{-# LANGUAGE CPP #-} {-# LANGUAGE DeriveFunctor #-} module GHC.Tc.Solver.Canonical( canonicalize, unifyDerived, makeSuperClasses, maybeSym, StopOrContinue(..), stopWith, continueWith, solveCallStack -- For GHC.Tc.Solver ) where #include "HsVersions.h" import GHC.Prelude import GHC.Tc.Types.Constraint import GHC.Core.Predicate import GHC.Tc.Types.Origin import GHC.Tc.Utils.Unify( swapOverTyVars, metaTyVarUpdateOK, MetaTyVarUpdateResult(..) ) import GHC.Tc.Utils.TcType import GHC.Core.Type import GHC.Tc.Solver.Flatten import GHC.Tc.Solver.Monad import GHC.Tc.Types.Evidence import GHC.Tc.Types.EvTerm import GHC.Core.Class import GHC.Core.TyCon import GHC.Core.Multiplicity import GHC.Core.TyCo.Rep -- cleverly decomposes types, good for completeness checking import GHC.Core.Coercion import GHC.Core import GHC.Types.Id( mkTemplateLocals ) import GHC.Core.FamInstEnv ( FamInstEnvs ) import GHC.Tc.Instance.Family ( tcTopNormaliseNewTypeTF_maybe ) import GHC.Types.Var import GHC.Types.Var.Env( mkInScopeSet ) import GHC.Types.Var.Set( delVarSetList ) import GHC.Utils.Outputable import GHC.Driver.Session( DynFlags ) import GHC.Types.Name.Set import GHC.Types.Name.Reader import GHC.Hs.Type( HsIPName(..) ) import GHC.Data.Pair import GHC.Utils.Misc import GHC.Data.Bag import GHC.Utils.Monad import Control.Monad import Data.Maybe ( isJust ) import Data.List ( zip4 ) import GHC.Types.Basic import Data.Bifunctor ( bimap ) import Data.Foldable ( traverse_ ) {- ************************************************************************ * * * 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. 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 GHC.Tc.Types). 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 pred 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 pred) canIrred OtherCIS ev ForAllPred tvs theta p -> do traceTcS "canEvNC:forall" (ppr pred) canForAllNC ev tvs theta p where pred = ctEvPred ev canonicalize (CQuantCan (QCI { qci_ev = ev, qci_pend_sc = pend_sc })) = canForAll ev pend_sc canonicalize (CIrredCan { cc_ev = ev, cc_status = status }) | EqPred eq_rel ty1 ty2 <- classifyPredType (ctEvPred ev) = -- For insolubles (all of which are equalities, do /not/ flatten the arguments -- In #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 = canIrred status 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 {- ************************************************************************ * * * 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 GHC.Tc.Types.Evidence -- and Note [Solving CallStack constraints] in GHC.Tc.Solver.Monad = 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 GHC.Tc.Solver 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)) setEvBindIfWanted ev 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, _kind_co) <- flattenArgsNom ev cls_tc tys ; MASSERT( isTcReflCo _kind_co ) ; let co = mkTcTyConAppCo Nominal cls_tc 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) } where cls_tc = classTyCon cls {- 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 (#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 using mkStrictSuperClasses in canClassNC, when we take a non-canonical Given constraint and cannonicalise it. However stop if you encounter the same class twice. That is, mkStrictSuperClasses expands eagerly, but has a conservative termination condition: see Note [Expanding superclasses] in GHC.Tc.Utils.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 superclasses for new /given/ /non-canonical/ constraints (canClassNC does this). As #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 GHC.Tc.Types.Constraint) try harder: take both the Givens and Wanteds, and expand superclasses again. See the calls to expandSuperClasses in GHC.Tc.Solver.simpl_loop and solveWanteds. This may succeed in generating (a finite number of) extra Givens, and extra Deriveds. Both may help the proof. 3a An important wrinkle: only expand Givens from the current level. Two reasons: - We only want to expand it once, and that is best done at the level it is bound, rather than repeatedly at the leaves of the implication tree - We may be inside a type where we can't create term-level evidence anyway, so we can't superclass-expand, say, (a ~ b) to get (a ~# b). This happened in #15290. 4. Go round to (2) again. This loop (2,3,4) is implemented in GHC.Tc.Solver.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 [Superclass loops] ~~~~~~~~~~~~~~~~~~~~~~~ Suppose we have class C a => D a class D a => C a Then, when we expand superclasses, we'll get back to the self-same predicate, so we have reached a fixpoint in expansion and there is no point in fruitlessly expanding further. This case just falls out from our strategy. Consider f :: C a => a -> Bool f x = x==x Then canClassNC gets the [G] d1: C a constraint, and eager emits superclasses G] d2: D a, [G] d3: C a (psc). (The "psc" means it has its sc_pend flag set.) When processing d3 we find a match with d1 in the inert set, and we always keep the inert item (d1) if possible: see Note [Replacement vs keeping] in GHC.Tc.Solver.Interact. So d3 dies a quick, happy death. Note [Eagerly expand given superclasses] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In step (1) of Note [The superclass story], why do we eagerly expand Given superclasses by one layer? (By "one layer" we mean expand transitively until you meet the same class again -- the conservative criterion embodied in expandSuperClasses. So a "layer" might be a whole stack of superclasses.) We do this eagerly for Givens 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 GHC.Tc.Solver.Interact. 