module TcCanonical( canonicalize, emitWorkNC, StopOrContinue (..) ) where #include "HsVersions.h" import TcRnTypes import TcType import Type import Kind import TcEvidence import Class import TyCon import TypeRep import Var import VarEnv import OccName( OccName ) import Outputable import Control.Monad ( when ) import DynFlags( DynFlags ) import VarSet import TcSMonad import FastString import Util import BasicTypes import Maybes( catMaybes )\end{code} %************************************************************************ %* * %* The Canonicaliser * %* * %************************************************************************ Note [Canonicalization] ~~~~~~~~~~~~~~~~~~~~~~~ Canonicalization converts a flat constraint to a canonical form. It is unary (i.e. treats individual constraints one at a time), does not do any zonking, but lives in TcS monad because it needs to create fresh variables (for flattening) and consult the inerts (for efficiency). The execution plan for canonicalization is the following: 1) Decomposition of equalities happens as necessary until we reach a variable or type family in one side. There is no decomposition step for other forms of constraints. 2) If, when we decompose, we discover a variable on the head then we look at inert_eqs from the current inert for a substitution for this variable and contine decomposing. Hence we lazily apply the inert substitution if it is needed. 3) If no more decomposition is possible, we deeply apply the substitution from the inert_eqs and continue with flattening. 4) During flattening, we examine whether we have already flattened some function application by looking at all the CTyFunEqs with the same function in the inert set. The reason for deeply applying the inert substitution at step (3) is to maximise our chances of matching an already flattened family application in the inert. The net result is that a constraint coming out of the canonicalization phase cannot be rewritten any further from the inerts (but maybe /it/ can rewrite an inert or still interact with an inert in a further phase in the simplifier. \begin{code}

-- Informative results of canonicalization data StopOrContinue = ContinueWith Ct -- Either no canonicalization happened, or if some did -- happen, it is still safe to just keep going with this -- work item. | Stop -- Some canonicalization happened, extra work is now in -- the TcS WorkList. instance Outputable StopOrContinue where ppr Stop = ptext (sLit "Stop") ppr (ContinueWith w) = ptext (sLit "ContinueWith") <+> ppr w continueWith :: Ct -> TcS StopOrContinue continueWith = return . ContinueWith andWhenContinue :: TcS StopOrContinue -> (Ct -> TcS StopOrContinue) -> TcS StopOrContinue andWhenContinue tcs1 tcs2 = do { r <- tcs1 ; case r of Stop -> return Stop ContinueWith ct -> tcs2 ct }\end{code} Note [Caching for canonicals] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Our plan with pre-canonicalization is to be able to solve a constraint really fast from existing bindings in TcEvBinds. So one may think that the condition (isCNonCanonical) is not necessary. However consider the following setup: InertSet = { [W] d1 : Num t } WorkList = { [W] d2 : Num t, [W] c : t ~ Int} Now, we prioritize equalities, but in our concrete example (should_run/mc17.hs) the first (d2) constraint is dealt with first, because (t ~ Int) is an equality that only later appears in the worklist since it is pulled out from a nested implication constraint. So, let's examine what happens: - We encounter work item (d2 : Num t) - Nothing is yet in EvBinds, so we reach the interaction with inerts and set: d2 := d1 and we discard d2 from the worklist. The inert set remains unaffected. - Now the equation ([W] c : t ~ Int) is encountered and kicks-out (d1 : Num t) from the inerts. Then that equation gets spontaneously solved, perhaps. We end up with: InertSet : { [G] c : t ~ Int } WorkList : { [W] d1 : Num t} - Now we examine (d1), we observe that there is a binding for (Num t) in the evidence binds and we set: d1 := d2 and end up in a loop! Now, the constraints that get kicked out from the inert set are always Canonical, so by restricting the use of the pre-canonicalizer to NonCanonical constraints we eliminate this danger. Moreover, for canonical constraints we already have good caching mechanisms (effectively the interaction solver) and we are interested in reducing things like superclasses of the same non-canonical constraint being generated hence I don't expect us to lose a lot by introducing the (isCNonCanonical) restriction. A similar situation can arise in TcSimplify, at the end of the solve_wanteds function, where constraints from the inert set are returned as new work -- our substCt ensures however that if they are not rewritten by subst, they remain canonical and hence we will not attempt to solve them from the EvBinds. If on the other hand they did get rewritten and are now non-canonical they will still not match the EvBinds, so we are again good. \begin{code}

