{-# LANGUAGE CPP #-}
module TcCanonical(
canonicalize,
unifyDerived,
makeSuperClasses, maybeSym,
StopOrContinue(..), stopWith, continueWith,
solveCallStack -- For TcSimplify
) where
#include "HsVersions.h"
import GhcPrelude
import TcRnTypes
import TcUnify( swapOverTyVars, metaTyVarUpdateOK )
import TcType
import Type
import TcFlatten
import TcSMonad
import TcEvidence
import TcEvTerm
import Class
import TyCon
import TyCoRep -- cleverly decomposes types, good for completeness checking
import Coercion
import FamInstEnv ( FamInstEnvs )
import FamInst ( tcTopNormaliseNewTypeTF_maybe )
import Var
import VarEnv( mkInScopeSet )
import Outputable
import DynFlags( DynFlags )
import NameSet
import RdrName
import HsTypes( HsIPName(..) )
import Pair
import Util
import Bag
import MonadUtils
import Control.Monad
import Data.Maybe ( isJust )
import Data.List ( zip4, foldl' )
import BasicTypes
import Data.Bifunctor ( bimap )
{-
************************************************************************
* *
* The Canonicaliser *
* *
************************************************************************
Note [Canonicalization]
~~~~~~~~~~~~~~~~~~~~~~~
Canonicalization converts a simple constraint to a canonical form. It is
unary (i.e. treats individual constraints one at a time).
Constraints originating from user-written code come into being as
CNonCanonicals (except for CHoleCans, arising from holes). We know nothing
about these constraints. So, first:
Classify CNonCanoncal constraints, depending on whether they
are equalities, class predicates, or other.
Then proceed depending on the shape of the constraint. Generally speaking,
each constraint gets flattened and then decomposed into one of several forms
(see type Ct in TcRnTypes).
When an already-canonicalized constraint gets kicked out of the inert set,
it must be recanonicalized. But we know a bit about its shape from the
last time through, so we can skip the classification step.
-}
-- Top-level canonicalization
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
canonicalize :: Ct -> TcS (StopOrContinue Ct)
canonicalize (CNonCanonical { cc_ev = ev })
= {-# SCC "canNC" #-}
case classifyPredType (ctEvPred ev) of
ClassPred cls tys -> do traceTcS "canEvNC:cls" (ppr cls <+> ppr tys)
canClassNC ev cls tys
EqPred eq_rel ty1 ty2 -> do traceTcS "canEvNC:eq" (ppr ty1 $$ ppr ty2)
canEqNC ev eq_rel ty1 ty2
IrredPred {} -> do traceTcS "canEvNC:irred" (ppr (ctEvPred ev))
canIrred ev
canonicalize (CIrredCan { cc_ev = ev })
= canIrred ev
canonicalize (CDictCan { cc_ev = ev, cc_class = cls
, cc_tyargs = xis, cc_pend_sc = pend_sc })
= {-# SCC "canClass" #-}
canClass ev cls xis pend_sc
canonicalize (CTyEqCan { cc_ev = ev
, cc_tyvar = tv
, cc_rhs = xi
, cc_eq_rel = eq_rel })
= {-# SCC "canEqLeafTyVarEq" #-}
canEqNC ev eq_rel (mkTyVarTy tv) xi
-- NB: Don't use canEqTyVar because that expects flattened types,
-- and tv and xi may not be flat w.r.t. an updated inert set
canonicalize (CFunEqCan { cc_ev = ev
, cc_fun = fn
, cc_tyargs = xis1
, cc_fsk = fsk })
= {-# SCC "canEqLeafFunEq" #-}
canCFunEqCan ev fn xis1 fsk
canonicalize (CHoleCan { cc_ev = ev, cc_hole = hole })
= canHole ev hole
{-
************************************************************************
* *
* Class Canonicalization
* *
************************************************************************
-}
canClassNC :: CtEvidence -> Class -> [Type] -> TcS (StopOrContinue Ct)
-- "NC" means "non-canonical"; that is, we have got here
-- from a NonCanonical constraint, not from a CDictCan
-- Precondition: EvVar is class evidence
canClassNC ev cls tys
| isGiven ev -- See Note [Eagerly expand given superclasses]
= do { sc_cts <- mkStrictSuperClasses ev cls tys
; emitWork sc_cts
; canClass ev cls tys False }
| isWanted ev
, Just ip_name <- isCallStackPred cls tys
, OccurrenceOf func <- ctLocOrigin loc
-- If we're given a CallStack constraint that arose from a function
-- call, we need to push the current call-site onto the stack instead
-- of solving it directly from a given.
-- See Note [Overview of implicit CallStacks] in TcEvidence
-- and Note [Solving CallStack constraints] in TcSMonad
= do { -- First we emit a new constraint that will capture the
-- given CallStack.
; let new_loc = setCtLocOrigin loc (IPOccOrigin (HsIPName ip_name))
-- We change the origin to IPOccOrigin so
-- this rule does not fire again.
-- See Note [Overview of implicit CallStacks]
; new_ev <- newWantedEvVarNC new_loc pred
-- Then we solve the wanted by pushing the call-site
-- onto the newly emitted CallStack
; let ev_cs = EvCsPushCall func (ctLocSpan loc) (ctEvExpr new_ev)
; solveCallStack ev ev_cs
; canClass new_ev cls tys False }
| otherwise
= canClass ev cls tys (has_scs cls)
where
has_scs cls = not (null (classSCTheta cls))
loc = ctEvLoc ev
pred = ctEvPred ev
solveCallStack :: CtEvidence -> EvCallStack -> TcS ()
-- Also called from TcSimplify when defaulting call stacks
solveCallStack ev ev_cs = do
-- We're given ev_cs :: CallStack, but the evidence term should be a
-- dictionary, so we have to coerce ev_cs to a dictionary for
-- `IP ip CallStack`. See Note [Overview of implicit CallStacks]
cs_tm <- evCallStack ev_cs
let ev_tm = mkEvCast cs_tm (wrapIP (ctEvPred ev))
setWantedEvBind (ctEvEvId ev) (EvExpr ev_tm)
canClass :: CtEvidence
-> Class -> [Type]
-> Bool -- True <=> un-explored superclasses
-> TcS (StopOrContinue Ct)
-- Precondition: EvVar is class evidence
canClass ev cls tys pend_sc
= -- all classes do *nominal* matching
ASSERT2( ctEvRole ev == Nominal, ppr ev $$ ppr cls $$ ppr tys )
do { (xis, cos) <- flattenManyNom ev tys
; let co = mkTcTyConAppCo Nominal (classTyCon cls) cos
xi = mkClassPred cls xis
mk_ct new_ev = CDictCan { cc_ev = new_ev
, cc_tyargs = xis
, cc_class = cls
, cc_pend_sc = pend_sc }
; mb <- rewriteEvidence ev xi co
; traceTcS "canClass" (vcat [ ppr ev
, ppr xi, ppr mb ])
; return (fmap mk_ct mb) }
{- Note [The superclass story]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We need to add superclass constraints for two reasons:
* For givens [G], they give us a route to proof. E.g.
f :: Ord a => a -> Bool
f x = x == x
We get a Wanted (Eq a), which can only be solved from the superclass
of the Given (Ord a).
* For wanteds [W], and deriveds [WD], [D], they may give useful
functional dependencies. E.g.
class C a b | a -> b where ...
class C a b => D a b where ...
Now a [W] constraint (D Int beta) has (C Int beta) as a superclass
and that might tell us about beta, via C's fundeps. We can get this
by generating a [D] (C Int beta) constraint. It's derived because
we don't actually have to cough up any evidence for it; it's only there
to generate fundep equalities.
See Note [Why adding superclasses can help].
For these reasons we want to generate superclass constraints for both
Givens and Wanteds. But:
* (Minor) they are often not needed, so generating them aggressively
is a waste of time.
* (Major) if we want recursive superclasses, there would be an infinite
number of them. Here is a real-life example (Trac #10318);
class (Frac (Frac a) ~ Frac a,
Fractional (Frac a),
IntegralDomain (Frac a))
=> IntegralDomain a where
type Frac a :: *
Notice that IntegralDomain has an associated type Frac, and one
of IntegralDomain's superclasses is another IntegralDomain constraint.
So here's the plan:
1. Eagerly generate superclasses for given (but not wanted)
constraints; see Note [Eagerly expand given superclasses].
This is done in canClassNC, when we take a non-canonical constraint
and cannonicalise it.
However stop if you encounter the same class twice. That is,
expand eagerly, but have a conservative termination condition: see
Note [Expanding superclasses] in TcType.
2. Solve the wanteds as usual, but do no further expansion of
superclasses for canonical CDictCans in solveSimpleGivens or
solveSimpleWanteds; Note [Danger of adding superclasses during solving]
However, /do/ continue to eagerly expand superlasses for /given/
non-canonical constraints (canClassNC does this). As Trac #12175
showed, a type-family application can expand to a class constraint,
and we want to see its superclasses for just the same reason as
Note [Eagerly expand given superclasses].
3. If we have any remaining unsolved wanteds
(see Note [When superclasses help] in TcRnTypes)
try harder: take both the Givens and Wanteds, and expand
superclasses again. This may succeed in generating (a finite
number of) extra Givens, and extra Deriveds. Both may help the
proof. This is done in TcSimplify.expandSuperClasses.
4. Go round to (2) again. This loop (2,3,4) is implemented
in TcSimplify.simpl_loop.
The cc_pend_sc flag in a CDictCan records whether the superclasses of
this constraint have been expanded. Specifically, in Step 3 we only
expand superclasses for constraints with cc_pend_sc set to true (i.e.
isPendingScDict holds).
Why do we do this? Two reasons:
* To avoid repeated work, by repeatedly expanding the superclasses of
same constraint,
* To terminate the above loop, at least in the -XNoRecursiveSuperClasses
case. If there are recursive superclasses we could, in principle,
expand forever, always encountering new constraints.
When we take a CNonCanonical or CIrredCan, but end up classifying it
as a CDictCan, we set the cc_pend_sc flag to False.
Note [Eagerly expand given superclasses]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In step (1) of Note [The superclass story], why do we eagerly expand
Given superclasses by one layer? Mainly because of some very obscure
cases like this:
instance Bad a => Eq (T a)
f :: (Ord (T a)) => blah
f x = ....needs Eq (T a), Ord (T a)....
Here if we can't satisfy (Eq (T a)) from the givens we'll use the
instance declaration; but then we are stuck with (Bad a). Sigh.
This is really a case of non-confluent proofs, but to stop our users
complaining we expand one layer in advance.
Note [Instance and Given overlap] in TcInteract.
