{-# LANGUAGE CPP #-} module TcFlatten( FlattenEnv(..), FlattenMode(..), mkFlattenEnv, flatten, flattenMany, flatten_many, flattenFamApp, flattenTyVarOuter, unflatten, eqCanRewrite, eqCanRewriteFR, canRewriteOrSame, CtFlavourRole, ctEvFlavourRole, ctFlavourRole ) where #include "HsVersions.h" import TcRnTypes import TcType import Type import TcEvidence import TyCon import TypeRep import Kind( isSubKind ) import Coercion ( tyConRolesX ) import Var import VarEnv import NameEnv import Outputable import VarSet import TcSMonad as TcS import DynFlags( DynFlags ) import Util import Bag import FastString import Control.Monad( when, liftM ) import MonadUtils ( zipWithAndUnzipM ) import GHC.Exts ( inline ) {- Note [The flattening story] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~ * A CFunEqCan is either of form [G] <F xis> : F xis ~ fsk -- fsk is a FlatSkol [W] x : F xis ~ fmv -- fmv is a unification variable, -- but untouchable, -- with MetaInfo = FlatMetaTv where x is the witness variable fsk/fmv is a flatten skolem xis are function-free CFunEqCans are always [Wanted], or [Given], never [Derived] fmv untouchable just means that in a CTyVarEq, say, fmv ~ Int we do NOT unify fmv. * KEY INSIGHTS: - A given flatten-skolem, fsk, is known a-priori to be equal to F xis (the LHS), with <F xis> evidence - A unification flatten-skolem, fmv, stands for the as-yet-unknown type to which (F xis) will eventually reduce * Inert set invariant: if F xis1 ~ fsk1, F xis2 ~ fsk2 then xis1 /= xis2 i.e. at most one CFunEqCan with a particular LHS * Each canonical CFunEqCan x : F xis ~ fsk/fmv has its own distinct evidence variable x and flatten-skolem fsk/fmv. Why? We make a fresh fsk/fmv when the constraint is born; and we never rewrite the RHS of a CFunEqCan. * Function applications can occur in the RHS of a CTyEqCan. No reason not allow this, and it reduces the amount of flattening that must occur. * Flattening a type (F xis): - If we are flattening in a Wanted/Derived constraint then create new [W] x : F xis ~ fmv else create new [G] x : F xis ~ fsk with fresh evidence variable x and flatten-skolem fsk/fmv - Add it to the work list - Replace (F xis) with fsk/fmv in the type you are flattening - You can also add the CFunEqCan to the "flat cache", which simply keeps track of all the function applications you have flattened. - If (F xis) is in the cache already, just use its fsk/fmv and evidence x, and emit nothing. - No need to substitute in the flat-cache. It's not the end of the world if we start with, say (F alpha ~ fmv1) and (F Int ~ fmv2) and then find alpha := Int. Athat will simply give rise to fmv1 := fmv2 via [Interacting rule] below * Canonicalising a CFunEqCan [G/W] x : F xis ~ fsk/fmv - Flatten xis (to substitute any tyvars; there are already no functions) cos :: xis ~ flat_xis - New wanted x2 :: F flat_xis ~ fsk/fmv - Add new wanted to flat cache - Discharge x = F cos ; x2 * Unification flatten-skolems, fmv, ONLY get unified when either a) The CFunEqCan takes a step, using an axiom b) During un-flattening They are never unified in any other form of equality. For example [W] ffmv ~ Int is stuck; it does not unify with fmv. * We *never* substitute in the RHS (i.e. the fsk/fmv) of a CFunEqCan. That would destroy the invariant about the shape of a CFunEqCan, and it would risk wanted/wanted interactions. The only way we learn information about fsk is when the CFunEqCan takes a step. However we *do* substitute in the LHS of a CFunEqCan (else it would never get to fire!) * [Interacting rule] (inert) [W] x1 : F tys ~ fmv1 (work item) [W] x2 : F tys ~ fmv2 Just solve one from the other: x2 := x1 fmv2 := fmv1 This just unites the two fsks into one. Always solve given from wanted if poss. * [Firing rule: wanteds] (work item) [W] x : F tys ~ fmv instantiate axiom: ax_co : F tys ~ rhs Dischard fmv: fmv := alpha x := ax_co ; sym x2 [W] x2 : alpha ~ rhs (Non-canonical) discharging the work item. This is the way that fmv's get unified; even though they are "untouchable". NB: this deals with the case where fmv appears in xi, which can happen; it just happens through the non-canonical stuff Possible short cut (shortCutReduction) if rhs = G rhs_tys, where G is a type function. Then - Flatten rhs_tys (cos : rhs_tys ~ rhs_xis) - Add G rhs_xis ~ fmv to flat cache - New wanted [W] x2 : G rhs_xis ~ fmv - Discharge x := co ; G cos ; x2 * [Firing rule: givens] (work item) [G] g : F tys ~ fsk instantiate axiom: co : F tys ~ rhs Now add non-canonical (since rhs is not flat) [G] (sym g ; co) : fsk ~ rhs Short cut (shortCutReduction) for when rhs = G rhs_tys and G is a type function [G] (co ; g) : G tys ~ fsk But need to flatten tys: flat_cos : tys ~ flat_tys [G] (sym (G flat_cos) ; co ; g) : G flat_tys ~ fsk Why given-fsks, alone, doesn't work ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Could we get away with only flatten meta-tyvars, with no flatten-skolems? No. [W] w : alpha ~ [F alpha Int] ---> flatten w = ...w'... [W] w' : alpha ~ [fsk] [G] <F alpha Int> : F alpha Int ~ fsk --> unify (no occurs check) alpha := [fsk] But since fsk = F alpha Int, this is really an occurs check error. If that is all we know about alpha, we will succeed in constraint solving, producing a program with an infinite type. Even if we did finally get (g : fsk ~ Boo)l by solving (F alpha Int ~ fsk) using axiom, zonking would not see it, so (x::alpha) sitting in the tree will get zonked to an infinite type. (Zonking always only does refl stuff.) Why flatten-meta-vars, alone doesn't work ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Look at Simple13, with unification-fmvs only [G] g : a ~ [F a] ---> Flatten given g' = g;[x] [G] g' : a ~ [fmv] [W] x : F a ~ fmv --> subst a in x x = F g' ; x2 [W] x2 : F [fmv] ~ fmv And now we have an evidence cycle between g' and x! If we used a given instead (ie current story) [G] g : a ~ [F a] ---> Flatten given g' = g;[x] [G] g' : a ~ [fsk] [G] <F a> : F a ~ fsk ---> Substitute for a [G] g' : a ~ [fsk] [G] F (sym g'); <F a> : F [fsk] ~ fsk Why is it right to treat fmv's differently to ordinary unification vars? ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ f :: forall a. a -> a -> Bool g :: F Int -> F Int -> Bool Consider f (x:Int) (y:Bool) This gives alpha~Int, alpha~Bool. There is an inconsistency, but really only one error. SherLoc may tell you which location is most likely, based on other occurrences of alpha. Consider g (x:Int) (y:Bool) Here we get (F Int ~ Int, F Int ~ Bool), which flattens to (fmv ~ Int, fmv ~ Bool) But there are really TWO separate errors. We must not complain about Int~Bool. Moreover these two errors could arise in entirely unrelated parts of the code. (In the alpha case, there must be *some* connection (eg v:alpha in common envt).) Note [Orient equalities with flatten-meta-vars on the left] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ This example comes from IndTypesPerfMerge From the ambiguity check for f :: (F a ~ a) => a we get: [G] F a ~ a [W] F alpha ~ alpha, alpha ~ a From Givens we get [G] F a ~ fsk, fsk ~ a Now if we flatten we get [W] alpha ~ fmv, F alpha ~ fmv, alpha ~ a Now, processing the first one first, choosing alpha := fmv [W] F fmv ~ fmv, fmv ~ a And now we are stuck. We must either *unify* fmv := a, or use the fmv ~ a to rewrite F fmv ~ fmv, so we can make it meet up with the given F a ~ blah. Solution: always put fmvs on the left, so we get [W] fmv ~ alpha, F alpha ~ fmv, alpha ~ a The point is that fmvs are very uninformative, so doing alpha := fmv is a bad idea. We want to use other constraints on alpha first. Note [Derived constraints from wanted CTyEqCans] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Is this type ambiguous: (Foo e ~ Maybe e) => Foo e (indexed-types/should_fail/T4093a) [G] Foo e ~ Maybe e [W] Foo e ~ Foo ee -- ee is a unification variable [W] Foo ee ~ Maybe ee) --- [G] Foo e ~ fsk [G] fsk ~ Maybe e [W] Foo e ~ fmv1 [W] Foo ee ~ fmv2 [W] fmv1 ~ fmv2 [W] fmv2 ~ Maybe ee ---> fmv1 := fsk by matching LHSs [W] Foo ee ~ fmv2 [W] fsk ~ fmv2 [W] fmv2 ~ Maybe ee ---> [W] Foo ee ~ fmv2 [W] fmv2 ~ Maybe e [W] fmv2 ~ Maybe ee Now maybe we shuld get [D] e ~ ee, and then we'd solve it entirely. But if in a smilar situation we got [D] Int ~ Bool we'd be back to complaining about wanted/wanted interactions. Maybe this arises also for fundeps? Here's another example: f :: [a] -> [b] -> blah f (e1 :: F Int) (e2 :: F Int) we get F Int ~ fmv fmv ~ [alpha] fmv ~ [beta] We want: alpha := beta (which might unlock something else). If we generated [D] [alpha] ~ [beta] we'd be good here. Current story: we don't generate these derived constraints. We could, but we'd want to make them very weak, so we didn't get the Int~Bool complaint. ************************************************************************ * * * Other notes (Oct 14) I have not revisted these, but I didn't want to discard them * * ************************************************************************ Try: rewrite wanted with wanted only for fmvs (not all meta-tyvars) But: fmv ~ alpha[0] alpha[0] ~ fmv’ Now we don’t see that fmv ~ fmv’, which is a problem for injectivity detection. Conclusion: rewrite wanteds with wanted for all untouchables. skol ~ untch, must re-orieint to untch ~ skol, so that we can use it to rewrite. ************************************************************************ * * * Examples Here is a long series of examples I had to work through * * ************************************************************************ Simple20 ~~~~~~~~ axiom F [a] = [F a] [G] F [a] ~ a --> [G] fsk ~ a [G] [F a] ~ fsk (nc) --> [G] F a ~ fsk2 [G] fsk ~ [fsk2] [G] fsk ~ a --> [G] F a ~ fsk2 [G] a ~ [fsk2] [G] fsk ~ a ----------------------------------- ---------------------------------------- indexed-types/should_compile/T44984 [W] H (F Bool) ~ H alpha [W] alpha ~ F Bool --> F Bool ~ fmv0 H fmv0 ~ fmv1 H alpha ~ fmv2 fmv1 ~ fmv2 fmv0 ~ alpha flatten ~~~~~~~ fmv0 := F Bool fmv1 := H (F Bool) fmv2 := H alpha alpha := F Bool plus fmv1 ~ fmv2 But these two are equal under the above assumptions. Solve by Refl. --- under plan B, namely solve fmv1:=fmv2 eagerly --- [W] H (F Bool) ~ H alpha [W] alpha ~ F Bool --> F Bool ~ fmv0 H fmv0 ~ fmv1 H alpha ~ fmv2 fmv1 ~ fmv2 fmv0 ~ alpha --> F Bool ~ fmv0 H fmv0 ~ fmv1 H alpha ~ fmv2 fmv2 := fmv1 fmv0 ~ alpha flatten fmv0 := F Bool fmv1 := H fmv0 = H (F Bool) retain H alpha ~ fmv2 because fmv2 has been filled alpha := F Bool ---------------------------- indexed-types/should_failt/T4179 after solving [W] fmv_1 ~ fmv_2 [W] A3 (FCon x) ~ fmv_1 (CFunEqCan) [W] A3 (x (aoa -> fmv_2)) ~ fmv_2 (CFunEqCan) ---------------------------------------- indexed-types/should_fail/T7729a a) [W] BasePrimMonad (Rand m) ~ m1 b) [W] tt m1 ~ BasePrimMonad (Rand m) ---> process (b) first BasePrimMonad (Ramd m) ~ fmv_atH fmv_atH ~ tt m1 ---> now process (a) m1 ~ s_atH ~ tt m1 -- An obscure occurs check ---------------------------------------- typecheck/TcTypeNatSimple Original constraint [W] x + y ~ x + alpha (non-canonical) ==> [W] x + y ~ fmv1 (CFunEqCan) [W] x + alpha ~ fmv2 (CFuneqCan) [W] fmv1 ~ fmv2 (CTyEqCan) (sigh) ---------------------------------------- indexed-types/should_fail/GADTwrong1 [G] Const a ~ () ==> flatten [G] fsk ~ () work item: Const a ~ fsk ==> fire top rule [G] fsk ~ () work item fsk ~ () Surely the work item should rewrite to () ~ ()? Well, maybe not; it'a very special case. More generally, our givens look like F a ~ Int, where (F a) is not reducible. ---------------------------------------- indexed_types/should_fail/T8227: Why using a different can-rewrite rule in CFunEqCan heads does not work. Assuming NOT rewriting wanteds with wanteds Inert: [W] fsk_aBh ~ fmv_aBk -> fmv_aBk [W] fmv_aBk ~ fsk_aBh [G] Scalar fsk_aBg ~ fsk_aBh [G] V a ~ f_aBg Worklist includes [W] Scalar fmv_aBi ~ fmv_aBk fmv_aBi, fmv_aBk are flatten unificaiton variables Work item: [W] V fsk_aBh ~ fmv_aBi Note that the inert wanteds are cyclic, because we do not rewrite wanteds with wanteds. Then we go into a loop when normalise the work-item, because we use rewriteOrSame on the argument of V. Conclusion: Don't make canRewrite context specific; instead use [W] a ~ ty to rewrite a wanted iff 'a' is a unification variable. ---------------------------------------- Here is a somewhat similar case: type family G a :: * blah :: (G a ~ Bool, Eq (G a)) => a -> a blah = error "urk" foo x = blah x For foo we get [W] Eq (G a), G a ~ Bool Flattening [W] G a ~ fmv, Eq fmv, fmv ~ Bool We can't simplify away the Eq Bool unless we substitute for fmv. Maybe