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. This same scenario happens with quantified constraints, whose superclasses are also eagerly expanded. Test case: typecheck/should_compile/T16502b These are handled in canForAllNC, analogously to 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 #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. Note [Nested quantified constraint superclasses] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider (typecheck/should_compile/T17202) class C1 a class (forall c. C1 c) => C2 a class (forall b. (b ~ F a) => C2 a) => C3 a Elsewhere in the code, we get a [G] g1 :: C3 a. We expand its superclass to get [G] g2 :: (forall b. (b ~ F a) => C2 a). This constraint has a superclass, as well. But we now must be careful: we cannot just add (forall c. C1 c) as a Given, because we need to remember g2's context. That new constraint is Given only when forall b. (b ~ F a) is true. It's tempting to make the new Given be (forall b. (b ~ F a) => forall c. C1 c), but that's problematic, because it's nested, and ForAllPred is not capable of representing a nested quantified constraint. (We could change ForAllPred to allow this, but the solution in this Note is much more local and simpler.) So, we swizzle it around to get (forall b c. (b ~ F a) => C1 c). More generally, if we are expanding the superclasses of g0 :: forall tvs. theta => cls tys and find a superclass constraint forall sc_tvs. sc_theta => sc_inner_pred we must have a selector sel_id :: forall cls_tvs. cls cls_tvs -> forall sc_tvs. sc_theta => sc_inner_pred and thus build g_sc :: forall tvs sc_tvs. theta => sc_theta => sc_inner_pred g_sc = /\ tvs. /\ sc_tvs. \ theta_ids. \ sc_theta_ids. sel_id tys (g0 tvs theta_ids) sc_tvs sc_theta_ids Actually, we cheat a bit by eta-reducing: note that sc_theta_ids are both the last bound variables and the last arguments. This avoids the need to produce the sc_theta_ids at all. So our final construction is g_sc = /\ tvs. /\ sc_tvs. \ theta_ids. sel_id tys (g0 tvs theta_ids) sc_tvs -} 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 GHC.Tc.Utils.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 (CQuantCan (QCI { qci_pred = pred, qci_ev = ev })) = ASSERT2( isClassPred pred, ppr pred ) -- The cts should all have -- class pred heads mkStrictSuperClasses ev tvs theta cls tys where (tvs, theta, cls, tys) = tcSplitDFunTy (ctEvPred ev) go ct = pprPanic "makeSuperClasses" (ppr ct) mkStrictSuperClasses :: CtEvidence -> [TyVar] -> ThetaType -- These two args are non-empty only when taking -- superclasses of a /quantified/ constraint -> Class -> [Type] -> TcS [Ct] -- Return constraints for the strict superclasses of -- ev :: forall as. theta => cls tys mkStrictSuperClasses ev tvs theta cls tys = mk_strict_superclasses (unitNameSet (className cls)) ev tvs theta cls tys mk_strict_superclasses :: NameSet -> CtEvidence -> [TyVar] -> ThetaType -> 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 (CtGiven { ctev_evar = evar, ctev_loc = loc }) tvs theta cls tys = concatMapM (do_one_given (mk_given_loc loc)) $ classSCSelIds cls where dict_ids = mkTemplateLocals theta size = sizeTypes tys do_one_given given_loc sel_id | isUnliftedType sc_pred , not (null tvs && null theta) = -- See Note [Equality superclasses in quantified constraints] return [] | otherwise = do { given_ev <- newGivenEvVar given_loc $ mk_given_desc sel_id sc_pred ; mk_superclasses rec_clss given_ev tvs theta sc_pred } where sc_pred = classMethodInstTy sel_id tys -- See Note [Nested quantified constraint superclasses] mk_given_desc :: Id -> PredType -> (PredType, EvTerm) mk_given_desc sel_id sc_pred = (swizzled_pred, swizzled_evterm) where (sc_tvs, sc_rho) = splitForAllTys sc_pred (sc_theta, sc_inner_pred) = splitFunTys sc_rho all_tvs = tvs `chkAppend` sc_tvs all_theta = theta `chkAppend` (map scaledThing sc_theta) swizzled_pred = mkInfSigmaTy all_tvs all_theta sc_inner_pred -- evar :: forall tvs. theta => cls tys -- sel_id :: forall cls_tvs. cls cls_tvs -- -> forall sc_tvs. sc_theta => sc_inner_pred -- swizzled_evterm :: forall tvs sc_tvs. theta => sc_theta => sc_inner_pred swizzled_evterm = EvExpr $ mkLams all_tvs $ mkLams dict_ids $ Var sel_id `mkTyApps` tys `App` (evId evar `mkVarApps` (tvs ++ dict_ids)) `mkVarApps` sc_tvs 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. -- #10335 | GivenOrigin skol_info <- ctLocOrigin loc -- See Note [Solving superclass constraints] in GHC.Tc.TyCl.Instance -- 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 mk_strict_superclasses rec_clss ev tvs theta cls tys | 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. = ASSERT2( null tvs && null theta, ppr tvs $$ ppr theta ) concatMapM do_one_derived (immSuperClasses cls tys) where loc = ctEvLoc ev do_one_derived sc_pred = do { sc_ev <- newDerivedNC loc sc_pred ; mk_superclasses rec_clss sc_ev [] [] sc_pred } {- Note [Improvement from Ground Wanteds] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Suppose class C b a => D a b and consider [W] D Int Bool Is there any point in emitting [D] C Bool Int? No! The only point of emitting superclass constraints for W/D constraints is to get improvement, extra unifications that result from functional dependencies. See Note [Why adding superclasses can help] above. But no variables means no improvement; case closed. -} mk_superclasses :: NameSet -> CtEvidence -> [TyVar] -> ThetaType -> PredType -> TcS [Ct] -- Return this constraint, plus its superclasses, if any mk_superclasses rec_clss ev tvs theta pred | ClassPred cls tys <- classifyPredType pred = mk_superclasses_of rec_clss ev tvs theta cls tys | otherwise -- Superclass is not a class predicate = return [mkNonCanonical ev] mk_superclasses_of :: NameSet -> CtEvidence -> [TyVar] -> ThetaType -> Class -> [Type] -> TcS [Ct] -- Always return this class constraint, -- and expand its superclasses mk_superclasses_of rec_clss ev tvs theta 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 tvs theta 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 | null tvs, null theta = 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 | otherwise = CQuantCan (QCI { qci_tvs = tvs, qci_pred = mkClassPred cls tys , qci_ev = ev , qci_pend_sc = loop_found }) {- Note [Equality superclasses in quantified constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider (#15359, #15593, #15625) f :: (forall a. theta => a ~ b) => stuff It's a bit odd to have a local, quantified constraint for `(a~b)`, but some people want such a thing (see the tickets). And for Coercible it is definitely useful f :: forall m. (forall p q. Coercible p q => Coercible (m p) (m q))) => stuff Moreover it's not hard to arrange; we just need to look up /equality/ constraints in the quantified-constraint environment, which we do in GHC.Tc.Solver.Interact.doTopReactOther. There is a wrinkle though, in the case where 'theta' is empty, so we have f :: (forall a. a~b) => stuff Now, potentially, the superclass machinery kicks in, in makeSuperClasses, giving us a a second quantified constraint (forall a. a ~# b) BUT this is an unboxed value! And nothing has prepared us for dictionary "functions" that are unboxed. Actually it does just about work, but the simplifier ends up with stuff like case (/\a. eq_sel d) of df -> ...(df @Int)... and fails to simplify that any further. And it doesn't satisfy isPredTy any more. So for now we simply decline to take superclasses in the quantified case. Instead we have a special case in GHC.Tc.Solver.Interact.doTopReactOther, which looks for primitive equalities specially in the quantified constraints. See also Note [Evidence for quantified constraints] in GHC.Core.Predicate. ************************************************************************ * * * Irreducibles canonicalization * * ************************************************************************ -} canIrred :: CtIrredStatus -> CtEvidence -> TcS (StopOrContinue Ct) -- Precondition: ty not a tuple and no other evidence form canIrred status ev = do { let pred = ctEvPred ev ; 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 status new_ev } } {- ********************************************************************* * * * Quantified predicates * * ********************************************************************* -} {- Note [Quantified constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The -XQuantifiedConstraints extension allows type-class contexts like this: data Rose f x = Rose x (f (Rose f x)) instance (Eq a, forall b. Eq b => Eq (f b)) => Eq (Rose f a) where (Rose x1 rs1) == (Rose x2 rs2) = x1==x2 && rs1 == rs2 Note the (forall b. Eq b => Eq (f b)) in the instance contexts. This quantified constraint is needed to solve the [W] (Eq (f (Rose f x))) constraint which arises form the (==) definition. The wiki page is https://gitlab.haskell.org/ghc/ghc/wikis/quantified-constraints which in turn contains a link to the GHC Proposal where the change is specified, and a Haskell Symposium paper about it. We implement two main extensions to the design in the paper: 1. We allow a variable in the instance head, e.g. f :: forall m a. (forall b. m b) => D (m a) Notice the 'm' in the head of the quantified constraint, not a class. 2. We support superclasses to quantified constraints. For example (contrived): f :: (Ord b, forall b. Ord b => Ord (m b)) => m a -> m a -> Bool f x y = x==y Here we need (Eq (m a)); but the quantified constraint deals only with Ord. But we can make it work by using its superclass. Here are the moving parts * Language extension {-# LANGUAGE QuantifiedConstraints #-} and add it to ghc-boot-th:GHC.LanguageExtensions.Type.Extension * A new form of evidence, EvDFun, that is used to discharge such wanted constraints * checkValidType gets some changes to accept forall-constraints only in the right places. * Predicate.Pred gets a new constructor ForAllPred, and and classifyPredType analyses a PredType to decompose the new forall-constraints * GHC.Tc.Solver.Monad.InertCans gets an extra field, inert_insts, which holds all the Given forall-constraints. In effect, such Given constraints are like local instance decls. * When trying to solve a class constraint, via GHC.Tc.Solver.Interact.matchInstEnv, use the InstEnv from inert_insts so that we include the local Given forall-constraints in the lookup. (See GHC.Tc.Solver.Monad.getInstEnvs.) * GHC.Tc.Solver.Canonical.canForAll deals with solving a forall-constraint. See Note [Solving a Wanted forall-constraint] * We augment the kick-out code to kick out an inert forall constraint if it can be rewritten by a new type equality; see GHC.Tc.Solver.Monad.kick_out_rewritable Note that a quantified constraint is never /inferred/ (by GHC.Tc.Solver.simplifyInfer). A function can only have a quantified constraint in its type if it is given an explicit type signature. -} canForAllNC :: CtEvidence -> [TyVar] -> TcThetaType -> TcPredType -> TcS (StopOrContinue Ct) canForAllNC ev tvs theta pred | isGiven ev -- See Note [Eagerly expand given superclasses] , Just (cls, tys) <- cls_pred_tys_maybe = do { sc_cts <- mkStrictSuperClasses ev tvs theta cls tys ; emitWork sc_cts ; canForAll ev False } | otherwise = canForAll ev (isJust cls_pred_tys_maybe) where cls_pred_tys_maybe = getClassPredTys_maybe pred canForAll :: CtEvidence -> Bool -> TcS (StopOrContinue Ct) -- We have a constraint (forall as. blah => C tys) canForAll ev