We also want to do this if we have
f :: F (T a) => blah
where
type instance F (T a) = Ord (T a)
So we may need to do a little work on the givens to expose the
class that has the superclasses. That's why the superclass
expansion for Givens happens in canClassNC.
Note [Why adding superclasses can help]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Examples of how adding superclasses can help:
--- Example 1
class C a b | a -> b
Suppose we want to solve
[G] C a b
[W] C a beta
Then adding [D] beta~b will let us solve it.
-- Example 2 (similar but using a type-equality superclass)
class (F a ~ b) => C a b
And try to sllve:
[G] C a b
[W] C a beta
Follow the superclass rules to add
[G] F a ~ b
[D] F a ~ beta
Now we get [D] beta ~ b, and can solve that.
-- Example (tcfail138)
class L a b | a -> b
class (G a, L a b) => C a b
instance C a b' => G (Maybe a)
instance C a b => C (Maybe a) a
instance L (Maybe a) a
When solving the superclasses of the (C (Maybe a) a) instance, we get
[G] C a b, and hance by superclasses, [G] G a, [G] L a b
[W] G (Maybe a)
Use the instance decl to get
[W] C a beta
Generate its derived superclass
[D] L a beta. Now using fundeps, combine with [G] L a b to get
[D] beta ~ b
which is what we want.
Note [Danger of adding superclasses during solving]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Here's a serious, but now out-dated example, from Trac #4497:
class Num (RealOf t) => Normed t
type family RealOf x
Assume the generated wanted constraint is:
[W] RealOf e ~ e
[W] Normed e
If we were to be adding the superclasses during simplification we'd get:
[W] RealOf e ~ e
[W] Normed e
[D] RealOf e ~ fuv
[D] Num fuv
==>
e := fuv, Num fuv, Normed fuv, RealOf fuv ~ fuv
While looks exactly like our original constraint. If we add the
superclass of (Normed fuv) again we'd loop. By adding superclasses
definitely only once, during canonicalisation, this situation can't
happen.
Mind you, now that Wanteds cannot rewrite Derived, I think this particular
situation can't happen.
-}
makeSuperClasses :: [Ct] -> TcS [Ct]
-- Returns strict superclasses, transitively, see Note [The superclasses story]
-- See Note [The superclass story]
-- The loop-breaking here follows Note [Expanding superclasses] in TcType
-- Specifically, for an incoming (C t) constraint, we return all of (C t)'s
-- superclasses, up to /and including/ the first repetition of C
--
-- Example: class D a => C a
-- class C [a] => D a
-- makeSuperClasses (C x) will return (D x, C [x])
--
-- NB: the incoming constraints have had their cc_pend_sc flag already
-- flipped to False, by isPendingScDict, so we are /obliged/ to at
-- least produce the immediate superclasses
makeSuperClasses cts = concatMapM go cts
where
go (CDictCan { cc_ev = ev, cc_class = cls, cc_tyargs = tys })
= mkStrictSuperClasses ev cls tys
go ct = pprPanic "makeSuperClasses" (ppr ct)
mkStrictSuperClasses :: CtEvidence -> Class -> [Type] -> TcS [Ct]
-- Return constraints for the strict superclasses of (c tys)
mkStrictSuperClasses ev cls tys
= mk_strict_superclasses (unitNameSet (className cls)) ev cls tys
mk_superclasses :: NameSet -> CtEvidence -> TcS [Ct]
-- Return this constraint, plus its superclasses, if any
mk_superclasses rec_clss ev
| ClassPred cls tys <- classifyPredType (ctEvPred ev)
= mk_superclasses_of rec_clss ev cls tys
| otherwise -- Superclass is not a class predicate
= return [mkNonCanonical ev]
mk_superclasses_of :: NameSet -> CtEvidence -> Class -> [Type] -> TcS [Ct]
-- Always return this class constraint,
-- and expand its superclasses
mk_superclasses_of rec_clss ev cls tys
| loop_found = do { traceTcS "mk_superclasses_of: loop" (ppr cls <+> ppr tys)
; return [this_ct] } -- cc_pend_sc of this_ct = True
| otherwise = do { traceTcS "mk_superclasses_of" (vcat [ ppr cls <+> ppr tys
, ppr (isCTupleClass cls)
, ppr rec_clss
])
; sc_cts <- mk_strict_superclasses rec_clss' ev cls tys
; return (this_ct : sc_cts) }
-- cc_pend_sc of this_ct = False
where
cls_nm = className cls
loop_found = not (isCTupleClass cls) && cls_nm `elemNameSet` rec_clss
-- Tuples never contribute to recursion, and can be nested
rec_clss' = rec_clss `extendNameSet` cls_nm
this_ct = CDictCan { cc_ev = ev, cc_class = cls, cc_tyargs = tys
, cc_pend_sc = loop_found }
-- NB: If there is a loop, we cut off, so we have not
-- added the superclasses, hence cc_pend_sc = True
mk_strict_superclasses :: NameSet -> CtEvidence -> Class -> [Type] -> TcS [Ct]
-- Always return the immediate superclasses of (cls tys);
-- and expand their superclasses, provided none of them are in rec_clss
-- nor are repeated
mk_strict_superclasses rec_clss ev cls tys
| CtGiven { ctev_evar = evar, ctev_loc = loc } <- ev
= do { sc_evs <- newGivenEvVars (mk_given_loc loc)
(mkEvScSelectors (evId evar) cls tys)
; concatMapM (mk_superclasses rec_clss) sc_evs }
| all noFreeVarsOfType tys
= return [] -- Wanteds with no variables yield no deriveds.
-- See Note [Improvement from Ground Wanteds]
| otherwise -- Wanted/Derived case, just add Derived superclasses
-- that can lead to improvement.
= do { let loc = ctEvLoc ev
; sc_evs <- mapM (newDerivedNC loc) (immSuperClasses cls tys)
; concatMapM (mk_superclasses rec_clss) sc_evs }
where
size = sizeTypes tys
mk_given_loc loc
| isCTupleClass cls
= loc -- For tuple predicates, just take them apart, without
-- adding their (large) size into the chain. When we
-- get down to a base predicate, we'll include its size.
-- Trac #10335
| GivenOrigin skol_info <- ctLocOrigin loc
-- See Note [Solving superclass constraints] in TcInstDcls
-- for explantation of this transformation for givens
= case skol_info of
InstSkol -> loc { ctl_origin = GivenOrigin (InstSC size) }
InstSC n -> loc { ctl_origin = GivenOrigin (InstSC (n `max` size)) }
_ -> loc
| otherwise -- Probably doesn't happen, since this function
= loc -- is only used for Givens, but does no harm
{-
************************************************************************
* *
* Irreducibles canonicalization
* *
************************************************************************
-}
canIrred :: CtEvidence -> TcS (StopOrContinue Ct)
-- Precondition: ty not a tuple and no other evidence form
canIrred ev
| EqPred eq_rel ty1 ty2 <- classifyPredType pred
= -- For insolubles (all of which are equalities, do /not/ flatten the arguments
-- In Trac #14350 doing so led entire-unnecessary and ridiculously large
-- type function expansion. Instead, canEqNC just applies
-- the substitution to the predicate, and may do decomposition;
-- e.g. a ~ [a], where [G] a ~ [Int], can decompose
canEqNC ev eq_rel ty1 ty2
| otherwise
= do { traceTcS "can_pred" (text "IrredPred = " <+> ppr pred)
; (xi,co) <- flatten FM_FlattenAll ev pred -- co :: xi ~ pred
; rewriteEvidence ev xi co `andWhenContinue` \ new_ev ->
do { -- Re-classify, in case flattening has improved its shape
; case classifyPredType (ctEvPred new_ev) of
ClassPred cls tys -> canClassNC new_ev cls tys
EqPred eq_rel ty1 ty2 -> canEqNC new_ev eq_rel ty1 ty2
_ -> continueWith $
mkIrredCt new_ev } }
where
pred = ctEvPred ev
canHole :: CtEvidence -> Hole -> TcS (StopOrContinue Ct)
canHole ev hole
= do { let ty = ctEvPred ev
; (xi,co) <- flatten FM_SubstOnly ev ty -- co :: xi ~ ty
; rewriteEvidence ev xi co `andWhenContinue` \ new_ev ->
do { updInertIrreds (`snocCts` (CHoleCan { cc_ev = new_ev
, cc_hole = hole }))
; stopWith new_ev "Emit insoluble hole" } }
{-
************************************************************************
* *
* Equalities
* *
************************************************************************
Note [Canonicalising equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In order to canonicalise an equality, we look at the structure of the
two types at hand, looking for similarities. A difficulty is that the
types may look dissimilar before flattening but similar after flattening.
However, we don't just want to jump in and flatten right away, because
this might be wasted effort. So, after looking for similarities and failing,
we flatten and then try again. Of course, we don't want to loop, so we
track whether or not we've already flattened.
It is conceivable to do a better job at tracking whether or not a type
is flattened, but this is left as future work. (Mar '15)
Note [FunTy and decomposing tycon applications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When can_eq_nc' attempts to decompose a tycon application we haven't yet zonked.
This means that we may very well have a FunTy containing a type of some unknown
kind. For instance, we may have,
FunTy (a :: k) Int
Where k is a unification variable. tcRepSplitTyConApp_maybe panics in the event
that it sees such a type as it cannot determine the RuntimeReps which the (->)
is applied to. Consequently, it is vital that we instead use
tcRepSplitTyConApp_maybe', which simply returns Nothing in such a case.
When this happens can_eq_nc' will fail to decompose, zonk, and try again.
Zonking should fill the variable k, meaning that decomposition will succeed the
second time around.
