) }
-- NB: ct is Derived then fmode must be also, hence
-- we are not going to touch the returned coercion
-- so ctEvCoercion is fine.
_other -> Left `liftM` flattenTyVarFinal tv
} } }
flattenTyVarFinal :: TcTyVar -> FlatM TyVar
flattenTyVarFinal tv
= -- Done, but make sure the kind is zonked
do { let kind = tyVarKind tv
; (new_knd, _kind_co) <- setMode FM_SubstOnly $ flatten_one kind
; return (setVarType tv new_knd) }
{-
Note [An alternative story for the inert substitution]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
(This entire note is just background, left here in case we ever want
to return the the previousl state of affairs)
We used (GHC 7.8) to have this story for the inert substitution inert_eqs
* 'a' is not in fvs(ty)
* They are *inert* in the weaker sense that there is no infinite chain of
(i1 `eqCanRewrite` i2), (i2 `eqCanRewrite` i3), etc
This means that flattening must be recursive, but it does allow
[G] a ~ [b]
[G] b ~ Maybe c
This avoids "saturating" the Givens, which can save a modest amount of work.
It is easy to implement, in TcInteract.kick_out, by only kicking out an inert
only if (a) the work item can rewrite the inert AND
(b) the inert cannot rewrite the work item
This is signifcantly harder to think about. It can save a LOT of work
in occurs-check cases, but we don't care about them much. Trac #5837
is an example; all the constraints here are Givens
[G] a ~ TF (a,Int)
-->
work TF (a,Int) ~ fsk
inert fsk ~ a
--->
work fsk ~ (TF a, TF Int)
inert fsk ~ a
--->
work a ~ (TF a, TF Int)
inert fsk ~ a
---> (attempting to flatten (TF a) so that it does not mention a
work TF a ~ fsk2
inert a ~ (fsk2, TF Int)
inert fsk ~ (fsk2, TF Int)
---> (substitute for a)
work TF (fsk2, TF Int) ~ fsk2
inert a ~ (fsk2, TF Int)
inert fsk ~ (fsk2, TF Int)
---> (top-level reduction, re-orient)
work fsk2 ~ (TF fsk2, TF Int)
inert a ~ (fsk2, TF Int)
inert fsk ~ (fsk2, TF Int)
---> (attempt to flatten (TF fsk2) to get rid of fsk2
work TF fsk2 ~ fsk3
work fsk2 ~ (fsk3, TF Int)
inert a ~ (fsk2, TF Int)
inert fsk ~ (fsk2, TF Int)
--->
work TF fsk2 ~ fsk3
inert fsk2 ~ (fsk3, TF Int)
inert a ~ ((fsk3, TF Int), TF Int)
inert fsk ~ ((fsk3, TF Int), TF Int)
Because the incoming given rewrites all the inert givens, we get more and
more duplication in the inert set. But this really only happens in pathalogical
casee, so we don't care.
-}
eqCanRewrite :: CtEvidence -> CtEvidence -> Bool
eqCanRewrite ev1 ev2 = ctEvFlavourRole ev1 `eqCanRewriteFR` ctEvFlavourRole ev2
-- | Whether or not one 'Ct' can rewrite another is determined by its
-- flavour and its equality relation
type CtFlavourRole = (CtFlavour, EqRel)
-- | Extract the flavour and role from a 'CtEvidence'
ctEvFlavourRole :: CtEvidence -> CtFlavourRole
ctEvFlavourRole ev = (ctEvFlavour ev, ctEvEqRel ev)
-- | Extract the flavour and role from a 'Ct'
ctFlavourRole :: Ct -> CtFlavourRole
ctFlavourRole = ctEvFlavourRole . cc_ev
eqCanRewriteFR :: CtFlavourRole -> CtFlavourRole -> Bool
-- Very important function!