pend_sc = do { -- First rewrite it to apply the current substitution -- Do not bother with type-family reductions; we can't -- do them under a forall anyway (c.f. Flatten.flatten_one -- on a forall type) let pred = ctEvPred ev ; (xi,co) <- flatten FM_SubstOnly ev pred -- co :: xi ~ pred ; rewriteEvidence ev xi co `andWhenContinue` \ new_ev -> do { -- Now decompose into its pieces and solve it -- (It takes a lot less code to flatten before decomposing.) ; case classifyPredType (ctEvPred new_ev) of ForAllPred tvs theta pred -> solveForAll new_ev tvs theta pred pend_sc _ -> pprPanic "canForAll" (ppr new_ev) } } solveForAll :: CtEvidence -> [TyVar] -> TcThetaType -> PredType -> Bool -> TcS (StopOrContinue Ct) solveForAll ev tvs theta pred pend_sc | CtWanted { ctev_dest = dest } <- ev = -- See Note [Solving a Wanted forall-constraint] do { let skol_info = QuantCtxtSkol empty_subst = mkEmptyTCvSubst $ mkInScopeSet $ tyCoVarsOfTypes (pred:theta) `delVarSetList` tvs ; (subst, skol_tvs) <- tcInstSkolTyVarsX empty_subst tvs ; given_ev_vars <- mapM newEvVar (substTheta subst theta) ; (lvl, (w_id, wanteds)) <- pushLevelNoWorkList (ppr skol_info) $ do { wanted_ev <- newWantedEvVarNC loc $ substTy subst pred ; return ( ctEvEvId wanted_ev , unitBag (mkNonCanonical wanted_ev)) } ; ev_binds <- emitImplicationTcS lvl skol_info skol_tvs given_ev_vars wanteds ; setWantedEvTerm dest $ EvFun { et_tvs = skol_tvs, et_given = given_ev_vars , et_binds = ev_binds, et_body = w_id } ; stopWith ev "Wanted forall-constraint" } | isGiven ev -- See Note [Solving a Given forall-constraint] = do { addInertForAll qci ; stopWith ev "Given forall-constraint" } | otherwise = do { traceTcS "discarding derived forall-constraint" (ppr ev) ; stopWith ev "Derived forall-constraint" } where loc = ctEvLoc ev qci = QCI { qci_ev = ev, qci_tvs = tvs , qci_pred = pred, qci_pend_sc = pend_sc } {- Note [Solving a Wanted forall-constraint] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Solving a wanted forall (quantified) constraint [W] df :: forall ab. (Eq a, Ord b) => C x a b is delightfully easy. Just build an implication constraint forall ab. (g1::Eq a, g2::Ord b) => [W] d :: C x a and discharge df thus: df = /\ab. \g1 g2. let <binds> in d where <binds> is filled in by solving the implication constraint. All the machinery is to hand; there is little to do. Note [Solving a Given forall-constraint] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ For a Given constraint [G] df :: forall ab. (Eq a, Ord b) => C x a b we just add it to TcS's local InstEnv of known instances, via addInertForall. Then, if we look up (C x Int Bool), say, we'll find a match in the InstEnv. ************************************************************************ * * * 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 [Decomposing FunTy] ~~~~~~~~~~~~~~~~~~~~~~~~ can_eq_nc' may attempt to decompose a FunTy that is un-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. So the calls to getRuntimeRep_maybe may fail (returning Nothing). In that case we'll fall through, zonk, and try again. Zonking should fill the variable k, meaning that decomposition will succeed the second time around. Also note that we require the AnonArgFlag to match. This will stop us decomposing (Int -> Bool) ~ (Show a => blah) It's as if we treat (->) and (=>) as different type constructors. -} 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 rdr_env envs ev eq_rel ty1' ps_ty1 ty2 ps_ty2 | Just ty2' <- tcView ty2 = can_eq_nc' flat rdr_env envs 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. -- See Note [Unwrap newtypes first] -- This must be above the TyVarTy case, in order to guarantee (TyEq:N) can_eq_nc' _flat rdr_env envs ev eq_rel ty1 ps_ty1 ty2 ps_ty2 | ReprEq <- eq_rel , Just stuff1 <- tcTopNormaliseNewTypeTF_maybe envs rdr_env ty1 = can_eq_newtype_nc ev NotSwapped ty1 stuff1 ty2 ps_ty2 | ReprEq <- eq_rel , Just stuff2 <- tcTopNormaliseNewTypeTF_maybe envs rdr_env ty2 = can_eq_newtype_nc ev IsSwapped ty2 stuff2 ty1 ps_ty1 -- Then, get rid of casts can_eq_nc' flat _rdr_env _envs ev eq_rel (CastTy ty1 co1) _ ty2 ps_ty2 | not (isTyVarTy ty2) -- See (3) in Note [Equalities with incompatible kinds] = 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) _ | not (isTyVarTy ty1) -- See (3) in Note [Equalities with incompatible kinds] = canEqCast flat ev eq_rel IsSwapped ty2 co2 ty1 ps_ty1 -- 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 (TyVarTy tv1) ps_ty1 ty2 ps_ty2 = canEqTyVar ev eq_rel NotSwapped tv1 ps_ty1 ty2 ps_ty2 can_eq_nc' True _rdr_env _envs ev eq_rel ty1 ps_ty1 (TyVarTy tv2) ps_ty2 = canEqTyVar ev eq_rel IsSwapped tv2 ps_ty2 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 { setEvBindIfWanted ev (evCoercion $ mkReflCo (eqRelRole eq_rel) ty1) ; stopWith ev "Equal LitTy" } -- Decompose FunTy: (s -> t) and (c => t) -- NB: don't decompose (Int -> blah) ~ (Show a => blah) can_eq_nc' _flat _rdr_env _envs ev eq_rel (FunTy { ft_mult = am1, ft_af = af1, ft_arg = ty1a, ft_res = ty1b }) _ (FunTy { ft_mult = am2, ft_af = af2, ft_arg = ty2a, ft_res = ty2b }) _ | af1 == af2 -- Don't decompose (Int -> blah) ~ (Show a => blah) , Just ty1a_rep <- getRuntimeRep_maybe ty1a -- getRutimeRep_maybe: , Just ty1b_rep <- getRuntimeRep_maybe ty1b -- see Note [Decomposing FunTy] , Just ty2a_rep <- getRuntimeRep_maybe ty2a , Just ty2b_rep <- getRuntimeRep_maybe ty2b = canDecomposableTyConAppOK ev eq_rel funTyCon [am1, ty1a_rep, ty1b_rep, ty1a, ty1b] [am2, ty2a_rep, ty2b_rep, ty2a, ty2b] -- Decompose type constructor applications -- NB: e have expanded type synonyms already can_eq_nc' _flat _rdr_env _envs ev eq_rel (TyConApp tc1 tys1) _ (TyConApp tc2 tys2) _ | 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 _ | NomEq <- eq_rel , Just (t2, s2) <- tcSplitAppTy_maybe ty2 = can_eq_app ev t1 s1 t2 s2 can_eq_nc' True _rdr_env _envs ev eq_rel ty1 _ (AppTy t2 s2) _ | NomEq <- eq_rel , Just (t1, s1) <- tcSplitAppTy_maybe ty1 = can_eq_app ev 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 ; new_ev <- rewriteEqEvidence ev NotSwapped xi1 xi2 co1 co2 ; 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) ; case eq_rel of -- See Note [Unsolved equalities] ReprEq -> continueWith (mkIrredCt OtherCIS ev) NomEq -> continueWith (mkIrredCt InsolubleCIS ev) } -- No need to call canEqFailure/canEqHardFailure because they -- flatten, and the types involved here are already flat {- Note [Unsolved equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ If we have an unsolved equality like (a b ~R# Int) that is not necessarily insoluble! Maybe 'a' will turn out to be a newtype. So we want to make it a potentially-soluble Irred not an insoluble one. Missing this point is what caused #15431 -} --------------------------------- 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 (#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) = tcSplitForAllVarBndrs s1 (bndrs2, phi2) = tcSplitForAllVarBndrs 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 :: [TcTyVar] -> TCvSubst -> [TyVarBinder] -> TcS (TcCoercion, Cts) go (skol_tv:skol_tvs) subst (bndr2:bndrs2) = do { let tv2 = binderVar bndr2 ; (kind_co, wanteds1) <- unify loc Nominal (tyVarKind skol_tv) (substTy subst (tyVarKind tv2)) ; let subst' = extendTvSubstAndInScope subst tv2 (mkCastTy (mkTyVarTy skol_tv) kind_co) -- skol_tv is already in the in-scope set, but the -- free vars of kind_co are not; hence "...AndInScope" ; (co, wanteds2) <- go skol_tvs subst' bndrs2 ; return ( mkTcForAllCo skol_tv kind_co co , wanteds1 `unionBags` wanteds2 ) } -- Done: unify phi1 ~ phi2 go [] subst bndrs2 = ASSERT( null bndrs2 ) unify loc (eqRelRole eq_rel) phi1' (substTyUnchecked subst phi2) go _ _ _ = panic "cna_eq_nc_forall" -- case (s:ss) [] empty_subst2 = mkEmptyTCvSubst (getTCvInScope subst1) ; (lvl, (all_co, wanteds)) <- pushLevelNoWorkList (ppr skol_info) $ go skol_tvs empty_subst2 bndrs2 ; emitTvImplicationTcS lvl skol_info skol_tvs wanteds ; 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" } where unify :: CtLoc -> Role -> TcType -> TcType -> TcS (TcCoercion, Cts) -- This version returns the wanted constraint rather -- than putting it in the work list unify loc role ty1 ty2 | ty1 `tcEqType` ty2 = return (mkTcReflCo role ty1, emptyBag) | otherwise = do { (wanted, co) <- newWantedEq loc role ty1 ty2 ; return (co, unitBag (mkNonCanonical wanted)) } --------------------------------- -- | 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 (Scaled w1 arg1, res1) <- split1 , Just (Scaled w2 arg2, res2) <- split2 , eqType w1 w2 = do { res_a <- go arg1 arg2 ; res_b <- go res1 res2 ; return $ combine_rev (mkVisFunTy w1) 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) <- repSplitTyConApp_maybe ty1 , Just (tc2, tys2) <- repSplitTyConApp_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 #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) {- See Note [Unwrap newtypes first] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider newtype N m a = MkN (m a) Then N will get a conservative, Nominal role for its second parameter 'a', because it appears as an argument to the unknown 'm'. Now consider [W] N Maybe a ~R# N Maybe b If we decompose, we'll get [W] a ~N# b But if instead we unwrap we'll get [W] Maybe a ~R# Maybe b which in turn gives us [W] a ~R# b which is easier to satisfy. Bottom line: unwrap newtypes before decomposing them! c.f. #9123 comment:52,53 for a compelling example. 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 -- Next, we record uses of newtype constructors, since coercing -- through newtypes is tantamount to using their constructors. ; addUsedGREs gre_list -- If a newtype constructor was imported, don't warn about not -- importing it... ; traverse_ keepAlive $ map gre_name gre_list -- ...and similarly, if a newtype constructor was defined in the same -- module, don't warn about it being unused. -- See Note [Tracking unused binding and imports] in GHC.Tc.Utils. ; new_ev <- rewriteEqEvidence ev swapped ty1' ps_ty2 (mkTcSymCo co) (mkTcReflCo Representational ps_ty2) ; can_eq_nc False new_ev ReprEq ty1' ty1' ty2 ps_ty2 } where gre_list = bagToList gres --------- -- ^ Decompose a type application. -- All input types must be flat. See Note [Canonicalising type applications] -- Nominal equality only! can_eq_app :: CtEvidence -- :: s1 t1 ~N s2 t2 -> 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 s1 t1 s2 t2 | CtDerived {} <- ev = do { unifyDeriveds loc [Nominal, Nominal] [s1, t1] [s2, t2] ; stopWith ev "Decomposed [D] AppTy" } | CtWanted { ctev_dest = dest } <- 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" } -- If there is a ForAll/(->) mismatch, the use of the Left coercion -- below is ill-typed, potentially leading to a panic in splitTyConApp -- Test case: typecheck/should_run/Typeable1 -- We could also include this mismatch check above (for W and D), but it's slow -- and we'll get a better error message not doing it | s1k `mismatches` s2k = canEqHardFailure ev (s1 `mkAppTy` t1) (s2 `mkAppTy` t2) | CtGiven { ctev_evar = evar } <- 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 } where loc = ctEvLoc ev s1k = tcTypeKind s1 s2k = tcTypeKind s2 k1 `mismatches` k2 = isForAllTy k1 && not (isForAllTy k2) || not (isForAllTy k1) && isForAllTy k2 ----------------------- -- | 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 ]) ; new_ev <- rewriteEqEvidence ev swapped ty1 ps_ty2 (mkTcGReflRightCo role ty1 co1) (mkTcReflCo role ps_ty2) ; 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] -- Neither tc1 nor tc2 is a saturated funTyCon canTyConApp ev eq_rel tc1 tys1 tc2 tys2 | tc1 == tc2 , tys1 `equalLength` tys2 = do { inerts <- getTcSInerts ; if can_decompose inerts then canDecomposableTyConAppOK ev eq_rel tc1 tys1 tys2 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 OtherCIS 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 -- Reconstruct the types for error messages. This would do -- the wrong thing (from a pretty printing point of view) -- for functions, because we've lost the AnonArgFlag; but -- in fact we never call canTyConApp on a saturated FunTyCon 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 GHC.Core.TyCo.Rep, 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 GHC.Tc.Solver.Interact. 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 (StopOrContinue Ct) -- Precondition: tys1 and tys2 are the same length, hence "OK" canDecomposableTyConAppOK ev eq_rel tc tys1 tys2 = ASSERT( tys1 `equalLength` tys2 ) do { traceTcS "canDecomposableTyConAppOK" (ppr ev $$ ppr eq_rel $$ ppr tc $$ ppr tys1 $$ ppr tys2) ; case ev of CtDerived {} -> unifyDeriveds loc tc_roles tys1 tys2 CtWanted { ctev_dest = dest } -- new_locs and tc_roles are both infinite, so -- we are guaranteed that cos has the same length -- as tys1 and tys2 -> 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 r i ev_co ) | (r, ty1, ty2, i) <- zip4 tc_roles tys1 tys2 [0..] , r /= Phantom , not (isCoercionTy ty1) && not (isCoercionTy ty2) ] ; emitWorkNC given_evs } ; stopWith ev "Decomposed TyConApp" } where loc = ctEvLoc ev role = eqRelRole eq_rel -- infinite, as tyConRolesX returns an infinite tail of Nominal tc_roles = tyConRolesX role tc -- Add nuances to the location during decomposition: -- * if the argument is a kind argument, remember this, so that error -- messages say "kind", not "type". This is determined based on whether -- the corresponding tyConBinder is named (that is, dependent) -- * if the argument is invisible, note this as well, again by -- looking at the corresponding binder -- For oversaturated tycons, we need the (repeat loc) tail, which doesn't -- do either of these changes. (Forgetting to do so led to #16188) -- -- NB: infinite in length new_locs = [ new_loc | bndr <- tyConBinders tc , let new_loc0 | isNamedTyConBinder bndr = toKindLoc loc | otherwise = loc new_loc | isVisibleTyConBinder bndr = updateCtLocOrigin new_loc0 toInvisibleOrigin | otherwise = new_loc0 ] ++ 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 ] ; new_ev <- rewriteEqEvidence ev NotSwapped xi1 xi2 co1 co2 ; continueWith (mkIrredCt OtherCIS 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 { traceTcS "canEqHardFailure" (ppr ty1 $$ ppr ty2) ; (s1, co1) <- flatten FM_SubstOnly ev ty1 ; (s2, co2) <- flatten FM_SubstOnly ev ty2 ; new_ev <- rewriteEqEvidence ev NotSwapped s1 s2 co1 co2 ; continueWith (mkIrredCt InsolubleCIS 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 GHC.Tc.Solver.Monad. And if we don't do this there is a bad danger that GHC.Tc.Solver.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. Note [canCFunEqCan] ~~~~~~~~~~~~~~~~~~~ Flattening the arguments to a type family can change the kind of the type family application. As an easy example, consider (Any k) where (k ~ Type) is in the inert set. The original (Any k :: k) becomes (Any Type :: Type). The problem here is that the fsk in the CFunEqCan will have the old kind. The solution is to come up with a new fsk/fmv of the right kind. For givens, this is easy: just introduce a new fsk and update the flat-cache with the new one. For wanteds, we want to solve the old one if favor of the new one, so we use dischargeFmv. This also kicks out constraints from the inert set; this behavior is correct, as the kind-change may allow more constraints to be solved. We use `isTcReflexiveCo`, to ensure that we only use the hetero-kinded case if we really need to. Of course `flattenArgsNom` should return `Refl` whenever possible, but #15577 was an infinite loop because even though the coercion was homo-kinded, `kind_co` was not `Refl`, so we made a new (identical) CFunEqCan, and then the entire process repeated. -} 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, kind_co) <- flattenArgsNom ev fn tys -- cos :: tys' ~ tys ; let lhs_co = mkTcTyConAppCo Nominal fn cos -- :: F tys' ~ F tys new_lhs = mkTyConApp fn tys' flav = ctEvFlavour ev ; (ev', fsk') <- if isTcReflexiveCo kind_co -- See Note [canCFunEqCan] then do { traceTcS "canCFunEqCan: refl" (ppr new_lhs) ; let fsk_ty = mkTyVarTy fsk ; ev' <- rewriteEqEvidence ev NotSwapped new_lhs fsk_ty lhs_co (mkTcNomReflCo fsk_ty) ; return (ev', fsk) } else do { traceTcS "canCFunEqCan: non-refl" $ vcat [ text "Kind co:" <+> ppr kind_co , text "RHS:" <+> ppr fsk <+> dcolon <+> ppr (tyVarKind fsk) , text "LHS:" <+> hang (ppr (mkTyConApp fn tys)) 2 (dcolon <+> ppr (tcTypeKind (mkTyConApp fn tys))) , text "New LHS" <+> hang (ppr new_lhs) 2 (dcolon <+> ppr (tcTypeKind new_lhs)) ] ; (ev', new_co, new_fsk) <- newFlattenSkolem flav (ctEvLoc ev) fn tys' ; let xi = mkTyVarTy new_fsk `mkCastTy` kind_co -- sym lhs_co :: F tys ~ F tys' -- new_co :: F tys' ~ new_fsk -- co :: F tys ~ (new_fsk |> kind_co) co = mkTcSymCo lhs_co `mkTcTransCo` mkTcCoherenceRightCo Nominal (mkTyVarTy new_fsk) kind_co new_co ; traceTcS "Discharging fmv/fsk due to hetero flattening" (ppr ev) ; dischargeFunEq ev fsk co xi ; return (ev', new_fsk) } ; extendFlatCache fn tys' (ctEvCoercion ev', mkTyVarTy fsk', ctEvFlavour ev') ; continueWith (CFunEqCan { cc_ev = ev', cc_fun = fn , cc_tyargs = tys', cc_fsk = fsk' }) } --------------------- canEqTyVar :: CtEvidence -- ev :: lhs ~ rhs -> EqRel -> SwapFlag -> TcTyVar -- tv1 -> TcType -- lhs: pretty lhs, already flat -> TcType -> TcType -- rhs: already flat -> TcS (StopOrContinue Ct) canEqTyVar ev eq_rel swapped tv1 ps_xi1 xi2 ps_xi2 | k1 `tcEqType` k2 = canEqTyVarHomo ev eq_rel swapped tv1 ps_xi1 xi2 ps_xi2 | otherwise = canEqTyVarHetero ev eq_rel swapped tv1 ps_xi1 k1 xi2 ps_xi2 k2 where k1 = tyVarKind tv1 k2 = tcTypeKind xi2 canEqTyVarHetero :: CtEvidence -- :: (tv1 :: ki1) ~ (xi2 :: ki2) -> EqRel -> SwapFlag -> TcTyVar -> TcType -- tv1, pretty tv1 -> TcKind -- ki1 -> TcType -> TcType -- xi2, pretty xi2 :: ki2 -> TcKind -- ki2 -> TcS (StopOrContinue Ct) canEqTyVarHetero ev eq_rel swapped tv1 ps_tv1 ki1 xi2 ps_xi2 ki2 -- See Note [Equalities with incompatible kinds] = do { kind_co <- emit_kind_co -- :: ki2 ~N ki1 ; let -- kind_co :: (ki2 :: *) ~N (ki1 :: *) (whether swapped or not) -- co1 :: kind(tv1) ~N ki1 rhs' = xi2 `mkCastTy` kind_co -- :: ki1 ps_rhs' = ps_xi2 `mkCastTy` kind_co -- :: ki1 rhs_co = mkTcGReflLeftCo role xi2 kind_co -- rhs_co :: (xi2 |> kind_co) ~ xi2 lhs' = mkTyVarTy tv1 -- same as old lhs lhs_co = mkTcReflCo role lhs' ; traceTcS "Hetero equality gives rise to kind equality" (ppr kind_co <+> dcolon <+> sep [ ppr ki2, text "~#", ppr ki1 ]) ; type_ev <- rewriteEqEvidence ev swapped lhs' rhs' lhs_co rhs_co -- rewriteEqEvidence carries out the swap, so we're NotSwapped any more ; canEqTyVarHomo type_ev eq_rel NotSwapped tv1 ps_tv1 rhs' ps_rhs' } where emit_kind_co :: TcS CoercionN emit_kind_co | CtGiven { ctev_evar = evar } <- ev = do { let kind_co = maybe_sym $ mkTcKindCo (mkTcCoVarCo evar) -- :: k2 ~ k1 ; kind_ev <- newGivenEvVar kind_loc (kind_pty, evCoercion kind_co) ; emitWorkNC [kind_ev] ; return (ctEvCoercion kind_ev) } | otherwise = unifyWanted kind_loc Nominal ki2 ki1 loc = ctev_loc ev role = eqRelRole eq_rel kind_loc = mkKindLoc (mkTyVarTy tv1) xi2 loc kind_pty = mkHeteroPrimEqPred liftedTypeKind liftedTypeKind ki2 ki1 maybe_sym = case swapped of IsSwapped -> id -- if the input is swapped, then we already -- will have k2 ~ k1 NotSwapped -> mkTcSymCo -- guaranteed that tcTypeKind lhs == tcTypeKind rhs canEqTyVarHomo :: CtEvidence -> EqRel -> SwapFlag -> TcTyVar -- lhs: tv1 -> TcType -- pretty lhs, flat -> TcType -> TcType -- rhs, flat -> TcS (StopOrContinue Ct) canEqTyVarHomo ev eq_rel swapped tv1 ps_xi1 xi2 _ | Just (tv2, _) <- tcGetCastedTyVar_maybe xi2 , tv1 == tv2 = canEqReflexive ev eq_rel (mkTyVarTy tv1) -- we don't need to check co because it must be reflexive -- this guarantees (TyEq:TV) | Just (tv2, co2) <- tcGetCastedTyVar_maybe xi2 , swapOverTyVars (isGiven ev) tv1 tv2 = do { traceTcS "canEqTyVar swapOver" (ppr tv1 $$ ppr tv2 $$ ppr swapped) ; let role = eqRelRole eq_rel sym_co2 = mkTcSymCo co2 ty1 = mkTyVarTy tv1 new_lhs = ty1 `mkCastTy` sym_co2 lhs_co = mkTcGReflLeftCo role ty1 sym_co2 new_rhs = mkTyVarTy tv2 rhs_co = mkTcGReflRightCo role new_rhs co2 ; new_ev <- rewriteEqEvidence ev swapped new_lhs new_rhs lhs_co rhs_co ; dflags <- getDynFlags ; canEqTyVar2 dflags new_ev eq_rel IsSwapped tv2 (ps_xi1 `mkCastTy` sym_co2) } canEqTyVarHomo ev eq_rel swapped tv1 _ _ ps_xi2 = do { dflags <- getDynFlags ; canEqTyVar2 dflags ev eq_rel swapped tv1 ps_xi2 } -- The RHS here is either not a casted tyvar, or it's a tyvar but we want -- to rewrite the LHS to the RHS (as per swapOverTyVars) canEqTyVar2 :: DynFlags -> CtEvidence -- lhs ~ rhs (or, if swapped, orhs ~ olhs) -> EqRel -> SwapFlag -> TcTyVar -- lhs = tv, flat -> TcType -- rhs, flat -> TcS (StopOrContinue Ct) -- LHS is an inert type variable, -- and RHS is fully rewritten, but with type synonyms -- preserved as much as possible -- guaranteed that tyVarKind lhs == typeKind rhs, for (TyEq:K) -- the "flat" requirement guarantees (TyEq:AFF) -- (TyEq:N) is checked in can_eq_nc', and (TyEq:TV) is handled in canEqTyVarHomo canEqTyVar2 dflags ev eq_rel swapped tv1 rhs -- this next line checks also for coercion holes; see -- Note [Equalities with incompatible kinds] | MTVU_OK rhs' <- mtvu -- No occurs check -- Must do the occurs check even on tyvar/tyvar -- equalities, in case have x ~ (y :: ..x...) -- #12593 -- guarantees (TyEq:OC), (TyEq:F), and (TyEq:H) = do { new_ev <- rewriteEqEvidence ev swapped lhs rhs' rewrite_co1 rewrite_co2 ; continueWith (CTyEqCan { cc_ev = new_ev, cc_tyvar = tv1 , cc_rhs = rhs', cc_eq_rel = eq_rel }) } | otherwise -- For some reason (occurs check, or forall) we can't unify -- We must not use it for further rewriting! = do { traceTcS "canEqTyVar2 can't unify" (ppr tv1 $$ ppr rhs) ; new_ev <- rewriteEqEvidence ev swapped lhs rhs rewrite_co1 rewrite_co2 ; let status | isInsolubleOccursCheck eq_rel tv1 rhs = InsolubleCIS -- 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 | MTVU_HoleBlocker <- mtvu = BlockedCIS -- This is the case detailed in -- Note [Equalities with incompatible kinds] | otherwise = OtherCIS -- 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. ; continueWith (mkIrredCt status new_ev) } where mtvu = metaTyVarUpdateOK dflags tv1 rhs role = eqRelRole eq_rel lhs = mkTyVarTy tv1 rewrite_co1 = mkTcReflCo role lhs rewrite_co2 = mkTcReflCo role rhs -- | 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 [Equalities with incompatible kinds] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ What do we do when we have an equality (tv :: k1) ~ (rhs :: k2) where k1 and k2 differ? Easy: we create a coercion that relates k1 and k2 and use this to cast. To wit, from [X] (tv :: k1) ~ (rhs :: k2) we go to [noDerived X] co :: k2 ~ k1 [X] (tv :: k1) ~ ((rhs |> co) :: k1) where