-}
canEqNC :: CtEvidence -> EqRel -> Type -> Type -> TcS (StopOrContinue Ct)
canEqNC ev eq_rel ty1 ty2
= do { result <- zonk_eq_types ty1 ty2
; case result of
Left (Pair ty1' ty2') -> can_eq_nc False ev eq_rel ty1' ty1 ty2' ty2
Right ty -> canEqReflexive ev eq_rel ty }
can_eq_nc
:: Bool -- True => both types are flat
-> CtEvidence
-> EqRel
-> Type -> Type -- LHS, after and before type-synonym expansion, resp
-> Type -> Type -- RHS, after and before type-synonym expansion, resp
-> TcS (StopOrContinue Ct)
can_eq_nc flat ev eq_rel ty1 ps_ty1 ty2 ps_ty2
= do { traceTcS "can_eq_nc" $
vcat [ ppr flat, ppr ev, ppr eq_rel, ppr ty1, ppr ps_ty1, ppr ty2, ppr ps_ty2 ]
; rdr_env <- getGlobalRdrEnvTcS
; fam_insts <- getFamInstEnvs
; can_eq_nc' flat rdr_env fam_insts ev eq_rel ty1 ps_ty1 ty2 ps_ty2 }
can_eq_nc'
:: Bool -- True => both input types are flattened
-> GlobalRdrEnv -- needed to see which newtypes are in scope
-> FamInstEnvs -- needed to unwrap data instances
-> CtEvidence
-> EqRel
-> Type -> Type -- LHS, after and before type-synonym expansion, resp
-> Type -> Type -- RHS, after and before type-synonym expansion, resp
-> TcS (StopOrContinue Ct)
-- Expand synonyms first; see Note [Type synonyms and canonicalization]
can_eq_nc' flat _rdr_env _envs ev eq_rel ty1 ps_ty1 ty2 ps_ty2
| Just ty1' <- tcView ty1 = can_eq_nc flat ev eq_rel ty1' ps_ty1 ty2 ps_ty2
| Just ty2' <- tcView ty2 = can_eq_nc flat ev eq_rel ty1 ps_ty1 ty2' ps_ty2
-- need to check for reflexivity in the ReprEq case.
-- See Note [Eager reflexivity check]
-- Check only when flat because the zonk_eq_types check in canEqNC takes
-- care of the non-flat case.
can_eq_nc' True _rdr_env _envs ev ReprEq ty1 _ ty2 _
| ty1 `tcEqType` ty2
= canEqReflexive ev ReprEq ty1
-- When working with ReprEq, unwrap newtypes.
can_eq_nc' _flat rdr_env envs ev ReprEq ty1 _ ty2 ps_ty2
| Just stuff1 <- tcTopNormaliseNewTypeTF_maybe envs rdr_env ty1
= can_eq_newtype_nc ev NotSwapped ty1 stuff1 ty2 ps_ty2
can_eq_nc' _flat rdr_env envs ev ReprEq ty1 ps_ty1 ty2 _
| Just stuff2 <- tcTopNormaliseNewTypeTF_maybe envs rdr_env ty2
= can_eq_newtype_nc ev IsSwapped ty2 stuff2 ty1 ps_ty1
-- Now, check for tyvars. This must happen before CastTy because we need
-- to catch casted tyvars, as the flattener might produce these,
-- due to the fact that flattened types have flattened kinds.
-- See Note [Flattening].
-- Note that there can be only one cast on the tyvar because this will
-- run after the "get rid of casts" case of can_eq_nc' function on the
-- not-yet-flattened types.
-- NB: pattern match on True: we want only flat types sent to canEqTyVar.
-- See also Note [No top-level newtypes on RHS of representational equalities]
can_eq_nc' True _rdr_env _envs ev eq_rel ty1 ps_ty1 ty2 ps_ty2
| Just (tv1, co1) <- getCastedTyVar_maybe ty1
= canEqTyVar ev eq_rel NotSwapped tv1 co1 ps_ty1 ty2 ps_ty2
can_eq_nc' True _rdr_env _envs ev eq_rel ty1 ps_ty1 ty2 ps_ty2
| Just (tv2, co2) <- getCastedTyVar_maybe ty2
= canEqTyVar ev eq_rel IsSwapped tv2 co2 ps_ty2 ty1 ps_ty1
-- Then, get rid of casts
can_eq_nc' flat _rdr_env _envs ev eq_rel (CastTy ty1 co1) _ ty2 ps_ty2
= canEqCast flat ev eq_rel NotSwapped ty1 co1 ty2 ps_ty2
can_eq_nc' flat _rdr_env _envs ev eq_rel ty1 ps_ty1 (CastTy ty2 co2) _
= canEqCast flat ev eq_rel IsSwapped ty2 co2 ty1 ps_ty1
----------------------
-- Otherwise try to decompose
----------------------
-- Literals
can_eq_nc' _flat _rdr_env _envs ev eq_rel ty1@(LitTy l1) _ (LitTy l2) _
| l1 == l2
= do { setEqIfWanted ev (mkReflCo (eqRelRole eq_rel) ty1)
; stopWith ev "Equal LitTy" }
-- Try to decompose type constructor applications
-- Including FunTy (s -> t)
can_eq_nc' _flat _rdr_env _envs ev eq_rel ty1 _ ty2 _
--- See Note [FunTy and decomposing type constructor applications].
| Just (tc1, tys1) <- tcRepSplitTyConApp_maybe' ty1
, Just (tc2, tys2) <- tcRepSplitTyConApp_maybe' ty2
, not (isTypeFamilyTyCon tc1)
, not (isTypeFamilyTyCon tc2)
= canTyConApp ev eq_rel tc1 tys1 tc2 tys2
can_eq_nc' _flat _rdr_env _envs ev eq_rel
s1@(ForAllTy {}) _ s2@(ForAllTy {}) _
= can_eq_nc_forall ev eq_rel s1 s2
-- See Note [Canonicalising type applications] about why we require flat types
can_eq_nc' True _rdr_env _envs ev eq_rel (AppTy t1 s1) _ ty2 _
| Just (t2, s2) <- tcSplitAppTy_maybe ty2
= can_eq_app ev eq_rel t1 s1 t2 s2
can_eq_nc' True _rdr_env _envs ev eq_rel ty1 _ (AppTy t2 s2) _
| Just (t1, s1) <- tcSplitAppTy_maybe ty1
= can_eq_app ev eq_rel t1 s1 t2 s2
-- No similarity in type structure detected. Flatten and try again.
can_eq_nc' False rdr_env envs ev eq_rel _ ps_ty1 _ ps_ty2
= do { (xi1, co1) <- flatten FM_FlattenAll ev ps_ty1
; (xi2, co2) <- flatten FM_FlattenAll ev ps_ty2
; rewriteEqEvidence ev NotSwapped xi1 xi2 co1 co2
`andWhenContinue` \ new_ev ->
can_eq_nc' True rdr_env envs new_ev eq_rel xi1 xi1 xi2 xi2 }
-- We've flattened and the types don't match. Give up.
can_eq_nc' True _rdr_env _envs ev _eq_rel _ ps_ty1 _ ps_ty2
= do { traceTcS "can_eq_nc' catch-all case" (ppr ps_ty1 $$ ppr ps_ty2)
; canEqHardFailure ev ps_ty1 ps_ty2 }
---------------------------------
can_eq_nc_forall :: CtEvidence -> EqRel
-> Type -> Type -- LHS and RHS
-> TcS (StopOrContinue Ct)
-- (forall as. phi1) ~ (forall bs. phi2)
-- Check for length match of as, bs
-- Then build an implication constraint: forall as. phi1 ~ phi2[as/bs]
-- But remember also to unify the kinds of as and bs
-- (this is the 'go' loop), and actually substitute phi2[as |> cos / bs]
-- Remember also that we might have forall z (a:z). blah
-- so we must proceed one binder at a time (Trac #13879)
can_eq_nc_forall ev eq_rel s1 s2
| CtWanted { ctev_loc = loc, ctev_dest = orig_dest } <- ev
= do { let free_tvs = tyCoVarsOfTypes [s1,s2]
(bndrs1, phi1) = tcSplitForAllTyVarBndrs s1
(bndrs2, phi2) = tcSplitForAllTyVarBndrs s2
; if not (equalLength bndrs1 bndrs2)
then do { traceTcS "Forall failure" $
vcat [ ppr s1, ppr s2, ppr bndrs1, ppr bndrs2
, ppr (map binderArgFlag bndrs1)
, ppr (map binderArgFlag bndrs2) ]
; canEqHardFailure ev s1 s2 }
else
do { traceTcS "Creating implication for polytype equality" $ ppr ev
; let empty_subst1 = mkEmptyTCvSubst $ mkInScopeSet free_tvs
; (subst1, skol_tvs) <- tcInstSkolTyVarsX empty_subst1 $
binderVars bndrs1
; let skol_info = UnifyForAllSkol phi1
phi1' = substTy subst1 phi1
-- Unify the kinds, extend the substitution
go (skol_tv:skol_tvs) subst (bndr2:bndrs2)
= do { let tv2 = binderVar bndr2
; kind_co <- unifyWanted loc Nominal
(tyVarKind skol_tv)
(substTy subst (tyVarKind tv2))
; let subst' = extendTvSubst subst tv2
(mkCastTy (mkTyVarTy skol_tv) kind_co)
; co <- go skol_tvs subst' bndrs2
; return (mkForAllCo skol_tv kind_co co) }
-- Done: unify phi1 ~ phi2
go [] subst bndrs2
= ASSERT( null bndrs2 )
unifyWanted loc (eqRelRole eq_rel)
phi1' (substTy subst phi2)
go _ _ _ = panic "cna_eq_nc_forall" -- case (s:ss) []
empty_subst2 = mkEmptyTCvSubst (getTCvInScope subst1)
; (implic, _ev_binds, all_co) <- buildImplication skol_info skol_tvs [] $
go skol_tvs empty_subst2 bndrs2
-- We have nowhere to put these bindings
-- but TcSimplify.setImplicationStatus
-- checks that we don't actually use them
-- when skol_info = UnifyForAllSkol
; updWorkListTcS (extendWorkListImplic implic)
; setWantedEq orig_dest all_co
; stopWith ev "Deferred polytype equality" } }
| otherwise
= do { traceTcS "Omitting decomposition of given polytype equality" $
pprEq s1 s2 -- See Note [Do not decompose given polytype equalities]
; stopWith ev "Discard given polytype equality" }
---------------------------------
-- | Compare types for equality, while zonking as necessary. Gives up
-- as soon as it finds that two types are not equal.
-- This is quite handy when some unification has made two
-- types in an inert wanted to be equal. We can discover the equality without
-- flattening, which is sometimes very expensive (in the case of type functions).
-- In particular, this function makes a ~20% improvement in test case
-- perf/compiler/T5030.
--
-- Returns either the (partially zonked) types in the case of
-- inequality, or the one type in the case of equality. canEqReflexive is
-- a good next step in the 'Right' case. Returning 'Left' is always safe.