-- See Note [eqCanRewrite]
-- See Note [Wanteds do not rewrite Wanteds]
-- See Note [Deriveds do rewrite Deriveds]
eqCanRewriteFR (Given, NomEq) (_, _) = True
eqCanRewriteFR (Given, ReprEq) (_, ReprEq) = True
eqCanRewriteFR (Derived, NomEq) (Derived, _) = True
eqCanRewriteFR _ _ = False
canRewriteOrSame :: CtEvidence -> CtEvidence -> Bool
-- See Note [canRewriteOrSame]
canRewriteOrSame ev1 ev2 = ev1 `eqCanRewrite` ev2 ||
ctEvFlavourRole ev1 == ctEvFlavourRole ev2
canRewriteOrSameFR :: CtFlavourRole -> CtFlavourRole -> Bool
canRewriteOrSameFR fr1 fr2 = fr1 `eqCanRewriteFR` fr2 || fr1 == fr2
{- Note [eqCanRewrite]
~~~~~~~~~~~~~~~~~~~
(eqCanRewrite ct1 ct2) holds if the constraint ct1 (a CTyEqCan of form
tv ~ ty) can be used to rewrite ct2. It must satisfy the properties of
a can-rewrite relation, see Definition [Can-rewrite relation]
With the solver handling Coercible constraints like equality constraints,
the rewrite conditions must take role into account, never allowing
a representational equality to rewrite a nominal one.
Note [Wanteds do not rewrite Wanteds]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We don't allow Wanteds to rewrite Wanteds, because that can give rise
to very confusing type error messages. A good example is Trac #8450.
Here's another
f :: a -> Bool
f x = ( [x,'c'], [x,True] ) `seq` True
Here we get
[W] a ~ Char
[W] a ~ Bool
but we do not want to complain about Bool ~ Char!
Note [Deriveds do rewrite Deriveds]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
However we DO allow Deriveds to rewrite Deriveds, because that's how
improvement works; see Note [The improvement story] in TcInteract.
Note [canRewriteOrSame]
~~~~~~~~~~~~~~~~~~~~~~~
canRewriteOrSame is similar but
* returns True for Wanted/Wanted.
* works for all kinds of constraints, not just CTyEqCans
See the call sites for explanations.
************************************************************************
* *
Unflattening
* *
************************************************************************
An unflattening example:
[W] F a ~ alpha
flattens to
[W] F a ~ fmv (CFunEqCan)
[W] fmv ~ alpha (CTyEqCan)
We must solve both!
-}
unflatten :: Cts -> Cts -> TcS Cts
unflatten tv_eqs funeqs
= do { dflags <- getDynFlags
; tclvl <- getTcLevel
; traceTcS "Unflattening" $ braces $
vcat [ ptext (sLit "Funeqs =") <+> pprCts funeqs
, ptext (sLit "Tv eqs =") <+> pprCts tv_eqs ]
-- Step 1: unflatten the CFunEqCans, except if that causes an occurs check
-- Occurs check: consider [W] alpha ~ [F alpha]
-- ==> (flatten) [W] F alpha ~ fmv, [W] alpha ~ [fmv]
-- ==> (unify) [W] F [fmv] ~ fmv
-- See Note [Unflatten using funeqs first]
; funeqs <- foldrBagM (unflatten_funeq dflags) emptyCts funeqs
; traceTcS "Unflattening 1" $ braces (pprCts funeqs)
-- Step 2: unify the irreds, if possible
; tv_eqs <- foldrBagM (unflatten_eq dflags tclvl) emptyCts tv_eqs
; traceTcS "Unflattening 2" $ braces (pprCts tv_eqs)
-- Step 3: fill any remaining fmvs with fresh unification variables
; funeqs <- mapBagM finalise_funeq funeqs
; traceTcS "Unflattening 3" $ braces (pprCts funeqs)
-- Step 4: remove any irreds that look like ty ~ ty
; tv_eqs <- foldrBagM finalise_eq emptyCts tv_eqs
; let all_flat = tv_eqs `andCts` funeqs