that doesn't matter: we would still be left with unsolved G a ~ Bool. -------------------------- Trac #9318 has a very simple program leading to [W] F Int ~ Int [W] F Int ~ Bool We don't want to get "Error Int~Bool". But if fmv's can rewrite wanteds, we will [W] fmv ~ Int [W] fmv ~ Bool ---> [W] Int ~ Bool ************************************************************************ * * * The main flattening functions * * ************************************************************************ Note [Flattening] ~~~~~~~~~~~~~~~~~~~~ flatten ty ==> (xi, cc) where xi has no type functions, unless they appear under ForAlls cc = Auxiliary given (equality) constraints constraining the fresh type variables in xi. Evidence for these is always the identity coercion, because internally the fresh flattening skolem variables are actually identified with the types they have been generated to stand in for. Note that it is flatten's job to flatten *every type function it sees*. flatten is only called on *arguments* to type functions, by canEqGiven. Recall that in comments we use alpha[flat = ty] to represent a flattening skolem variable alpha which has been generated to stand in for ty. ----- Example of flattening a constraint: ------ flatten (List (F (G Int))) ==> (xi, cc) where xi = List alpha cc = { G Int ~ beta[flat = G Int], F beta ~ alpha[flat = F beta] } Here * alpha and beta are 'flattening skolem variables'. * All the constraints in cc are 'given', and all their coercion terms are the identity. NB: Flattening Skolems only occur in canonical constraints, which are never zonked, so we don't need to worry about zonking doing accidental unflattening. Note that we prefer to leave type synonyms unexpanded when possible, so when the flattener encounters one, it first asks whether its transitive expansion contains any type function applications. If so, it expands the synonym and proceeds; if not, it simply returns the unexpanded synonym. Note [Flattener EqRels] ~~~~~~~~~~~~~~~~~~~~~~~ When flattening, we need to know which equality relation -- nominal or representation -- we should be respecting. The only difference is that we rewrite variables by representational equalities when fe_eq_rel is ReprEq. -} data FlattenEnv = FE { fe_mode :: FlattenMode , fe_loc :: CtLoc , fe_flavour :: CtFlavour , fe_eq_rel :: EqRel } -- See Note [Flattener EqRels] data FlattenMode -- Postcondition for all three: inert wrt the type substitution = FM_FlattenAll -- Postcondition: function-free | FM_Avoid TcTyVar Bool -- See Note [Lazy flattening] -- Postcondition: -- * tyvar is only mentioned in result under a rigid path -- e.g. [a] is ok, but F a won't happen -- * If flat_top is True, top level is not a function application -- (but under type constructors is ok e.g. [F a]) | FM_SubstOnly -- See Note [Flattening under a forall] mkFlattenEnv :: FlattenMode -> CtEvidence -> FlattenEnv mkFlattenEnv fm ctev = FE { fe_mode = fm , fe_loc = ctEvLoc ctev , fe_flavour = ctEvFlavour ctev , fe_eq_rel = ctEvEqRel ctev } feRole :: FlattenEnv -> Role feRole = eqRelRole . fe_eq_rel {- Note [Lazy flattening] ~~~~~~~~~~~~~~~~~~~~~~ The idea of FM_Avoid mode is to flatten less aggressively. If we have a ~ [F Int] there seems to be no great merit in lifting out (F Int). But if it was a ~ [G a Int] then we *do* want to lift it out, in case (G a Int) reduces to Bool, say, which gets rid of the occurs-check problem. (For the flat_top Bool, see comments above and at call sites.) HOWEVER, the lazy flattening actually seems to make type inference go *slower*, not faster. perf/compiler/T3064 is a case in point; it gets *dramatically* worse with FM_Avoid. I think it may be because floating the types out means we normalise them, and that often makes them smaller and perhaps allows more re-use of previously solved goals. But to be honest I'm not absolutely certain, so I am leaving FM_Avoid in the code base. What I'm removing is the unique place where it is *used*, namely in TcCanonical.canEqTyVar. See also Note [Conservative unification check] in TcUnify, which gives other examples where lazy flattening caused problems. Bottom line: FM_Avoid is unused for now (Nov 14). Note: T5321Fun got faster when I disabled FM_Avoid T5837 did too, but it's pathalogical anyway Note [Phantoms in the flattener] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Suppose we have data Proxy p = Proxy and we're flattening (Proxy ty) w.r.t. ReprEq. Then, we know that `ty` is really irrelevant -- it will be ignored when solving for representational equality later on. So, we omit flattening `ty` entirely. This may violate the expectation of "xi"s for a bit, but the canonicaliser will soon throw out the phantoms when decomposing a TyConApp. (Or, the canonicaliser will emit an insoluble, in which case the unflattened version yields a better error message anyway.) Note [flatten_many performance] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In programs with lots of type-level evaluation, flatten_many becomes part of a tight loop. For example, see test perf/compiler/T9872a, which calls flatten_many a whopping 7,106,808 times. It is thus important that flatten_many be efficient. Performance testing showed that the current implementation is indeed efficient. It's critically important that zipWithAndUnzipM be specialized to TcS, and it's also quite helpful to actually `inline` it. On test T9872a, here are the allocation stats (Dec 16, 2014): * Unspecialized, uninlined: 8,472,613,440 bytes allocated in the heap * Specialized, uninlined: 6,639,253,488 bytes allocated in the heap * Specialized, inlined: 6,281,539,792 bytes allocated in the heap To improve performance even further, flatten_many_nom is split off from flatten_many, as nominal equality is the common case. This would be natural to write using mapAndUnzipM, but even inlined, that function is not as performant as a hand-written loop. * mapAndUnzipM, inlined: 7,463,047,432 bytes allocated in the heap * hand-written recursion: 5,848,602,848 bytes allocated in the heap If you make any change here, pay close attention to the T9872{a,b,c} tests and T5321Fun. If we need to make this yet more performant, a possible way forward is to duplicate the flattener code for the nominal case, and make that case faster. This doesn't seem