--
-- NB: This does *not* look through type synonyms. In fact, it treats type
-- synonyms as rigid constructors. In the future, it might be convenient
-- to look at only those arguments of type synonyms that actually appear
-- in the synonym RHS. But we're not there yet.
zonk_eq_types :: TcType -> TcType -> TcS (Either (Pair TcType) TcType)
zonk_eq_types = go
where
go (TyVarTy tv1) (TyVarTy tv2) = tyvar_tyvar tv1 tv2
go (TyVarTy tv1) ty2 = tyvar NotSwapped tv1 ty2
go ty1 (TyVarTy tv2) = tyvar IsSwapped tv2 ty1
-- We handle FunTys explicitly here despite the fact that they could also be
-- treated as an application. Why? Well, for one it's cheaper to just look
-- at two types (the argument and result types) than four (the argument,
-- result, and their RuntimeReps). Also, we haven't completely zonked yet,
-- so we may run into an unzonked type variable while trying to compute the
-- RuntimeReps of the argument and result types. This can be observed in
-- testcase tc269.
go ty1 ty2
| Just (arg1, res1) <- split1
, Just (arg2, res2) <- split2
= do { res_a <- go arg1 arg2
; res_b <- go res1 res2
; return $ combine_rev mkFunTy res_b res_a
}
| isJust split1 || isJust split2
= bale_out ty1 ty2
where
split1 = tcSplitFunTy_maybe ty1
split2 = tcSplitFunTy_maybe ty2
go ty1 ty2
| Just (tc1, tys1) <- tcRepSplitTyConApp_maybe ty1
, Just (tc2, tys2) <- tcRepSplitTyConApp_maybe ty2
= if tc1 == tc2 && tys1 `equalLength` tys2
-- Crucial to check for equal-length args, because
-- we cannot assume that the two args to 'go' have
-- the same kind. E.g go (Proxy * (Maybe Int))
-- (Proxy (*->*) Maybe)
-- We'll call (go (Maybe Int) Maybe)
-- See Trac #13083
then tycon tc1 tys1 tys2
else bale_out ty1 ty2
go ty1 ty2
| Just (ty1a, ty1b) <- tcRepSplitAppTy_maybe ty1
, Just (ty2a, ty2b) <- tcRepSplitAppTy_maybe ty2
= do { res_a <- go ty1a ty2a
; res_b <- go ty1b ty2b
; return $ combine_rev mkAppTy res_b res_a }
go ty1@(LitTy lit1) (LitTy lit2)
| lit1 == lit2
= return (Right ty1)
go ty1 ty2 = bale_out ty1 ty2
-- We don't handle more complex forms here
bale_out ty1 ty2 = return $ Left (Pair ty1 ty2)
tyvar :: SwapFlag -> TcTyVar -> TcType
-> TcS (Either (Pair TcType) TcType)
-- Try to do as little as possible, as anything we do here is redundant
-- with flattening. In particular, no need to zonk kinds. That's why
-- we don't use the already-defined zonking functions
tyvar swapped tv ty
= case tcTyVarDetails tv of
MetaTv { mtv_ref = ref }
-> do { cts <- readTcRef ref
; case cts of
Flexi -> give_up
Indirect ty' -> do { trace_indirect tv ty'
; unSwap swapped go ty' ty } }
_ -> give_up
where
give_up = return $ Left $ unSwap swapped Pair (mkTyVarTy tv) ty
tyvar_tyvar tv1 tv2
| tv1 == tv2 = return (Right (mkTyVarTy tv1))
| otherwise = do { (ty1', progress1) <- quick_zonk tv1
; (ty2', progress2) <- quick_zonk tv2
; if progress1 || progress2
then go ty1' ty2'
else return $ Left (Pair (TyVarTy tv1) (TyVarTy tv2)) }
trace_indirect tv ty
= traceTcS "Following filled tyvar (zonk_eq_types)"
(ppr tv <+> equals <+> ppr ty)
quick_zonk tv = case tcTyVarDetails tv of
MetaTv { mtv_ref = ref }
-> do { cts <- readTcRef ref
; case cts of
Flexi -> return (TyVarTy tv, False)
Indirect ty' -> do { trace_indirect tv ty'
; return (ty', True) } }
_ -> return (TyVarTy tv, False)
-- This happens for type families, too. But recall that failure
-- here just means to try harder, so it's OK if the type function
-- isn't injective.
tycon :: TyCon -> [TcType] -> [TcType]
-> TcS (Either (Pair TcType) TcType)
tycon tc tys1 tys2
= do { results <- zipWithM go tys1 tys2
; return $ case combine_results results of
Left tys -> Left (mkTyConApp tc <$> tys)
Right tys -> Right (mkTyConApp tc tys) }
combine_results :: [Either (Pair TcType) TcType]
-> Either (Pair [TcType]) [TcType]
combine_results = bimap (fmap reverse) reverse .
foldl' (combine_rev (:)) (Right [])
-- combine (in reverse) a new result onto an already-combined result
combine_rev :: (a -> b -> c)
-> Either (Pair b) b
-> Either (Pair a) a
-> Either (Pair c) c
combine_rev f (Left list) (Left elt) = Left (f <$> elt <*> list)
combine_rev f (Left list) (Right ty) = Left (f <$> pure ty <*> list)
combine_rev f (Right tys) (Left elt) = Left (f <$> elt <*> pure tys)
combine_rev f (Right tys) (Right ty) = Right (f ty tys)
{-
Note [Newtypes can blow the stack]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have
newtype X = MkX (Int -> X)
newtype Y = MkY (Int -> Y)
and now wish to prove
[W] X ~R Y
This Wanted will loop, expanding out the newtypes ever deeper looking
for a solid match or a solid discrepancy. Indeed, there is something
appropriate to this looping, because X and Y *do* have the same representation,
in the limit -- they're both (Fix ((->) Int)). However, no finitely-sized
coercion will ever witness it. This loop won't actually cause GHC to hang,
though, because we check our depth when unwrapping newtypes.
Note [Eager reflexivity check]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have
newtype X = MkX (Int -> X)
and
[W] X ~R X
Naively, we would start unwrapping X and end up in a loop. Instead,
we do this eager reflexivity check. This is necessary only for representational
equality because the flattener technology deals with the similar case
(recursive type families) for nominal equality.
Note that this check does not catch all cases, but it will catch the cases
we're most worried about, types like X above that are actually inhabited.
Here's another place where this reflexivity check is key:
Consider trying to prove (f a) ~R (f a). The AppTys in there can't
be decomposed, because representational equality isn't congruent with respect
to AppTy. So, when canonicalising the equality above, we get stuck and
would normally produce a CIrredCan. However, we really do want to
be able to solve (f a) ~R (f a). So, in the representational case only,
we do a reflexivity check.
(This would be sound in the nominal case, but unnecessary, and I [Richard
E.] am worried that it would slow down the common case.)
-}
------------------------
-- | We're able to unwrap a newtype. Update the bits accordingly.
can_eq_newtype_nc :: CtEvidence -- ^ :: ty1 ~ ty2
-> SwapFlag
-> TcType -- ^ ty1
-> ((Bag GlobalRdrElt, TcCoercion), TcType) -- ^ :: ty1 ~ ty1'
-> TcType -- ^ ty2
-> TcType -- ^ ty2, with type synonyms
-> TcS (StopOrContinue Ct)
can_eq_newtype_nc ev swapped ty1 ((gres, co), ty1') ty2 ps_ty2
= do { traceTcS "can_eq_newtype_nc" $
vcat [ ppr ev, ppr swapped, ppr co, ppr gres, ppr ty1', ppr ty2 ]
-- check for blowing our stack:
-- See Note [Newtypes can blow the stack]
; checkReductionDepth (ctEvLoc ev) ty1
; addUsedGREs (bagToList gres)
-- we have actually used the newtype constructor here, so
-- make sure we don't warn about importing it!
; rewriteEqEvidence ev swapped ty1' ps_ty2
(mkTcSymCo co) (mkTcReflCo Representational ps_ty2)
`andWhenContinue` \ new_ev ->
can_eq_nc False new_ev ReprEq ty1' ty1' ty2 ps_ty2 }
---------
-- ^ Decompose a type application.
-- All input types must be flat. See Note [Canonicalising type applications]
can_eq_app :: CtEvidence -- :: s1 t1 ~r s2 t2
-> EqRel -- r
-> Xi -> Xi -- s1 t1
-> Xi -> Xi -- s2 t2
-> TcS (StopOrContinue Ct)
-- AppTys only decompose for nominal equality, so this case just leads
-- to an irreducible constraint; see typecheck/should_compile/T10494
-- See Note [Decomposing equality], note {4}
can_eq_app ev ReprEq _ _ _ _
= do { traceTcS "failing to decompose representational AppTy equality" (ppr ev)
; continueWith (mkIrredCt ev) }
-- no need to call canEqFailure, because that flattens, and the
-- types involved here are already flat
can_eq_app ev NomEq s1 t1 s2 t2
| CtDerived { ctev_loc = loc } <- ev
= do { unifyDeriveds loc [Nominal, Nominal] [s1, t1] [s2, t2]
; stopWith ev "Decomposed [D] AppTy" }
| CtWanted { ctev_dest = dest, ctev_loc = loc } <- ev
= do { co_s <- unifyWanted loc Nominal s1 s2
; let arg_loc
| isNextArgVisible s1 = loc
| otherwise = updateCtLocOrigin loc toInvisibleOrigin
; co_t <- unifyWanted arg_loc Nominal t1 t2
; let co = mkAppCo co_s co_t
; setWantedEq dest co
; stopWith ev "Decomposed [W] AppTy" }
| CtGiven { ctev_evar = evar, ctev_loc = loc } <- ev
= do { let co = mkTcCoVarCo evar
co_s = mkTcLRCo CLeft co
co_t = mkTcLRCo CRight co
; evar_s <- newGivenEvVar loc ( mkTcEqPredLikeEv ev s1 s2
, evCoercion co_s )
; evar_t <- newGivenEvVar loc ( mkTcEqPredLikeEv ev t1 t2
, evCoercion co_t )
; emitWorkNC [evar_t]
; canEqNC evar_s NomEq s1 s2 }
| otherwise -- Can't happen
= error "can_eq_app"
-----------------------
-- | Break apart an equality over a casted type
-- looking like (ty1 |> co1) ~ ty2 (modulo a swap-flag)
canEqCast :: Bool -- are both types flat?