; traceTcS "Unflattening done" $ braces (pprCts all_flat)
-- Step 5: zonk the result
-- Motivation: makes them nice and ready for the next step
-- (see TcInteract.solveSimpleWanteds)
; zonkSimples all_flat }
where
----------------
unflatten_funeq :: DynFlags -> Ct -> Cts -> TcS Cts
unflatten_funeq dflags ct@(CFunEqCan { cc_fun = tc, cc_tyargs = xis
, cc_fsk = fmv, cc_ev = ev }) rest
= do { -- fmv should be a flatten meta-tv; we now fix its final
-- value, and then zonking will eliminate it
filled <- tryFill dflags fmv (mkTyConApp tc xis) ev
; return (if filled then rest else ct `consCts` rest) }
unflatten_funeq _ other_ct _
= pprPanic "unflatten_funeq" (ppr other_ct)
----------------
finalise_funeq :: Ct -> TcS Ct
finalise_funeq (CFunEqCan { cc_fsk = fmv, cc_ev = ev })
= do { demoteUnfilledFmv fmv
; return (mkNonCanonical ev) }
finalise_funeq ct = pprPanic "finalise_funeq" (ppr ct)
----------------
unflatten_eq :: DynFlags -> TcLevel -> Ct -> Cts -> TcS Cts
unflatten_eq dflags tclvl ct@(CTyEqCan { cc_ev = ev, cc_tyvar = tv, cc_rhs = rhs }) rest
| isFmvTyVar tv
= do { lhs_elim <- tryFill dflags tv rhs ev
; if lhs_elim then return rest else
do { rhs_elim <- try_fill dflags tclvl ev rhs (mkTyVarTy tv)
; if rhs_elim then return rest else
return (ct `consCts` rest) } }
| otherwise
= return (ct `consCts` rest)
unflatten_eq _ _ ct _ = pprPanic "unflatten_irred" (ppr ct)
----------------
finalise_eq :: Ct -> Cts -> TcS Cts
finalise_eq (CTyEqCan { cc_ev = ev, cc_tyvar = tv
, cc_rhs = rhs, cc_eq_rel = eq_rel }) rest
| isFmvTyVar tv
= do { ty1 <- zonkTcTyVar tv
; ty2 <- zonkTcType rhs
; let is_refl = ty1 `tcEqType` ty2
; if is_refl then do { setEvBindIfWanted ev
(EvCoercion $
mkTcReflCo (eqRelRole eq_rel) rhs)
; return rest }
else return (mkNonCanonical ev `consCts` rest) }
| otherwise
= return (mkNonCanonical ev `consCts` rest)
finalise_eq ct _ = pprPanic "finalise_irred" (ppr ct)
----------------
try_fill dflags tclvl ev ty1 ty2
| Just tv1 <- tcGetTyVar_maybe ty1
, isTouchableOrFmv tclvl tv1
, typeKind ty1 `isSubKind` tyVarKind tv1
= tryFill dflags tv1 ty2 ev
| otherwise
= return False
tryFill :: DynFlags -> TcTyVar -> TcType -> CtEvidence -> TcS Bool
-- (tryFill tv rhs ev) sees if 'tv' is an un-filled MetaTv
-- If so, and if tv does not appear in 'rhs', set tv := rhs
-- bind the evidence (which should be a CtWanted) to Refl
-- and return True. Otherwise return False
tryFill dflags tv rhs ev
= ASSERT2( not (isGiven ev), ppr ev )
do { is_filled <- isFilledMetaTyVar tv
; if is_filled then return False else
do { rhs' <- zonkTcType rhs
; case occurCheckExpand dflags tv rhs' of
OC_OK rhs'' -- Normal case: fill the tyvar
-> do { setEvBindIfWanted ev
(EvCoercion (mkTcReflCo (ctEvRole ev) rhs''))
; unifyTyVar tv rhs''
; return True }
_ -> -- Occurs check
return False } }
{-
Note [Unflatten using funeqs first]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
[W] G a ~ Int
[W] F (G a) ~ G a
do not want to end up with
[W] F Int ~ Int
because that might actually hold! Better to end up with the two above
unsolved constraints. The flat form will be
G a ~ fmv1 (CFunEqCan)
F fmv1 ~ fmv2 (CFunEqCan)
fmv1 ~ Int (CTyEqCan)
fmv1 ~ fmv2 (CTyEqCan)
Flatten using the fun-eqs first.
-}