quite worth it, yet. -} ------------------ flatten :: FlattenMode -> CtEvidence -> TcType -> TcS (Xi, TcCoercion) flatten mode ev ty = runFlatten (flatten_one fmode ty) where fmode = mkFlattenEnv mode ev flattenMany :: FlattenMode -> CtEvidence -> [Role] -> [TcType] -> TcS ([Xi], [TcCoercion]) -- Flatten a bunch of types all at once. Roles on the coercions returned -- always match the corresponding roles passed in. flattenMany mode ev roles tys = runFlatten (flatten_many fmode roles tys) where fmode = mkFlattenEnv mode ev flattenFamApp :: FlattenMode -> CtEvidence -> TyCon -> [TcType] -> TcS (Xi, TcCoercion) flattenFamApp mode ev tc tys = runFlatten (flatten_fam_app fmode tc tys) where fmode = mkFlattenEnv mode ev ------------------ flatten_many :: FlattenEnv -> [Role] -> [Type] -> TcS ([Xi], [TcCoercion]) -- Coercions :: Xi ~ Type, at roles given -- Returns True iff (no flattening happened) -- NB: The EvVar inside the 'fe_ev :: CtEvidence' is unused, -- we merely want (a) Given/Solved/Derived/Wanted info -- (b) the GivenLoc/WantedLoc for when we create new evidence flatten_many fmode roles tys -- See Note [flatten_many performance] = inline zipWithAndUnzipM go roles tys where go Nominal ty = flatten_one (setFEEqRel fmode NomEq) ty go Representational ty = flatten_one (setFEEqRel fmode ReprEq) ty go Phantom ty = -- See Note [Phantoms in the flattener] return (ty, mkTcPhantomCo ty ty) -- | Like 'flatten_many', but assumes that every role is nominal. flatten_many_nom :: FlattenEnv -> [Type] -> TcS ([Xi], [TcCoercion]) flatten_many_nom _ [] = return ([], []) -- See Note [flatten_many performance] flatten_many_nom fmode (ty:tys) = ASSERT( fe_eq_rel fmode == NomEq ) do { (xi, co) <- flatten_one fmode ty ; (xis, cos) <- flatten_many_nom fmode tys ; return (xi:xis, co:cos) } ------------------ flatten_one :: FlattenEnv -> TcType -> TcS (Xi, TcCoercion) -- Flatten a type to get rid of type function applications, returning -- the new type-function-free type, and a collection of new equality -- constraints. See Note [Flattening] for more detail. -- -- Postcondition: Coercion :: Xi ~ TcType -- The role on the result coercion matches the EqRel in the FlattenEnv flatten_one fmode xi@(LitTy {}) = return (xi, mkTcReflCo (feRole fmode) xi) flatten_one fmode (TyVarTy tv) = flattenTyVar fmode tv flatten_one fmode (AppTy ty1 ty2) = do { (xi1,co1) <- flatten_one fmode ty1 ; case (fe_eq_rel fmode, nextRole xi1) of (NomEq, _) -> flatten_rhs xi1 co1 NomEq (ReprEq, Nominal) -> flatten_rhs xi1 co1 NomEq (ReprEq, Representational) -> flatten_rhs xi1 co1 ReprEq (ReprEq, Phantom) -> return (mkAppTy xi1 ty2, co1 `mkTcAppCo` mkTcNomReflCo ty2) } where flatten_rhs xi1 co1 eq_rel2 = do { (xi2,co2) <- flatten_one (setFEEqRel fmode eq_rel2) ty2 ; traceTcS "flatten/appty" (ppr ty1 $$ ppr ty2 $$ ppr xi1 $$ ppr co1 $$ ppr xi2 $$ ppr co2) ; let role1 = feRole fmode role2 = eqRelRole eq_rel2 ; return ( mkAppTy xi1 xi2 , mkTcTransAppCo role1 co1 xi1 ty1 role2 co2 xi2 ty2 role1 ) } -- output should match fmode flatten_one fmode (FunTy ty1 ty2) = do { (xi1,co1) <- flatten_one fmode ty1 ; (xi2,co2) <- flatten_one fmode ty2 ; return (mkFunTy xi1 xi2, mkTcFunCo (feRole fmode) co1 co2) } flatten_one fmode (TyConApp tc tys) -- Expand type synonyms that mention type families -- on the RHS; see Note [Flattening synonyms] | Just (tenv, rhs, tys') <- tcExpandTyCon_maybe tc tys , let expanded_ty = mkAppTys (substTy (mkTopTvSubst tenv) rhs) tys' = case fe_mode fmode of FM_FlattenAll | anyNameEnv isTypeFamilyTyCon (tyConsOfType rhs) -> flatten_one fmode expanded_ty | otherwise -> flattenTyConApp fmode tc tys _ -> flattenTyConApp fmode tc tys -- Otherwise, it's a type function application, and we have to -- flatten it away as well, and generate a new given equality constraint -- between the application and a newly generated flattening skolem variable. | isTypeFamilyTyCon tc = flatten_fam_app fmode tc tys -- For * a normal data type application -- * data family application -- we just recursively flatten the arguments. | otherwise -- FM_Avoid stuff commented out; see Note [Lazy flattening] -- , let fmode' = case fmode of -- Switch off the flat_top bit in FM_Avoid -- FE { fe_mode = FM_Avoid tv _ } -- -> fmode { fe_mode = FM_Avoid tv False } -- _ -> fmode = flattenTyConApp fmode tc tys flatten_one fmode ty@(ForAllTy {}) -- We allow for-alls when, but only when, no type function -- applications inside the forall involve the bound type variables. = do { let (tvs, rho) = splitForAllTys ty ; (rho', co) <- flatten_one (setFEMode fmode FM_SubstOnly) rho -- Substitute only under a forall -- See Note [Flattening under a forall] ; return (mkForAllTys tvs rho', foldr mkTcForAllCo co tvs) } flattenTyConApp :: FlattenEnv -> TyCon -> [TcType] -> TcS (Xi, TcCoercion) flattenTyConApp fmode tc tys = do { (xis, cos) <- case fe_eq_rel fmode of NomEq -> flatten_many_nom fmode tys ReprEq -> flatten_many fmode (tyConRolesX role tc) tys ; return (mkTyConApp tc xis, mkTcTyConAppCo role tc cos) } where role = feRole fmode {- Note [Flattening synonyms] ~~~~~~~~~~~~~~~~~~~~~~~~~~ Not expanding synonyms aggressively improves error messages, and keeps types smaller. But we need to take care. Suppose type T a = a -> a and we want to flatten the type (T (F a)). Then we can safely flatten the (F a) to a skolem, and return (T fsk). We don't need to expand the synonym. This works because TcTyConAppCo can deal with synonyms (unlike TyConAppCo), see Note [TcCoercions] in TcEvidence. But (Trac #8979) for type T a = (F a, a) where F is a type function we must expand the synonym in (say) T Int, to expose the type function to the flattener. Note [Flattening under a forall] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Under a forall, we (a) MUST apply the inert substitution (b) MUST NOT flatten type family applications Hence FMSubstOnly. For (a) consider c ~ a, a ~ T (forall b. (b, [c])) If we don't apply the c~a substitution to the second constraint we won't see the occurs-check error. For (b) consider (a ~ forall b. F a b), we don't want to flatten to (a ~ forall b.fsk, F a b ~ fsk) because now the 'b' has escaped its scope. We'd have to flatten to (a ~ forall b. fsk b, forall b. F a b ~ fsk b) and we have not begun to think about how to make that