-> CtEvidence
-> EqRel
-> SwapFlag
-> TcType -> Coercion -- LHS (res. RHS), ty1 |> co1
-> TcType -> TcType -- RHS (res. LHS), ty2 both normal and pretty
-> TcS (StopOrContinue Ct)
canEqCast flat ev eq_rel swapped ty1 co1 ty2 ps_ty2
= do { traceTcS "Decomposing cast" (vcat [ ppr ev
, ppr ty1 <+> text "|>" <+> ppr co1
, ppr ps_ty2 ])
; rewriteEqEvidence ev swapped ty1 ps_ty2
(mkTcReflCo role ty1
`mkTcCoherenceRightCo` co1)
(mkTcReflCo role ps_ty2)
`andWhenContinue` \ new_ev ->
can_eq_nc flat new_ev eq_rel ty1 ty1 ty2 ps_ty2 }
where
role = eqRelRole eq_rel
------------------------
canTyConApp :: CtEvidence -> EqRel
-> TyCon -> [TcType]
-> TyCon -> [TcType]
-> TcS (StopOrContinue Ct)
-- See Note [Decomposing TyConApps]
canTyConApp ev eq_rel tc1 tys1 tc2 tys2
| tc1 == tc2
, tys1 `equalLength` tys2
= do { inerts <- getTcSInerts
; if can_decompose inerts
then do { traceTcS "canTyConApp"
(ppr ev $$ ppr eq_rel $$ ppr tc1 $$ ppr tys1 $$ ppr tys2)
; canDecomposableTyConAppOK ev eq_rel tc1 tys1 tys2
; stopWith ev "Decomposed TyConApp" }
else canEqFailure ev eq_rel ty1 ty2 }
-- See Note [Skolem abstract data] (at tyConSkolem)
| tyConSkolem tc1 || tyConSkolem tc2
= do { traceTcS "canTyConApp: skolem abstract" (ppr tc1 $$ ppr tc2)
; continueWith (mkIrredCt ev) }
-- Fail straight away for better error messages
-- See Note [Use canEqFailure in canDecomposableTyConApp]
| eq_rel == ReprEq && not (isGenerativeTyCon tc1 Representational &&
isGenerativeTyCon tc2 Representational)
= canEqFailure ev eq_rel ty1 ty2
| otherwise
= canEqHardFailure ev ty1 ty2
where
ty1 = mkTyConApp tc1 tys1
ty2 = mkTyConApp tc2 tys2
loc = ctEvLoc ev
pred = ctEvPred ev
-- See Note [Decomposing equality]
can_decompose inerts
= isInjectiveTyCon tc1 (eqRelRole eq_rel)
|| (ctEvFlavour ev /= Given && isEmptyBag (matchableGivens loc pred inerts))
{-
Note [Use canEqFailure in canDecomposableTyConApp]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We must use canEqFailure, not canEqHardFailure here, because there is
the possibility of success if working with a representational equality.
Here is one case:
type family TF a where TF Char = Bool
data family DF a
newtype instance DF Bool = MkDF Int
Suppose we are canonicalising (Int ~R DF (TF a)), where we don't yet
know `a`. This is *not* a hard failure, because we might soon learn
that `a` is, in fact, Char, and then the equality succeeds.
Here is another case:
[G] Age ~R Int
where Age's constructor is not in scope. We don't want to report
an "inaccessible code" error in the context of this Given!
For example, see typecheck/should_compile/T10493, repeated here:
import Data.Ord (Down) -- no constructor
foo :: Coercible (Down Int) Int => Down Int -> Int
foo = coerce
That should compile, but only because we use canEqFailure and not
canEqHardFailure.
Note [Decomposing equality]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
If we have a constraint (of any flavour and role) that looks like
T tys1 ~ T tys2, what can we conclude about tys1 and tys2? The answer,
of course, is "it depends". This Note spells it all out.
In this Note, "decomposition" refers to taking the constraint
[fl] (T tys1 ~X T tys2)
(for some flavour fl and some role X) and replacing it with
[fls'] (tys1 ~Xs' tys2)
where that notation indicates a list of new constraints, where the
new constraints may have different flavours and different roles.
The key property to consider is injectivity. When decomposing a Given the
decomposition is sound if and only if T is injective in all of its type
arguments. When decomposing a Wanted, the decomposition is sound (assuming the
correct roles in the produced equality constraints), but it may be a guess --
that is, an unforced decision by the constraint solver. Decomposing Wanteds
over injective TyCons does not entail guessing. But sometimes we want to
decompose a Wanted even when the TyCon involved is not injective! (See below.)
So, in broad strokes, we want this rule:
(*) Decompose a constraint (T tys1 ~X T tys2) if and only if T is injective
at role X.
Pursuing the details requires exploring three axes:
* Flavour: Given vs. Derived vs. Wanted
* Role: Nominal vs. Representational
* TyCon species: datatype vs. newtype vs. data family vs. type family vs. type variable
(So a type variable isn't a TyCon, but it's convenient to put the AppTy case
in the same table.)
Right away, we can say that Derived behaves just as Wanted for the purposes
of decomposition. The difference between Derived and Wanted is the handling of
evidence. Since decomposition in these cases isn't a matter of soundness but of
guessing, we want the same behavior regardless of evidence.
Here is a table (discussion following) detailing where decomposition of
(T s1 ... sn) ~r (T t1 .. tn)
is allowed. The first four lines (Data types ... type family) refer
to TyConApps with various TyCons T; the last line is for AppTy, where
there is presumably a type variable at the head, so it's actually
(s s1 ... sn) ~r (t t1 .. tn)
NOMINAL GIVEN WANTED
Datatype YES YES
Newtype YES YES
Data family YES YES
Type family YES, in injective args{1} YES, in injective args{1}
Type variable YES YES
REPRESENTATIONAL GIVEN WANTED
Datatype YES YES
Newtype NO{2} MAYBE{2}
Data family NO{3} MAYBE{3}
Type family NO NO
Type variable NO{4} NO{4}
{1}: Type families can be injective in some, but not all, of their arguments,
so we want to do partial decomposition. This is quite different than the way
other decomposition is done, where the decomposed equalities replace the original
one. We thus proceed much like we do with superclasses: emitting new Givens
when "decomposing" a partially-injective type family Given and new Deriveds
when "decomposing" a partially-injective type family Wanted. (As of the time of
writing, 13 June 2015, the implementation of injective type families has not
been merged, but it should be soon. Please delete this parenthetical if the
implementation is indeed merged.)
{2}: See Note [Decomposing newtypes at representational role]
{3}: Because of the possibility of newtype instances, we must treat
data families like newtypes. See also Note [Decomposing newtypes at
representational role]. See #10534 and test case
typecheck/should_fail/T10534.
{4}: Because type variables can stand in for newtypes, we conservatively do not
decompose AppTys over representational equality.
In the implementation of can_eq_nc and friends, we don't directly pattern
match using lines like in the tables above, as those tables don't cover
all cases (what about PrimTyCon? tuples?). Instead we just ask about injectivity,
boiling the tables above down to rule (*). The exceptions to rule (*) are for
injective type families, which are handled separately from other decompositions,
and the MAYBE entries above.
Note [Decomposing newtypes at representational role]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
This note discusses the 'newtype' line in the REPRESENTATIONAL table
in Note [Decomposing equality]. (At nominal role, newtypes are fully
decomposable.)
Here is a representative example of why representational equality over
newtypes is tricky:
newtype Nt a = Mk Bool -- NB: a is not used in the RHS,
type role Nt representational -- but the user gives it an R role anyway
If we have [W] Nt alpha ~R Nt beta, we *don't* want to decompose to
[W] alpha ~R beta, because it's possible that alpha and beta aren't
representationally equal. Here's another example.
newtype Nt a = MkNt (Id a)
type family Id a where Id a = a
[W] Nt Int ~R Nt Age
Because of its use of a type family, Nt's parameter will get inferred to have
a nominal role. Thus, decomposing the wanted will yield [W] Int ~N Age, which
is unsatisfiable. Unwrapping, though, leads to a solution.
Conclusion:
* Unwrap newtypes before attempting to decompose them.
This is done in can_eq_nc'.
It all comes from the fact that newtypes aren't necessarily injective
w.r.t. representational equality.
Furthermore, as explained in Note [NthCo and newtypes] in TyCoRep, we can't use
NthCo on representational coercions over newtypes. NthCo comes into play
only when decomposing givens.
Conclusion:
* Do not decompose [G] N s ~R N t
Is it sensible to decompose *Wanted* constraints over newtypes? Yes!
It's the only way we could ever prove (IO Int ~R IO Age), recalling
that IO is a newtype.
However we must be careful. Consider
type role Nt representational
[G] Nt a ~R Nt b (1)
[W] NT alpha ~R Nt b (2)
[W] alpha ~ a (3)
If we focus on (3) first, we'll substitute in (2), and now it's
identical to the given (1), so we succeed. But if we focus on (2)
first, and decompose it, we'll get (alpha ~R b), which is not soluble.
This is exactly like the question of overlapping Givens for class
constraints: see Note [Instance and Given overlap] in TcInteract.
Conclusion:
* Decompose [W] N s ~R N t iff there no given constraint that could
later solve it.
-}
canDecomposableTyConAppOK :: CtEvidence -> EqRel
-> TyCon -> [TcType] -> [TcType]
-> TcS ()
-- Precondition: tys1 and tys2 are the same length, hence "OK"
canDecomposableTyConAppOK ev eq_rel tc tys1 tys2
= case ev of
CtDerived {}
-> unifyDeriveds loc tc_roles tys1 tys2
CtWanted { ctev_dest = dest }
-> do { cos <- zipWith4M unifyWanted new_locs tc_roles tys1 tys2
; setWantedEq dest (mkTyConAppCo role tc cos) }
CtGiven { ctev_evar = evar }
-> do { let ev_co = mkCoVarCo evar
; given_evs <- newGivenEvVars loc $
[ ( mkPrimEqPredRole r ty1 ty2
, evCoercion $ mkNthCo i ev_co )
| (r, ty1, ty2, i) <- zip4 tc_roles tys1 tys2 [0..]
, r /= Phantom
, not (isCoercionTy ty1) && not (isCoercionTy ty2) ]
; emitWorkNC given_evs }
where
loc = ctEvLoc ev
role = eqRelRole eq_rel
tc_roles = tyConRolesX role tc
-- the following makes a better distinction between "kind" and "type"
-- in error messages
bndrs = tyConBinders tc
is_kinds = map isNamedTyConBinder bndrs
is_viss = map isVisibleTyConBinder bndrs
kind_xforms = map (\is_kind -> if is_kind then toKindLoc else id) is_kinds
vis_xforms = map (\is_vis -> if is_vis then id
else flip updateCtLocOrigin toInvisibleOrigin)
is_viss
-- zipWith3 (.) composes its first two arguments and applies it to the third
new_locs = zipWith3 (.) kind_xforms vis_xforms (repeat loc)
-- | Call when canonicalizing an equality fails, but if the equality is
-- representational, there is some hope for the future.