noDerived G = G noDerived _ = W Wrinkles: (1) The noDerived step is because Derived equalities have no evidence. And yet we absolutely need evidence to be able to proceed here. Given evidence will use the KindCo coercion; Wanted evidence will be a coercion hole. Even a Derived hetero equality begets a Wanted kind equality. (2) Though it would be sound to do so, we must not mark the rewritten Wanted [W] (tv :: k1) ~ ((rhs |> co) :: k1) as canonical in the inert set. In particular, we must not unify tv. If we did, the Wanted becomes a Given (effectively), and then can rewrite other Wanteds. But that's bad: See Note [Wanteds to not rewrite Wanteds] in GHC.Tc.Types.Constraint. The problem is about poor error messages. See #11198 for tales of destruction. So, we have an invariant on CTyEqCan (TyEq:H) that the RHS does not have any coercion holes. This is checked in metaTyVarUpdateOK. We also must be sure to kick out any constraints that mention coercion holes when those holes get filled in. (2a) We don't want to do this for CoercionHoles that witness CFunEqCans (that are produced by the flattener), as these will disappear once we unflatten. So we remember in the CoercionHole structure whether the presence of the hole should block substitution or not. A bit gross, this. (2b) We must now absolutely make sure to kick out any constraints that mention a newly-filled-in coercion hole. This is done in kickOutAfterFillingCoercionHole. (3) Suppose we have [W] (a :: k1) ~ (rhs :: k2). We duly follow the algorithm detailed here, producing [W] co :: k2 ~ k1, and adding [W] (a :: k1) ~ ((rhs |> co) :: k1) to the irreducibles. Some time later, we solve co, and fill in co's coercion hole. This kicks out the irreducible as described in (2b). But now, during canonicalization, we see the cast and remove it, in canEqCast. By the time we get into canEqTyVar, the equality is heterogeneous again, and the process repeats. To avoid this, we don't strip casts off a type if the other type in the equality is a tyvar. And this is an improvement regardless: because tyvars can, generally, unify with casted types, there's no reason to go through the work of stripping off the cast when the cast appears opposite a tyvar. This is implemented in the cast case of can_eq_nc'. (4) Reporting an error for a constraint that is blocked only because of wrinkle (2) is hard: what would we say to users? And we don't really need to report, because if a constraint is blocked, then there is unsolved wanted blocking it; that unsolved wanted will be reported. We thus push such errors to the bottom of the queue in the error-reporting code; they should never be printed. (4a) It would seem possible to do this filtering just based on the presence of a blocking coercion hole. However, this is no good, as it suppresses e.g. no-instance-found errors. We thus record a CtIrredStatus in CIrredCan and filter based on this status. This happened in T14584. An alternative approach is to expressly look for *equalities* with blocking coercion holes, but actually recording the blockage in a status field seems nicer. (4b) The error message might be printed with -fdefer-type-errors, so it still must exist. This is the only reason why there is a message at all. Otherwise, we could simply do nothing. Historical note: We used to do this via emitting a Derived kind equality and then parking the heterogeneous equality as irreducible. But this new approach is much more direct. And it doesn't produce duplicate Deriveds (as the old one did). 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_xi1, ps_xi2 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 deriving (Functor) 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 ~ <type of old evidence> -> 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 #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_nosh = si , ctev_loc = loc }) new_pred co = do { mb_new_ev <- newWanted_SI si loc new_pred -- The "_SI" variant ensures that we make a new Wanted -- with the same shadow-info as the existing one -- with the same shadow-info as the existing one (#16735) ; MASSERT( tcCoercionRole co == ctEvRole ev ) ; setWantedEvTerm dest (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 -> TcCoercion -- lhs_co, of type :: nlhs ~ olhs -> TcCoercion -- rhs_co, of type :: nrhs ~ orhs -> TcS 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 = return (old_ev { ctev_pred = new_pred }) | NotSwapped <- swapped , isTcReflCo lhs_co -- See Note [Rewriting with Refl] , isTcReflCo rhs_co = return (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) ; newGivenEvVar loc' (new_pred, new_tm) } | CtWanted { ctev_dest = dest, ctev_nosh = si } <- old_ev = case dest of HoleDest hole -> do { (new_ev, hole_co) <- newWantedEq_SI (ch_blocker hole) si loc' (ctEvRole old_ev) nlhs nrhs -- The "_SI" variant ensures that we make a new Wanted -- with the same shadow-info as the existing one (#16735) ; 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]) ; return new_ev } _ -> panic "rewriteEqEvidence" #if __GLASGOW_HASKELL__ <= 810 | otherwise = panic "rewriteEvidence" #endif 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 (tcTypeKind ty1) (tcTypeKind 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 _ w1 s1 t1) (FunTy _ w2 s2 t2) = do { co_s <- unifyWanted loc role s1 s2 ; co_t <- unifyWanted loc role t1 t2 ; co_w <- unifyWanted loc Nominal w1 w2 ; return (mkFunCo role co_w 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 _ w1 s1 t1) (FunTy _ w2 s2 t2) = do { unify_derived loc role s1 s2 ; unify_derived loc role t1 t2 ; unify_derived loc Nominal w1 w2 } 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