work! ************************************************************************ * * Flattening a type-family application * * ************************************************************************ -} flatten_fam_app, flatten_exact_fam_app, flatten_exact_fam_app_fully :: FlattenEnv -> TyCon -> [TcType] -> TcS (Xi, TcCoercion) -- flatten_fam_app can be over-saturated -- flatten_exact_fam_app is exactly saturated -- flatten_exact_fam_app_fully lifts out the application to top level -- Postcondition: Coercion :: Xi ~ F tys flatten_fam_app fmode tc tys -- Can be over-saturated = ASSERT( tyConArity tc <= length tys ) -- Type functions are saturated -- The type function might be *over* saturated -- in which case the remaining arguments should -- be dealt with by AppTys do { let (tys1, tys_rest) = splitAt (tyConArity tc) tys ; (xi1, co1) <- flatten_exact_fam_app fmode tc tys1 -- co1 :: xi1 ~ F tys1 -- all Nominal roles b/c the tycon is oversaturated ; (xis_rest, cos_rest) <- flatten_many fmode (repeat Nominal) tys_rest -- cos_res :: xis_rest ~ tys_rest ; return ( mkAppTys xi1 xis_rest -- NB mkAppTys: rhs_xi might not be a type variable -- cf Trac #5655 , mkTcAppCos co1 cos_rest -- (rhs_xi :: F xis) ; (F cos :: F xis ~ F tys) ) } flatten_exact_fam_app fmode tc tys = case fe_mode fmode of FM_FlattenAll -> flatten_exact_fam_app_fully fmode tc tys FM_SubstOnly -> do { (xis, cos) <- flatten_many fmode roles tys ; return ( mkTyConApp tc xis , mkTcTyConAppCo (feRole fmode) tc cos ) } FM_Avoid tv flat_top -> do { (xis, cos) <- flatten_many fmode roles tys ; if flat_top || tv `elemVarSet` tyVarsOfTypes xis then flatten_exact_fam_app_fully fmode tc tys else return ( mkTyConApp tc xis , mkTcTyConAppCo (feRole fmode) tc cos ) } where -- These are always going to be Nominal for now, -- but not if #8177 is implemented roles = tyConRolesX (feRole fmode) tc flatten_exact_fam_app_fully fmode tc tys -- See Note [Reduce type family applications eagerly] = try_to_reduce tc tys False id $ do { (xis, cos) <- flatten_many_nom (setFEEqRel (setFEMode fmode FM_FlattenAll) NomEq) tys ; let ret_co = mkTcTyConAppCo (feRole fmode) tc cos -- ret_co :: F xis ~ F tys ; mb_ct <- lookupFlatCache tc xis ; case mb_ct of Just (co, rhs_ty, flav) -- co :: F xis ~ fsk | (flav, NomEq) `canRewriteOrSameFR` (feFlavourRole fmode) -> -- Usable hit in the flat-cache -- We certainly *can* use a Wanted for a Wanted do { traceTcS "flatten/flat-cache hit" $ (ppr tc <+> ppr xis $$ ppr rhs_ty $$ ppr co) ; (fsk_xi, fsk_co) <- flatten_one fmode rhs_ty -- The fsk may already have been unified, so flatten it -- fsk_co :: fsk_xi ~ fsk ; return (fsk_xi, fsk_co `mkTcTransCo` maybeTcSubCo (fe_eq_rel fmode) (mkTcSymCo co) `mkTcTransCo` ret_co) } -- :: fsk_xi ~ F xis -- Try to reduce the family application right now -- See Note [Reduce type family applications eagerly] _ -> try_to_reduce tc xis True (`mkTcTransCo` ret_co) $ do { let fam_ty = mkTyConApp tc xis ; (ev, fsk) <- newFlattenSkolem (fe_flavour fmode) (fe_loc fmode) fam_ty ; let fsk_ty = mkTyVarTy fsk co = ctEvCoercion ev ; extendFlatCache tc xis (co, fsk_ty, ctEvFlavour ev) -- The new constraint (F xis ~ fsk) is not necessarily inert -- (e.g. the LHS may be a redex) so we must put it in the work list ; let ct = CFunEqCan { cc_ev = ev , cc_fun = tc , cc_tyargs = xis , cc_fsk = fsk } ; emitFlatWork ct ; traceTcS "flatten/flat-cache miss" $ (ppr fam_ty $$ ppr fsk $$ ppr ev) ; return (fsk_ty, maybeTcSubCo (fe_eq_rel fmode) (mkTcSymCo co) `mkTcTransCo` ret_co) } } where try_to_reduce :: TyCon -- F, family tycon -> [Type] -- args, not necessarily flattened -> Bool -- add to the flat cache? -> ( TcCoercion -- :: xi ~ F args -> TcCoercion ) -- what to return from outer function -> TcS (Xi, TcCoercion) -- continuation upon failure -> TcS (Xi, TcCoercion) try_to_reduce tc tys cache update_co k = do { mb_match <- matchFam tc tys ; case mb_match of Just (norm_co, norm_ty) -> do { traceTcS "Eager T.F. reduction success" $ vcat [ppr tc, ppr tys, ppr norm_ty, ppr cache] ; (xi, final_co) <- flatten_one fmode norm_ty ; let co = norm_co `mkTcTransCo` mkTcSymCo final_co ; when cache $ extendFlatCache tc tys (co, xi, fe_flavour fmode) ; return (xi, update_co $ mkTcSymCo co) } Nothing -> k } {- Note [Reduce type family applications eagerly] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ If we come across a type-family application like (Append (Cons x Nil) t), then, rather than flattening to a skolem etc, we may as well just reduce it on the spot to (Cons x t). This saves a lot of intermediate steps. Examples that are helped are tests T9872, and T5321Fun. Performance testing indicates that it's best to try this *twice*, once before flattening arguments and once after flattening arguments. Adding the extra reduction attempt before flattening arguments cut the allocation amounts for the T9872{a,b,c} tests by half. Testing also indicated that the early reduction should not use the flat-cache, but that the later reduction should. It's possible that with more examples, we might learn that these knobs should be set differently. Once we've got a flat rhs, we extend the flatten-cache to record the result. Doing so can save lots of work when the same redex shows up more than once. Note that we record the link from the redex all the way to its *final* value, not just the single step reduction. ************************************************************************ * * Flattening a type variable * * ************************************************************************ Note [The inert equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Definition [Can-rewrite relation] A "can-rewrite" relation between flavours, written f1 >= f2, is a binary relation with the following properties R1. >= is transitive R2. If f1 >= f, and f2 >= f, then either f1 >= f2 or f2 >= f1 Lemma. If f1 >= f then f1 >= f1 Proof. By property (R2), with f1=f2 Definition [Generalised substitution] A "generalised substitution" S is a set of triples (a -f-> t), where a is a type variable t is a type f is a flavour such that (WF1) if (a -f1-> t1) in S (a -f2-> t2) in S then neither (f1 >= f2) nor (f2 >= f1) hold (WF2) if (a -f-> t) is in S, then t /= a Definition [Applying a generalised substitution] If S is a generalised substitution S(f,a) = t, if (a -fs-> t) in S, and