-- Examples in Note [Use canEqFailure in canDecomposableTyConApp]
canEqFailure :: CtEvidence -> EqRel
-> TcType -> TcType -> TcS (StopOrContinue Ct)
canEqFailure ev NomEq ty1 ty2
= canEqHardFailure ev ty1 ty2
canEqFailure ev ReprEq ty1 ty2
= do { (xi1, co1) <- flatten FM_FlattenAll ev ty1
; (xi2, co2) <- flatten FM_FlattenAll ev ty2
-- We must flatten the types before putting them in the
-- inert set, so that we are sure to kick them out when
-- new equalities become available
; traceTcS "canEqFailure with ReprEq" $
vcat [ ppr ev, ppr ty1, ppr ty2, ppr xi1, ppr xi2 ]
; rewriteEqEvidence ev NotSwapped xi1 xi2 co1 co2
`andWhenContinue` \ new_ev ->
continueWith (mkIrredCt new_ev) }
-- | Call when canonicalizing an equality fails with utterly no hope.
canEqHardFailure :: CtEvidence
-> TcType -> TcType -> TcS (StopOrContinue Ct)
-- See Note [Make sure that insolubles are fully rewritten]
canEqHardFailure ev ty1 ty2
= do { (s1, co1) <- flatten FM_SubstOnly ev ty1
; (s2, co2) <- flatten FM_SubstOnly ev ty2
; rewriteEqEvidence ev NotSwapped s1 s2 co1 co2
`andWhenContinue` \ new_ev ->
continueWith (mkInsolubleCt new_ev) }
{-
Note [Decomposing TyConApps]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
If we see (T s1 t1 ~ T s2 t2), then we can just decompose to
(s1 ~ s2, t1 ~ t2)
and push those back into the work list. But if
s1 = K k1 s2 = K k2
then we will just decomopose s1~s2, and it might be better to
do so on the spot. An important special case is where s1=s2,
and we get just Refl.
So canDecomposableTyCon is a fast-path decomposition that uses
unifyWanted etc to short-cut that work.
Note [Canonicalising type applications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Given (s1 t1) ~ ty2, how should we proceed?
The simple things is to see if ty2 is of form (s2 t2), and
decompose. By this time s1 and s2 can't be saturated type
function applications, because those have been dealt with
by an earlier equation in can_eq_nc, so it is always sound to
decompose.
However, over-eager decomposition gives bad error messages
for things like
a b ~ Maybe c
e f ~ p -> q
Suppose (in the first example) we already know a~Array. Then if we
decompose the application eagerly, yielding
a ~ Maybe
b ~ c
we get an error "Can't match Array ~ Maybe",
but we'd prefer to get "Can't match Array b ~ Maybe c".
So instead can_eq_wanted_app flattens the LHS and RHS, in the hope of
replacing (a b) by (Array b), before using try_decompose_app to
decompose it.
Note [Make sure that insolubles are fully rewritten]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When an equality fails, we still want to rewrite the equality
all the way down, so that it accurately reflects
(a) the mutable reference substitution in force at start of solving
(b) any ty-binds in force at this point in solving
See Note [Rewrite insolubles] in TcSMonad.
And if we don't do this there is a bad danger that
TcSimplify.applyTyVarDefaulting will find a variable
that has in fact been substituted.
Note [Do not decompose Given polytype equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider [G] (forall a. t1 ~ forall a. t2). Can we decompose this?
No -- what would the evidence look like? So instead we simply discard
this given evidence.
Note [Combining insoluble constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
As this point we have an insoluble constraint, like Int~Bool.
* If it is Wanted, delete it from the cache, so that subsequent
Int~Bool constraints give rise to separate error messages
* But if it is Derived, DO NOT delete from cache. A class constraint
may get kicked out of the inert set, and then have its functional
dependency Derived constraints generated a second time. In that
case we don't want to get two (or more) error messages by
generating two (or more) insoluble fundep constraints from the same
class constraint.
Note [No top-level newtypes on RHS of representational equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we're in this situation:
work item: [W] c1 : a ~R b
inert: [G] c2 : b ~R Id a
where
newtype Id a = Id a
We want to make sure canEqTyVar sees [W] a ~R a, after b is flattened
and the Id newtype is unwrapped. This is assured by requiring only flat
types in canEqTyVar *and* having the newtype-unwrapping check above
the tyvar check in can_eq_nc.
Note [Occurs check error]
~~~~~~~~~~~~~~~~~~~~~~~~~
If we have an occurs check error, are we necessarily hosed? Say our
tyvar is tv1 and the type it appears in is xi2. Because xi2 is function
free, then if we're computing w.r.t. nominal equality, then, yes, we're
hosed. Nothing good can come from (a ~ [a]). If we're computing w.r.t.
representational equality, this is a little subtler. Once again, (a ~R [a])
is a bad thing, but (a ~R N a) for a newtype N might be just fine. This
means also that (a ~ b a) might be fine, because `b` might become a newtype.
So, we must check: does tv1 appear in xi2 under any type constructor
that is generative w.r.t. representational equality? That's what
isInsolubleOccursCheck does.
See also #10715, which induced this addition.
-}
canCFunEqCan :: CtEvidence
-> TyCon -> [TcType] -- LHS
-> TcTyVar -- RHS
-> TcS (StopOrContinue Ct)
-- ^ Canonicalise a CFunEqCan. We know that
-- the arg types are already flat,
-- and the RHS is a fsk, which we must *not* substitute.
-- So just substitute in the LHS
canCFunEqCan ev fn tys fsk
= do { (tys', cos) <- flattenManyNom ev tys
-- cos :: tys' ~ tys
; let lhs_co = mkTcTyConAppCo Nominal fn cos
-- :: F tys' ~ F tys
new_lhs = mkTyConApp fn tys'
fsk_ty = mkTyVarTy fsk
; rewriteEqEvidence ev NotSwapped new_lhs fsk_ty
lhs_co (mkTcNomReflCo fsk_ty)
`andWhenContinue` \ ev' ->
do { extendFlatCache fn tys' (ctEvCoercion ev', fsk_ty, ctEvFlavour ev')
; continueWith (CFunEqCan { cc_ev = ev', cc_fun = fn
, cc_tyargs = tys', cc_fsk = fsk }) } }
---------------------
canEqTyVar :: CtEvidence -- ev :: lhs ~ rhs
-> EqRel -> SwapFlag
-> TcTyVar -> CoercionN -- tv1 |> co1
-> TcType -- lhs: pretty lhs, already flat
-> TcType -> TcType -- rhs: already flat
-> TcS (StopOrContinue Ct)
canEqTyVar ev eq_rel swapped tv1 co1 ps_ty1 xi2 ps_xi2
| k1 `eqType` k2
= canEqTyVarHomo ev eq_rel swapped tv1 co1 ps_ty1 xi2 ps_xi2
-- See Note [Equalities with incompatible kinds]
| CtGiven { ctev_evar = evar } <- ev
-- unswapped: tm :: (lhs :: k1) ~ (rhs :: k2)
-- swapped : tm :: (rhs :: k2) ~ (lhs :: k1)
= do { kind_ev_id <- newBoundEvVarId kind_pty
(evCoercion $
if isSwapped swapped
then mkTcSymCo $ mkTcKindCo $ mkTcCoVarCo evar
else mkTcKindCo $ mkTcCoVarCo evar)
-- kind_ev_id :: (k1 :: *) ~ (k2 :: *) (whether swapped or not)
; let kind_ev = CtGiven { ctev_pred = kind_pty
, ctev_evar = kind_ev_id
, ctev_loc = kind_loc }
homo_co = mkSymCo $ mkCoVarCo kind_ev_id
rhs' = mkCastTy xi2 homo_co
ps_rhs' = mkCastTy ps_xi2 homo_co
; traceTcS "Hetero equality gives rise to given kind equality"
(ppr kind_ev_id <+> dcolon <+> ppr kind_pty)
; emitWorkNC [kind_ev]
; type_ev <- newGivenEvVar loc $
if isSwapped swapped
then ( mkTcEqPredLikeEv ev rhs' lhs
, evCoercion $
mkTcCoherenceLeftCo (mkTcCoVarCo evar) homo_co )
else ( mkTcEqPredLikeEv ev lhs rhs'
, evCoercion $
mkTcCoherenceRightCo (mkTcCoVarCo evar) homo_co )
-- unswapped: type_ev :: (lhs :: k1) ~ ((rhs |> sym kind_ev_id) :: k1)
-- swapped : type_ev :: ((rhs |> sym kind_ev_id) :: k1) ~ (lhs :: k1)
; canEqTyVarHomo type_ev eq_rel swapped tv1 co1 ps_ty1 rhs' ps_rhs' }
-- See Note [Equalities with incompatible kinds]
| otherwise -- Wanted and Derived
-- NB: all kind equalities are Nominal
= do { emitNewDerivedEq kind_loc Nominal k1 k2
-- kind_ev :: (k1 :: *) ~ (k2 :: *)
; traceTcS "Hetero equality gives rise to derived kind equality" $
ppr ev
; continueWith (mkIrredCt ev) }
where
lhs = mkTyVarTy tv1 `mkCastTy` co1
Pair _ k1 = coercionKind co1
k2 = typeKind xi2
kind_pty = mkHeteroPrimEqPred liftedTypeKind liftedTypeKind k1 k2
kind_loc = mkKindLoc lhs xi2 loc
loc = ctev_loc ev
-- guaranteed that typeKind lhs == typeKind rhs
canEqTyVarHomo :: CtEvidence
-> EqRel -> SwapFlag
-> TcTyVar -> CoercionN -- lhs: tv1 |> co1
-> TcType -- pretty lhs
-> TcType -> TcType -- rhs (might not be flat)
-> TcS (StopOrContinue Ct)
canEqTyVarHomo ev eq_rel swapped tv1 co1 ps_ty1 ty2 _
| Just (tv2, _) <- tcGetCastedTyVar_maybe ty2
, tv1 == tv2
= canEqReflexive ev eq_rel (mkTyVarTy tv1 `mkCastTy` co1)
-- we don't need to check co2 because its type must match co1
| Just (tv2, co2) <- tcGetCastedTyVar_maybe ty2
, swapOverTyVars tv1 tv2
= do { traceTcS "canEqTyVar" (ppr tv1 $$ ppr tv2 $$ ppr swapped)
-- FM_Avoid commented out: see Note [Lazy flattening] in TcFlatten
-- let fmode = FE { fe_ev = ev, fe_mode = FM_Avoid tv1' True }
-- Flatten the RHS less vigorously, to avoid gratuitous flattening
-- True <=> xi2 should not itself be a type-function application
; dflags <- getDynFlags
; canEqTyVar2 dflags ev eq_rel (flipSwap swapped) tv2 co2 ps_ty1 }
canEqTyVarHomo ev eq_rel swapped tv1 co1 _ _ ps_ty2
= do { dflags <- getDynFlags
; canEqTyVar2 dflags ev eq_rel swapped tv1 co1 ps_ty2 }
-- The RHS here is either not a casted tyvar, or it's a tyvar but we want
-- to rewrite the LHS to the RHS (as per swapOverTyVars)
canEqTyVar2 :: DynFlags
-> CtEvidence -- lhs ~ rhs (or, if swapped, orhs ~ olhs)
-> EqRel
-> SwapFlag
-> TcTyVar -> CoercionN -- lhs = tv |> co, flat
-> TcType -- rhs
-> TcS (StopOrContinue Ct)
-- LHS is an inert type variable,
-- and RHS is fully rewritten, but with type synonyms
-- preserved as much as possible
canEqTyVar2 dflags ev eq_rel swapped tv1 co1 orhs
| Just nrhs' <- metaTyVarUpdateOK dflags tv1 nrhs -- No occurs check
-- Must do the occurs check even on tyvar/tyvar
-- equalities, in case have x ~ (y :: ..x...)