fs >= f = a, otherwise Application extends naturally to types S(f,t), modulo roles. See Note [Flavours with roles]. Theorem: S(f,a) is well defined as a function. Proof: Suppose (a -f1-> t1) and (a -f2-> t2) are both in S, and f1 >= f and f2 >= f Then by (R2) f1 >= f2 or f2 >= f1, which contradicts (WF) Notation: repeated application. S^0(f,t) = t S^(n+1)(f,t) = S(f, S^n(t)) Definition: inert generalised substitution A generalised substitution S is "inert" iff (IG1) there is an n such that for every f,t, S^n(f,t) = S^(n+1)(f,t) (IG2) if (b -f-> t) in S, and f >= f, then S(f,t) = t that is, each individual binding is "self-stable" ---------------------------------------------------------------- Our main invariant: the inert CTyEqCans should be an inert generalised substitution ---------------------------------------------------------------- Note that inertness is not the same as idempotence. To apply S to a type, you may have to apply it recursive. But inertness does guarantee that this recursive use will terminate. ---------- The main theorem -------------- Suppose we have a "work item" a -fw-> t and an inert generalised substitution S, such that (T1) S(fw,a) = a -- LHS of work-item is a fixpoint of S(fw,_) (T2) S(fw,t) = t -- RHS of work-item is a fixpoint of S(fw,_) (T3) a not in t -- No occurs check in the work item (K1) if (a -fs-> s) is in S then not (fw >= fs) (K2) if (b -fs-> s) is in S, where b /= a, then (K2a) not (fs >= fs) or (K2b) not (fw >= fs) or (K2c) a not in s (K3) If (b -fs-> s) is in S with (fw >= fs), then (K3a) If the role of fs is nominal: s /= a (K3b) If the role of fs is representational: EITHER a not in s, OR the path from the top of s to a includes at least one non-newtype then the extended substition T = S+(a -fw-> t) is an inert generalised substitution. The idea is that * (T1-2) are guaranteed by exhaustively rewriting the work-item with S(fw,_). * T3 is guaranteed by a simple occurs-check on the work item. * (K1-3) are the "kick-out" criteria. (As stated, they are really the "keep" criteria.) If the current inert S contains a triple that does not satisfy (K1-3), then we remove it from S by "kicking it out", and re-processing it. * Note that kicking out is a Bad Thing, because it means we have to re-process a constraint. The less we kick out, the better. TODO: Make sure that kicking out really *is* a Bad Thing. We've assumed this but haven't done the empirical study to check. * Assume we have G>=G, G>=W, D>=D, and that's all. Then, when performing a unification we add a new given a -G-> ty. But doing so does NOT require us to kick out an inert wanted that mentions a, because of (K2a). This is a common case, hence good not to kick out. * Lemma (L1): The conditions of the Main Theorem imply that there is no (a fs-> t) in S, s.t. (fs >= fw). Proof. Suppose the contrary (fs >= fw). Then because of (T1), S(fw,a)=a. But since fs>=fw, S(fw,a) = s, hence s=a. But now we have (a -fs-> a) in S, which contradicts (WF2). * The extended substitution satisfies (WF1) and (WF2) - (K1) plus (L1) guarantee that the extended substiution satisfies (WF1). - (T3) guarantees (WF2). * (K2) is about inertness. Intuitively, any infinite chain T^0(f,t), T^1(f,t), T^2(f,T).... must pass through the new work item infnitely often, since the substution without the work item is inert; and must pass through at least one of the triples in S infnitely often. - (K2a): if not(fs>=fs) then there is no f that fs can rewrite (fs>=f), and hence this triple never plays a role in application S(f,a). It is always safe to extend S with such a triple. (NB: we could strengten K1) in this way too, but see K3. - (K2b): If this holds, we can't pass through this triple infinitely often, because if we did then fs>=f, fw>=f, hence fs>=fw, contradicting (L1), or fw>=fs contradicting K2b. - (K2c): if a not in s, we hae no further opportunity to apply the work item. NB: this reasoning isn't water tight. Key lemma to make it watertight. Under the conditions of the Main Theorem, forall f st fw >= f, a is not in S^k(f,t), for any k Also, consider roles more carefully. See Note [Flavours with roles]. Completeness ~~~~~~~~~~~~~ K3: completeness. (K3) is not necessary for the extended substitution to be inert. In fact K1 could be made stronger by saying ... then (not (fw >= fs) or not (fs >= fs)) But it's not enough for S to be inert; we also want completeness. That is, we want to be able to solve all soluble wanted equalities. Suppose we have work-item b -G-> a inert-item a -W-> b Assuming (G >= W) but not (W >= W), this fulfills all the conditions, so we could extend the inerts, thus: inert-items b -G-> a a -W-> b But if we kicked-out the inert item, we'd get work-item a -W-> b inert-item b -G-> a Then rewrite the work-item gives us (a -W-> a), which is soluble via Refl. So we add one more clause to the kick-out criteria Another way to understand (K3) is that we treat an inert item a -f-> b in the same way as b -f-> a So if we kick out one, we should kick out the other. The orientation is somewhat accidental. When considering roles, we also need the second clause (K3b). Consider inert-item a -W/R-> b c work-item c -G/N-> a The work-item doesn't get rewritten by the inert, because (>=) doesn't hold. We've satisfied conditions (T1)-(T3) and (K1) and (K2). If all we had were condition (K3a), then we would keep the inert around and add the work item. But then, consider if we hit the following: work-item2 b -G/N-> Id where newtype Id x = Id x For similar reasons, if we only had (K3a), we wouldn't kick the representational inert out. And then, we'd miss solving the inert, which now reduced to reflexivity. The solution here is to kick out representational inerts whenever the tyvar of a work item is "exposed", where exposed means not under some proper data-type constructor, like [] or Maybe. See isTyVarExposed in TcType. This is encoded in (K3b). Note [Flavours with roles] ~~~~~~~~~~~~~~~~~~~~~~~~~~ The system described in Note [The inert equalities] discusses an abstract set of flavours. In GHC, flavours have two components: the flavour proper, taken from {Wanted, Derived, Given}; and the equality relation (often called role), taken from {NomEq, ReprEq}. When substituting w.r.t. the inert set, as described in Note [The inert equalities], we must be careful to respect roles. For example, if we have inert set: a -G/R-> Int b -G/R-> Bool type role T nominal representational and we wish to compute S(W/R, T a b), the correct answer is T a Bool, NOT T Int Bool. The reason is that T's first parameter has a nominal role, and thus rewriting a to Int in T a b is wrong. Indeed, this non-congruence of subsitution means that the proof in Note [The inert equalities] may need to be revisited, but we don't think that the end conclusion is wrong. -} flattenTyVar :: FlattenEnv -> TcTyVar -> TcS (Xi, TcCoercion) -- "Flattening" a type variable means to apply the substitution to it -- The substitution is actually the union of -- * the unifications that have taken place (either before the -- solver started, or in TcInteract.solveByUnification) -- * the CTyEqCans held in the inert set -- -- Postcondition: co : xi ~ tv flattenTyVar fmode tv = do { mb_yes <- flattenTyVarOuter fmode tv ; case mb_yes of Left tv' -> -- Done do { traceTcS "flattenTyVar1" (ppr tv $$ ppr (tyVarKind tv')) ; return (ty', mkTcReflCo (feRole fmode) ty') } where ty' = mkTyVarTy tv' Right (ty1, co1) -- Recurse -> do { (ty2, co2) <- flatten_one fmode ty1 ; traceTcS "flattenTyVar3" (ppr tv $$ ppr ty2) ; return (ty2, co2 `mkTcTransCo` co1) } } flattenTyVarOuter :: FlattenEnv -> TcTyVar -> TcS (Either TyVar (TcType, TcCoercion)) -- Look up the tyvar in -- a) the internal MetaTyVar box -- b) the tyvar binds -- c) the inerts -- Return (Left tv') if it is not found, tv' has a properly zonked kind -- (Right (ty, co) if found, with co :: ty ~ tv; flattenTyVarOuter fmode tv | not (isTcTyVar tv) -- Happens when flatten under a (forall a. ty) = Left `liftM` flattenTyVarFinal fmode tv -- So ty contains refernces to the non-TcTyVar a | otherwise = do { mb_ty <- isFilledMetaTyVar_maybe tv ; case mb_ty of { Just ty -> do { traceTcS "Following filled tyvar" (ppr tv <+> equals <+> ppr ty) ; return (Right (ty, mkTcReflCo (feRole fmode) ty)) } ; Nothing -> -- Try in the inert equalities -- See Definition [Applying a generalised substitution] do { ieqs <- getInertEqs ; case lookupVarEnv ieqs tv of Just (ct:_) -- If the first doesn't work, -- the subsequent ones won't either | CTyEqCan { cc_ev = ctev, cc_tyvar = tv, cc_rhs = rhs_ty } <- ct , ctEvFlavourRole ctev `eqCanRewriteFR` feFlavourRole fmode -> do { traceTcS "Following inert tyvar" (ppr tv <+> equals <+> ppr rhs_ty $$ ppr ctev) ; let rewrite_co1 = mkTcSymCo (ctEvCoercion ctev) rewrite_co = case (ctEvEqRel ctev, fe_eq_rel fmode) of (ReprEq, _rel) -> ASSERT( _rel == ReprEq ) -- if this ASSERT fails, then -- eqCanRewriteFR answered incorrectly rewrite_co1 (NomEq, NomEq) -> rewrite_co1 (NomEq, ReprEq) -> mkTcSubCo rewrite_co1 ; return (Right (rhs_ty, rewrite_co)) } -- NB: ct is Derived then fmode must be also, hence -- we are not going to touch the returned coercion -- so ctEvCoercion is fine. _other -> Left `liftM` flattenTyVarFinal fmode tv } } } flattenTyVarFinal :: FlattenEnv -> TcTyVar -> TcS TyVar flattenTyVarFinal fmode tv = -- Done, but make sure the kind is zonked do { let kind = tyVarKind tv kind_fmode = setFEMode fmode FM_SubstOnly ; (new_knd, _kind_co) <- flatten_one kind_fmode kind ; return (setVarType tv new_knd) } {- Note [An alternative story for the inert substitution] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ (This entire note is just background, left here in case we ever want to return the the previousl state of affairs) We used (GHC 7.8) to have this story for the inert substitution inert_eqs * 'a' is not in fvs(ty) * They are *inert* in the weaker sense that there is no infinite chain of (i1 `eqCanRewrite` i2), (i2 `eqCanRewrite` i3), etc This means that flattening must be recursive, but it does allow [G] a ~ [b] [G] b ~ Maybe c This avoids "saturating" the Givens, which can save a modest amount of work. It is easy to implement, in TcInteract.kick_out, by only kicking out an inert only if (a) the work item can rewrite the inert AND (b) the inert cannot rewrite the work item This is signifcantly harder to think about. It can save a LOT of work in occurs-check cases, but we don't care about them much. Trac #5837 is an example; all the constraints here are Givens [G] a ~ TF (a,Int) --> work TF (a,Int) ~ fsk inert fsk ~ a ---> work fsk ~ (TF a, TF Int) inert fsk ~ a ---> work a ~ (TF a, TF Int) inert fsk ~ a ---> (attempting to flatten (TF a) so that it does not mention a work TF a ~ fsk2 inert a ~ (fsk2, TF Int) inert fsk ~ (fsk2, TF Int) ---> (substitute for a) work TF (fsk2, TF Int) ~ fsk2 inert a ~ (fsk2, TF Int) inert fsk ~ (fsk2, TF Int) ---> (top-level reduction, re-orient) work fsk2 ~ (TF fsk2, TF Int) inert a ~ (fsk2, TF Int) inert fsk ~ (fsk2, TF Int) ---> (attempt to flatten (TF fsk2) to get rid of fsk2 work TF fsk2 ~ fsk3 work fsk2 ~ (fsk3, TF Int) inert a ~ (fsk2, TF Int) inert fsk ~ (fsk2, TF Int) ---> work TF fsk2 ~ fsk3 inert fsk2 ~ (fsk3, TF Int) inert a ~ ((fsk3, TF Int), TF Int) inert fsk ~ ((fsk3, TF Int), TF Int) Because the incoming given rewrites all the inert givens, we get more and more duplication in the inert set. But this really only happens in pathalogical casee, so we don't care. -} eqCanRewrite :: CtEvidence -> CtEvidence -> Bool eqCanRewrite ev1 ev2 = ctEvFlavourRole ev1 `eqCanRewriteFR` ctEvFlavourRole ev2 -- | Whether or not one 'Ct' can rewrite another is determined by its -- flavour and its equality relation type CtFlavourRole = (CtFlavour, EqRel) -- | Extract the flavour and role from a 'CtEvidence' ctEvFlavourRole :: CtEvidence -> CtFlavourRole ctEvFlavourRole ev = (ctEvFlavour ev, ctEvEqRel ev) -- | Extract the flavour and role from a 'Ct' ctFlavourRole :: Ct -> CtFlavourRole ctFlavourRole = ctEvFlavourRole . cc_ev -- | Extract the flavour and role from a 'FlattenEnv' feFlavourRole :: FlattenEnv -> CtFlavourRole feFlavourRole (FE { fe_flavour = flav, fe_eq_rel = eq_rel }) = (flav, eq_rel) eqCanRewriteFR :: CtFlavourRole -> CtFlavourRole -> Bool -- Very important function! -- See Note [eqCanRewrite] eqCanRewriteFR (Given, NomEq) (_, _) = True eqCanRewriteFR (Given, ReprEq) (_, ReprEq) = True eqCanRewriteFR _ _ = False canRewriteOrSame :: CtEvidence -> CtEvidence -> Bool -- See Note [canRewriteOrSame] canRewriteOrSame ev1 ev2 = ev1 `eqCanRewrite` ev2 || ctEvFlavourRole ev1 == ctEvFlavourRole ev2 canRewriteOrSameFR :: CtFlavourRole -> CtFlavourRole -> Bool canRewriteOrSameFR fr1 fr2 = fr1 `eqCanRewriteFR` fr2 || fr1 == fr2 {- Note [eqCanRewrite] ~~~~~~~~~~~~~~~~~~~ (eqCanRewrite ct1 ct2) holds if the constraint ct1 (a CTyEqCan of form tv ~ ty) can be used to rewrite ct2. It must satisfy the properties of a can-rewrite relation, see Definition [Can-rewrite relation] At the moment we don't allow Wanteds to rewrite Wanteds, because that can give rise to very confusing type error messages. A good example is Trac #8450. Here's another f :: a -> Bool f x = ( [x,'c'], [x,True] ) `seq` True Here we get [W] a ~ Char [W] a ~ Bool but we do not want to complain about Bool ~ Char! Accordingly, we also don't let Deriveds rewrite Deriveds. With the solver handling Coercible constraints like equality constraints, the rewrite conditions must take role into account, never allowing a representational equality to rewrite a nominal one. Note [canRewriteOrSame] ~~~~~~~~~~~~~~~~~~~~~~~ canRewriteOrSame is similar but * returns True for Wanted/Wanted. * works for all kinds of constraints, not just CTyEqCans See the call sites for explanations. ************************************************************************ * * Unflattening * * ************************************************************************ An unflattening example: [W] F a ~ alpha flattens to [W] F a ~ fmv (CFunEqCan) [W] fmv ~ alpha (CTyEqCan) We must solve both! -} unflatten :: Cts -> Cts -> TcS Cts unflatten tv_eqs funeqs = do { dflags <- getDynFlags ; tclvl <- getTcLevel ; traceTcS "Unflattening" $ braces $ vcat [ ptext (sLit "Funeqs =") <+> pprCts funeqs , ptext (sLit "Tv eqs =") <+> pprCts tv_eqs ] -- Step 1: unflatten the CFunEqCans, except if that causes an occurs check -- See Note [Unflatten using funeqs first] ; funeqs <- foldrBagM (unflatten_funeq dflags) emptyCts funeqs ; traceTcS "Unflattening 1" $ braces (pprCts funeqs) -- Step 2: unify the irreds, if possible ; tv_eqs <- foldrBagM (unflatten_eq dflags tclvl) emptyCts tv_eqs ; traceTcS "Unflattening 2" $ braces (pprCts tv_eqs) -- Step 3: fill any remaining fmvs with fresh unification variables ; funeqs <- mapBagM finalise_funeq funeqs ; traceTcS "Unflattening 3" $ braces (pprCts funeqs) -- Step 4: remove any irreds that look like ty ~ ty ; tv_eqs <- foldrBagM finalise_eq emptyCts tv_eqs ; let all_flat = tv_eqs `andCts` funeqs ; traceTcS "Unflattening done" $ braces (pprCts all_flat) ; return all_flat } where ---------------- unflatten_funeq :: DynFlags -> Ct -> Cts -> TcS Cts unflatten_funeq dflags ct@(CFunEqCan { cc_fun = tc, cc_tyargs = xis , cc_fsk = fmv, cc_ev = ev }) rest = do { -- fmv should be a flatten meta-tv; we now fix its final -- value, and then zonking will eliminate it filled <- tryFill dflags fmv (mkTyConApp tc xis) ev ; return (if filled then rest else ct `consCts` rest) } unflatten_funeq _ other_ct _ = pprPanic "unflatten_funeq" (ppr other_ct) ---------------- finalise_funeq :: Ct -> TcS Ct finalise_funeq (CFunEqCan { cc_fsk = fmv, cc_ev = ev }) = do { demoteUnfilledFmv fmv ; return (mkNonCanonical ev) } finalise_funeq ct = pprPanic "finalise_funeq" (ppr ct) ---------------- unflatten_eq :: DynFlags -> TcLevel -> Ct -> Cts -> TcS Cts unflatten_eq dflags tclvl ct@(CTyEqCan { cc_ev = ev, cc_tyvar = tv, cc_rhs = rhs }) rest | isFmvTyVar tv = do { lhs_elim <- tryFill dflags tv rhs ev ; if lhs_elim then return rest else do { rhs_elim <- try_fill dflags tclvl ev rhs (mkTyVarTy tv) ; if rhs_elim then return rest else return (ct `consCts` rest) } } | otherwise = return (ct `consCts` rest) unflatten_eq _ _ ct _ = pprPanic "unflatten_irred" (ppr ct) ---------------- finalise_eq :: Ct -> Cts -> TcS Cts finalise_eq (CTyEqCan { cc_ev = ev, cc_tyvar = tv , cc_rhs = rhs, cc_eq_rel = eq_rel }) rest | isFmvTyVar tv = do { ty1 <- zonkTcTyVar tv ; ty2 <- zonkTcType rhs ; let is_refl = ty1 `tcEqType` ty2 ; if is_refl then do { when (isWanted ev) $ setEvBind (ctEvId ev) (EvCoercion $ mkTcReflCo (eqRelRole eq_rel) rhs) ; return rest } else return (mkNonCanonical ev `consCts` rest) } | otherwise = return (mkNonCanonical ev `consCts` rest) finalise_eq ct _ = pprPanic "finalise_irred" (ppr ct) ---------------- try_fill dflags tclvl ev ty1 ty2 | Just tv1 <- tcGetTyVar_maybe ty1 , isTouchableOrFmv tclvl tv1 , typeKind ty1 `isSubKind` tyVarKind tv1 = tryFill dflags tv1 ty2 ev | otherwise = return False tryFill :: DynFlags -> TcTyVar -> TcType -> CtEvidence -> TcS Bool -- (tryFill tv rhs ev) sees if 'tv' is an un-filled MetaTv -- If so, and if tv does not appear in 'rhs', set tv := rhs -- bind the evidence (which should be a CtWanted) to Refl<rhs> -- and return True. Otherwise return False tryFill dflags tv rhs ev = ASSERT2( not (isGiven ev), ppr ev ) do { is_filled <- isFilledMetaTyVar tv ; if is_filled then return False else do { rhs' <- zonkTcType rhs ; case occurCheckExpand dflags tv rhs' of OC_OK rhs'' -- Normal case: fill the tyvar -> do { when (isWanted ev) $ setEvBind (ctEvId ev) (EvCoercion (mkTcReflCo (ctEvRole ev) rhs'')) ; setWantedTyBind tv rhs'' ; return True } _ -> -- Occurs check return False } } {- Note [Unflatten using funeqs first] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ [W] G a ~ Int [W] F (G a) ~ G a do not want to end up with [W} F Int ~ Int because that might actually hold! Better to end up with the two above unsolved constraints. The flat form will be G a ~ fmv1 (CFunEqCan) F fmv1 ~ fmv2 (CFunEqCan) fmv1 ~ Int (CTyEqCan) fmv1 ~ fmv2 (CTyEqCan) Flatten using the fun-eqs first. -} -- | Change the 'EqRel' in a 'FlattenEnv'. Avoids allocating a -- new 'FlattenEnv' where possible. setFEEqRel :: FlattenEnv -> EqRel -> FlattenEnv setFEEqRel fmode@(FE { fe_eq_rel = old_eq_rel }) new_eq_rel | old_eq_rel == new_eq_rel = fmode | otherwise = fmode { fe_eq_rel = new_eq_rel } -- | Change the 'FlattenMode' in a 'FlattenEnv'. Avoids allocating -- a new 'FlattenEnv' where possible. setFEMode :: FlattenEnv -> FlattenMode -> FlattenEnv setFEMode fmode@(FE { fe_mode = old_mode }) new_mode | old_mode `eq` new_mode = fmode | otherwise = fmode { fe_mode = new_mode } where FM_FlattenAll `eq` FM_FlattenAll = True FM_SubstOnly `eq` FM_SubstOnly = True FM_Avoid tv1 b1 `eq` FM_Avoid tv2 b2 = tv1 == tv2 && b1 == b2 _ `eq` _ = False