-- Trac #12593
= rewriteEqEvidence ev swapped nlhs nrhs' rewrite_co1 rewrite_co2
`andWhenContinue` \ new_ev ->
continueWith (CTyEqCan { cc_ev = new_ev, cc_tyvar = tv1
, cc_rhs = nrhs', cc_eq_rel = eq_rel })
| otherwise -- For some reason (occurs check, or forall) we can't unify
-- We must not use it for further rewriting!
= do { traceTcS "canEqTyVar2 can't unify" (ppr tv1 $$ ppr nrhs)
; rewriteEqEvidence ev swapped nlhs nrhs rewrite_co1 rewrite_co2
`andWhenContinue` \ new_ev ->
if isInsolubleOccursCheck eq_rel tv1 nrhs
then continueWith (mkInsolubleCt new_ev)
-- If we have a ~ [a], it is not canonical, and in particular
-- we don't want to rewrite existing inerts with it, otherwise
-- we'd risk divergence in the constraint solver
else continueWith (mkIrredCt new_ev) }
-- A representational equality with an occurs-check problem isn't
-- insoluble! For example:
-- a ~R b a
-- We might learn that b is the newtype Id.
-- But, the occurs-check certainly prevents the equality from being
-- canonical, and we might loop if we were to use it in rewriting.
where
role = eqRelRole eq_rel
nlhs = mkTyVarTy tv1
nrhs = orhs `mkCastTy` mkTcSymCo co1
-- rewrite_co1 :: tv1 ~ (tv1 |> co1)
-- rewrite_co2 :: (rhs |> sym co1) ~ rhs
rewrite_co1 = mkTcReflCo role nlhs `mkTcCoherenceRightCo` co1
rewrite_co2 = mkTcReflCo role orhs `mkTcCoherenceLeftCo` mkTcSymCo co1
-- | Solve a reflexive equality constraint
canEqReflexive :: CtEvidence -- ty ~ ty
-> EqRel
-> TcType -- ty
-> TcS (StopOrContinue Ct) -- always Stop
canEqReflexive ev eq_rel ty
= do { setEvBindIfWanted ev (evCoercion $
mkTcReflCo (eqRelRole eq_rel) ty)
; stopWith ev "Solved by reflexivity" }
{-
Note [Canonical orientation for tyvar/tyvar equality constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When we have a ~ b where both 'a' and 'b' are TcTyVars, which way
round should be oriented in the CTyEqCan? The rules, implemented by
canEqTyVarTyVar, are these
* If either is a flatten-meta-variables, it goes on the left.
* Put a meta-tyvar on the left if possible
alpha[3] ~ r
* If both are meta-tyvars, put the more touchable one (deepest level
number) on the left, so there is the best chance of unifying it
alpha[3] ~ beta[2]
* If both are meta-tyvars and both at the same level, put a SigTv
on the right if possible
alpha[2] ~ beta[2](sig-tv)
That way, when we unify alpha := beta, we don't lose the SigTv flag.
* Put a meta-tv with a System Name on the left if possible so it
gets eliminated (improves error messages)
* If one is a flatten-skolem, put it on the left so that it is
substituted out Note [Elminate flat-skols]
fsk ~ a
Note [Eliminate flat-skols]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have [G] Num (F [a])
then we flatten to
[G] Num fsk
[G] F [a] ~ fsk
where fsk is a flatten-skolem (FlatSkolTv). Suppose we have
type instance F [a] = a
then we'll reduce the second constraint to
[G] a ~ fsk
and then replace all uses of 'a' with fsk. That's bad because
in error messages intead of saying 'a' we'll say (F [a]). In all
places, including those where the programmer wrote 'a' in the first
place. Very confusing! See Trac #7862.
Solution: re-orient a~fsk to fsk~a, so that we preferentially eliminate
the fsk.
Note [Equalities with incompatible kinds]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
What do we do when we have an equality
(tv :: k1) ~ (rhs :: k2)
where k1 and k2 differ? This Note explores this treacherous area.
First off, the question above is slightly the wrong question. Flattening
a tyvar will flatten its kind (Note [Flattening] in TcFlatten); flattening
the kind might introduce a cast. So we might have a casted tyvar on the
left. We thus revise our test case to
(tv |> co :: k1) ~ (rhs :: k2)
We must proceed differently here depending on whether we have a Wanted
or a Given. Consider this:
[W] w :: (alpha :: k) ~ (Int :: Type)
where k is a skolem. One possible way forward is this:
[W] co :: k ~ Type
[W] w :: (alpha :: k) ~ (Int |> sym co :: k)
The next step will be to unify
alpha := Int |> sym co
Now, consider what error we'll report if we can't solve the "co"
wanted. Its CtOrigin is the w wanted... which now reads (after zonking)
Int ~ Int. The user thus sees that GHC can't solve Int ~ Int, which
is embarrassing. See #11198 for more tales of destruction.
The reason for this odd behavior is much the same as
Note [Wanteds do not rewrite Wanteds] in TcRnTypes: note that the
new `co` is a Wanted. The solution is then not to use `co` to "rewrite"
-- that is, cast -- `w`, but instead to keep `w` heterogeneous and irreducible.
Given that we're not using `co`, there is no reason to collect evidence
for it, so `co` is born a Derived. When the Derived is solved (by unification),
the original wanted (`w`) will get kicked out.
Note that, if we had [G] co1 :: k ~ Type available, then none of this code would
trigger, because flattening would have rewritten k to Type. That is,
`w` would look like [W] (alpha |> co1 :: Type) ~ (Int :: Type), and the tyvar
case will trigger, correctly rewriting alpha to (Int |> sym co1).
Successive canonicalizations of the same Wanted may produce
duplicate Deriveds. Similar duplications can happen with fundeps, and there
seems to be no easy way to avoid. I expect this case to be rare.
For Givens, this problem doesn't bite, so a heterogeneous Given gives
rise to a Given kind equality. No Deriveds here. We thus homogenise
the Given (see the "homo_co" in the Given case in canEqTyVar) and
carry on with a homogeneous equality constraint.
Separately, I (Richard E) spent some time pondering what to do in the case
that we have [W] (tv |> co1 :: k1) ~ (tv |> co2 :: k2) where k1 and k2
differ. Note that the tv is the same. (This case is handled as the first
case in canEqTyVarHomo.) At one point, I thought we could solve this limited
form of heterogeneous Wanted, but I then reconsidered and now treat this case
just like any other heterogeneous Wanted.
Note [Type synonyms and canonicalization]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We treat type synonym applications as xi types, that is, they do not
count as type function applications. However, we do need to be a bit
careful with type synonyms: like type functions they may not be
generative or injective. However, unlike type functions, they are
parametric, so there is no problem in expanding them whenever we see
them, since we do not need to know anything about their arguments in
order to expand them; this is what justifies not having to treat them
as specially as type function applications. The thing that causes
some subtleties is that we prefer to leave type synonym applications
*unexpanded* whenever possible, in order to generate better error
messages.
If we encounter an equality constraint with type synonym applications
on both sides, or a type synonym application on one side and some sort
of type application on the other, we simply must expand out the type
synonyms in order to continue decomposing the equality constraint into
primitive equality constraints. For example, suppose we have
type F a = [Int]
and we encounter the equality
F a ~ [b]
In order to continue we must expand F a into [Int], giving us the
equality
[Int] ~ [b]
which we can then decompose into the more primitive equality
constraint
Int ~ b.
However, if we encounter an equality constraint with a type synonym
application on one side and a variable on the other side, we should
NOT (necessarily) expand the type synonym, since for the purpose of
good error messages we want to leave type synonyms unexpanded as much
as possible. Hence the ps_ty1, ps_ty2 argument passed to canEqTyVar.
-}
{-
************************************************************************
* *
Evidence transformation
* *
************************************************************************
-}
data StopOrContinue a
= ContinueWith a -- The constraint was not solved, although it may have
-- been rewritten
| Stop CtEvidence -- The (rewritten) constraint was solved
SDoc -- Tells how it was solved
-- Any new sub-goals have been put on the work list
instance Functor StopOrContinue where
fmap f (ContinueWith x) = ContinueWith (f x)
fmap _ (Stop ev s) = Stop ev s
instance Outputable a => Outputable (StopOrContinue a) where
ppr (Stop ev s) = text "Stop" <> parens s <+> ppr ev
ppr (ContinueWith w) = text "ContinueWith" <+> ppr w
continueWith :: a -> TcS (StopOrContinue a)
continueWith = return . ContinueWith
stopWith :: CtEvidence -> String -> TcS (StopOrContinue a)
stopWith ev s = return (Stop ev (text s))
andWhenContinue :: TcS (StopOrContinue a)
-> (a -> TcS (StopOrContinue b))
-> TcS (StopOrContinue b)
andWhenContinue tcs1 tcs2
= do { r <- tcs1
; case r of
Stop ev s -> return (Stop ev s)
ContinueWith ct -> tcs2 ct }
infixr 0 `andWhenContinue` -- allow chaining with ($)
rewriteEvidence :: CtEvidence -- old evidence
-> TcPredType -- new predicate
-> TcCoercion -- Of type :: new predicate ~
-> TcS (StopOrContinue CtEvidence)
-- Returns Just new_ev iff either (i) 'co' is reflexivity
-- or (ii) 'co' is not reflexivity, and 'new_pred' not cached
-- In either case, there is nothing new to do with new_ev
{-
rewriteEvidence old_ev new_pred co
Main purpose: create new evidence for new_pred;
unless new_pred is cached already
* Returns a new_ev : new_pred, with same wanted/given/derived flag as old_ev
* If old_ev was wanted, create a binding for old_ev, in terms of new_ev
* If old_ev was given, AND not cached, create a binding for new_ev, in terms of old_ev
* Returns Nothing if new_ev is already cached
Old evidence New predicate is Return new evidence
flavour of same flavor
-------------------------------------------------------------------
Wanted Already solved or in inert Nothing
or Derived Not Just new_evidence
Given Already in inert Nothing
Not Just new_evidence
Note [Rewriting with Refl]
~~~~~~~~~~~~~~~~~~~~~~~~~~
If the coercion is just reflexivity then you may re-use the same
variable. But be careful! Although the coercion is Refl, new_pred
may reflect the result of unification alpha := ty, so new_pred might
not _look_ the same as old_pred, and it's vital to proceed from now on
using new_pred.
qThe flattener preserves type synonyms, so they should appear in new_pred
as well as in old_pred; that is important for good error messages.
-}
rewriteEvidence old_ev@(CtDerived {}) new_pred _co
= -- If derived, don't even look at the coercion.
-- This is very important, DO NOT re-order the equations for
-- rewriteEvidence to put the isTcReflCo test first!
-- Why? Because for *Derived* constraints, c, the coercion, which
-- was produced by flattening, may contain suspended calls to
-- (ctEvExpr c), which fails for Derived constraints.
-- (Getting this wrong caused Trac #7384.)
continueWith (old_ev { ctev_pred = new_pred })
rewriteEvidence old_ev new_pred co
| isTcReflCo co -- See Note [Rewriting with Refl]
= continueWith (old_ev { ctev_pred = new_pred })
rewriteEvidence ev@(CtGiven { ctev_evar = old_evar, ctev_loc = loc }) new_pred co
= do { new_ev <- newGivenEvVar loc (new_pred, new_tm)
; continueWith new_ev }
where
-- mkEvCast optimises ReflCo
new_tm = mkEvCast (evId old_evar) (tcDowngradeRole Representational
(ctEvRole ev)
(mkTcSymCo co))
rewriteEvidence ev@(CtWanted { ctev_dest = dest
, ctev_loc = loc }) new_pred co
= do { mb_new_ev <- newWanted loc new_pred
; MASSERT( tcCoercionRole co == ctEvRole ev )
; setWantedEvTerm dest
(EvExpr $ mkEvCast (getEvExpr mb_new_ev)
(tcDowngradeRole Representational (ctEvRole ev) co))
; case mb_new_ev of
Fresh new_ev -> continueWith new_ev
Cached _ -> stopWith ev "Cached wanted" }
rewriteEqEvidence :: CtEvidence -- Old evidence :: olhs ~ orhs (not swapped)
-- or orhs ~ olhs (swapped)
-> SwapFlag
-> TcType -> TcType -- New predicate nlhs ~ nrhs
-- Should be zonked, because we use typeKind on nlhs/nrhs
-> TcCoercion -- lhs_co, of type :: nlhs ~ olhs
-> TcCoercion -- rhs_co, of type :: nrhs ~ orhs
-> TcS (StopOrContinue CtEvidence) -- Of type nlhs ~ nrhs
-- For (rewriteEqEvidence (Given g olhs orhs) False nlhs nrhs lhs_co rhs_co)
-- we generate
-- If not swapped
-- g1 : nlhs ~ nrhs = lhs_co ; g ; sym rhs_co
-- If 'swapped'
-- g1 : nlhs ~ nrhs = lhs_co ; Sym g ; sym rhs_co
--
-- For (Wanted w) we do the dual thing.
-- New w1 : nlhs ~ nrhs
-- If not swapped
-- w : olhs ~ orhs = sym lhs_co ; w1 ; rhs_co
-- If swapped
-- w : orhs ~ olhs = sym rhs_co ; sym w1 ; lhs_co
--
-- It's all a form of rewwriteEvidence, specialised for equalities
rewriteEqEvidence old_ev swapped nlhs nrhs lhs_co rhs_co
| CtDerived {} <- old_ev -- Don't force the evidence for a Derived
= continueWith (old_ev { ctev_pred = new_pred })
| NotSwapped <- swapped
, isTcReflCo lhs_co -- See Note [Rewriting with Refl]
, isTcReflCo rhs_co
= continueWith (old_ev { ctev_pred = new_pred })
| CtGiven { ctev_evar = old_evar } <- old_ev
= do { let new_tm = evCoercion (lhs_co
`mkTcTransCo` maybeSym swapped (mkTcCoVarCo old_evar)
`mkTcTransCo` mkTcSymCo rhs_co)
; new_ev <- newGivenEvVar loc' (new_pred, new_tm)
; continueWith new_ev }
| CtWanted { ctev_dest = dest } <- old_ev
= do { (new_ev, hole_co) <- newWantedEq loc' (ctEvRole old_ev) nlhs nrhs
; let co = maybeSym swapped $
mkSymCo lhs_co
`mkTransCo` hole_co
`mkTransCo` rhs_co
; setWantedEq dest co
; traceTcS "rewriteEqEvidence" (vcat [ppr old_ev, ppr nlhs, ppr nrhs, ppr co])
; continueWith new_ev }
| otherwise
= panic "rewriteEvidence"
where
new_pred = mkTcEqPredLikeEv old_ev nlhs nrhs
-- equality is like a type class. Bumping the depth is necessary because
-- of recursive newtypes, where "reducing" a newtype can actually make
-- it bigger. See Note [Newtypes can blow the stack].
loc = ctEvLoc old_ev
loc' = bumpCtLocDepth loc
{- Note [unifyWanted and unifyDerived]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When decomposing equalities we often create new wanted constraints for
(s ~ t). But what if s=t? Then it'd be faster to return Refl right away.
Similar remarks apply for Derived.
Rather than making an equality test (which traverses the structure of the
type, perhaps fruitlessly, unifyWanted traverses the common structure, and
bales out when it finds a difference by creating a new Wanted constraint.
But where it succeeds in finding common structure, it just builds a coercion
to reflect it.
-}
unifyWanted :: CtLoc -> Role
-> TcType -> TcType -> TcS Coercion
-- Return coercion witnessing the equality of the two types,
-- emitting new work equalities where necessary to achieve that
-- Very good short-cut when the two types are equal, or nearly so
-- See Note [unifyWanted and unifyDerived]
-- The returned coercion's role matches the input parameter
unifyWanted loc Phantom ty1 ty2
= do { kind_co <- unifyWanted loc Nominal (typeKind ty1) (typeKind ty2)
; return (mkPhantomCo kind_co ty1 ty2) }
unifyWanted loc role orig_ty1 orig_ty2
= go orig_ty1 orig_ty2
where
go ty1 ty2 | Just ty1' <- tcView ty1 = go ty1' ty2
go ty1 ty2 | Just ty2' <- tcView ty2 = go ty1 ty2'
go (FunTy s1 t1) (FunTy s2 t2)
= do { co_s <- unifyWanted loc role s1 s2
; co_t <- unifyWanted loc role t1 t2
; return (mkFunCo role co_s co_t) }
go (TyConApp tc1 tys1) (TyConApp tc2 tys2)
| tc1 == tc2, tys1 `equalLength` tys2
, isInjectiveTyCon tc1 role -- don't look under newtypes at Rep equality
= do { cos <- zipWith3M (unifyWanted loc)
(tyConRolesX role tc1) tys1 tys2
; return (mkTyConAppCo role tc1 cos) }
go ty1@(TyVarTy tv) ty2
= do { mb_ty <- isFilledMetaTyVar_maybe tv
; case mb_ty of
Just ty1' -> go ty1' ty2
Nothing -> bale_out ty1 ty2}
go ty1 ty2@(TyVarTy tv)
= do { mb_ty <- isFilledMetaTyVar_maybe tv
; case mb_ty of
Just ty2' -> go ty1 ty2'
Nothing -> bale_out ty1 ty2 }
go ty1@(CoercionTy {}) (CoercionTy {})
= return (mkReflCo role ty1) -- we just don't care about coercions!
go ty1 ty2 = bale_out ty1 ty2
bale_out ty1 ty2
| ty1 `tcEqType` ty2 = return (mkTcReflCo role ty1)
-- Check for equality; e.g. a ~ a, or (m a) ~ (m a)
| otherwise = emitNewWantedEq loc role orig_ty1 orig_ty2
unifyDeriveds :: CtLoc -> [Role] -> [TcType] -> [TcType] -> TcS ()
-- See Note [unifyWanted and unifyDerived]
unifyDeriveds loc roles tys1 tys2 = zipWith3M_ (unify_derived loc) roles tys1 tys2
unifyDerived :: CtLoc -> Role -> Pair TcType -> TcS ()
-- See Note [unifyWanted and unifyDerived]
unifyDerived loc role (Pair ty1 ty2) = unify_derived loc role ty1 ty2
unify_derived :: CtLoc -> Role -> TcType -> TcType -> TcS ()
-- Create new Derived and put it in the work list
-- Should do nothing if the two types are equal
-- See Note [unifyWanted and unifyDerived]
unify_derived _ Phantom _ _ = return ()
unify_derived loc role orig_ty1 orig_ty2
= go orig_ty1 orig_ty2
where
go ty1 ty2 | Just ty1' <- tcView ty1 = go ty1' ty2
go ty1 ty2 | Just ty2' <- tcView ty2 = go ty1 ty2'
go (FunTy s1 t1) (FunTy s2 t2)
= do { unify_derived loc role s1 s2
; unify_derived loc role t1 t2 }
go (TyConApp tc1 tys1) (TyConApp tc2 tys2)
| tc1 == tc2, tys1 `equalLength` tys2
, isInjectiveTyCon tc1 role
= unifyDeriveds loc (tyConRolesX role tc1) tys1 tys2
go ty1@(TyVarTy tv) ty2
= do { mb_ty <- isFilledMetaTyVar_maybe tv
; case mb_ty of
Just ty1' -> go ty1' ty2
Nothing -> bale_out ty1 ty2 }
go ty1 ty2@(TyVarTy tv)
= do { mb_ty <- isFilledMetaTyVar_maybe tv
; case mb_ty of
Just ty2' -> go ty1 ty2'
Nothing -> bale_out ty1 ty2 }
go ty1 ty2 = bale_out ty1 ty2
bale_out ty1 ty2
| ty1 `tcEqType` ty2 = return ()
-- Check for equality; e.g. a ~ a, or (m a) ~ (m a)
| otherwise = emitNewDerivedEq loc role orig_ty1 orig_ty2
maybeSym :: SwapFlag -> TcCoercion -> TcCoercion
maybeSym IsSwapped co = mkTcSymCo co
maybeSym NotSwapped co = co