-- Top-level canonicalization -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ canonicalize :: Ct -> TcS StopOrContinue canonicalize ct@(CNonCanonical { cc_ev = ev }) = do { traceTcS "canonicalize (non-canonical)" (ppr ct) ; {-# SCC "canEvVar" #-} canEvNC ev } canonicalize (CDictCan { cc_ev = ev , cc_class = cls , cc_tyargs = xis }) = {-# SCC "canClass" #-} canClass ev cls xis -- Do not add any superclasses canonicalize (CTyEqCan { cc_ev = ev , cc_tyvar = tv , cc_rhs = xi }) = {-# SCC "canEqLeafTyVarEq" #-} canEqTyVar ev NotSwapped tv xi xi canonicalize (CFunEqCan { cc_ev = ev , cc_fun = fn , cc_tyargs = xis1 , cc_rhs = xi2 }) = {-# SCC "canEqLeafFunEq" #-} canEqLeafFun ev NotSwapped fn xis1 xi2 xi2 canonicalize (CIrredEvCan { cc_ev = ev }) = canIrred ev canonicalize (CHoleCan { cc_ev = ev, cc_occ = occ }) = canHole ev occ canEvNC :: CtEvidence -> TcS StopOrContinue -- Called only for non-canonical EvVars canEvNC ev = case classifyPredType (ctEvPred ev) of ClassPred cls tys -> traceTcS "canEvNC:cls" (ppr cls <+> ppr tys) >> canClassNC ev cls tys EqPred ty1 ty2 -> traceTcS "canEvNC:eq" (ppr ty1 $$ ppr ty2) >> canEqNC ev ty1 ty2 TuplePred tys -> traceTcS "canEvNC:tup" (ppr tys) >> canTuple ev tys IrredPred {} -> traceTcS "canEvNC:irred" (ppr (ctEvPred ev)) >> canIrred ev\end{code} %************************************************************************ %* * %* Tuple Canonicalization %* * %************************************************************************ \begin{code}

canTuple :: CtEvidence -> [PredType] -> TcS StopOrContinue canTuple ev tys = do { traceTcS "can_pred" (text "TuplePred!") ; let xcomp = EvTupleMk xdecomp x = zipWith (\_ i -> EvTupleSel x i) tys [0..] ; ctevs <- xCtEvidence ev (XEvTerm tys xcomp xdecomp) ; canEvVarsCreated ctevs }\end{code} %************************************************************************ %* * %* Class Canonicalization %* * %************************************************************************ \begin{code}

canClass, canClassNC :: CtEvidence -> Class -> [Type] -> TcS StopOrContinue -- Precondition: EvVar is class evidence -- The canClassNC version is used on non-canonical constraints -- and adds superclasses. The plain canClass version is used -- for already-canonical class constraints (but which might have -- been subsituted or somthing), and hence do not need superclasses canClassNC ev cls tys = canClass ev cls tys `andWhenContinue` emitSuperclasses canClass ev cls tys = do { (xis, cos) <- flattenMany FMFullFlatten ev tys ; let co = mkTcTyConAppCo Nominal (classTyCon cls) cos xi = mkClassPred cls xis ; mb <- rewriteEvidence ev xi co ; traceTcS "canClass" (vcat [ ppr ev <+> ppr cls <+> ppr tys , ppr xi, ppr mb ]) ; case mb of Nothing -> return Stop Just new_ev -> continueWith $ CDictCan { cc_ev = new_ev , cc_tyargs = xis, cc_class = cls } } emitSuperclasses :: Ct -> TcS StopOrContinue emitSuperclasses ct@(CDictCan { cc_ev = ev , cc_tyargs = xis_new, cc_class = cls }) -- Add superclasses of this one here, See Note [Adding superclasses]. -- But only if we are not simplifying the LHS of a rule. = do { newSCWorkFromFlavored ev cls xis_new -- Arguably we should "seq" the coercions if they are derived, -- as we do below for emit_kind_constraint, to allow errors in -- superclasses to be executed if deferred to runtime! ; continueWith ct } emitSuperclasses _ = panic "emit_superclasses of non-class!"\end{code} Note [Adding superclasses] ~~~~~~~~~~~~~~~~~~~~~~~~~~ Since dictionaries are canonicalized only once in their lifetime, the place to add their superclasses is canonicalisation (The alternative would be to do it during constraint solving, but we'd have to be extremely careful to not repeatedly introduced the same superclass in our worklist). Here is what we do: For Givens: We add all their superclasses as Givens. For Wanteds: Generally speaking we want to be able to add superclasses of wanteds for two reasons: (1) Oportunities for improvement. Example: class (a ~ b) => C a b Wanted constraint is: C alpha beta We'd like to simply have C alpha alpha. Similar situations arise in relation to functional dependencies. (2) To have minimal constraints to quantify over: For instance, if our wanted constraint is (Eq a, Ord a) we'd only like to quantify over Ord a. To deal with (1) above we only add the superclasses of wanteds which may lead to improvement, that is: equality superclasses or superclasses with functional dependencies. We deal with (2) completely independently in TcSimplify. See Note [Minimize by SuperClasses] in TcSimplify. Moreover, in all cases the extra improvement constraints are Derived. Derived constraints have an identity (for now), but we don't do anything with their evidence. For instance they are never used to rewrite other constraints. See also [New Wanted Superclass Work] in TcInteract. For Deriveds: We do nothing. Here's an example that demonstrates why we chose to NOT add superclasses during simplification: [Comes from ticket #4497] class Num (RealOf t) => Normed t type family RealOf x Assume the generated wanted constraint is: RealOf e ~ e, Normed e If we were to be adding the superclasses during simplification we'd get: Num uf, Normed e, RealOf e ~ e, RealOf e ~ uf ==> e ~ uf, Num uf, Normed e, RealOf e ~ e ==> [Spontaneous solve] Num uf, Normed uf, RealOf uf ~ uf While looks exactly like our original constraint. If we add the superclass again we'd loop. By adding superclasses definitely only once, during canonicalisation, this situation can't happen. \begin{code}

newSCWorkFromFlavored :: CtEvidence -> Class -> [Xi] -> TcS () -- Returns superclasses, see Note [Adding superclasses] newSCWorkFromFlavored flavor cls xis | isDerived flavor = return () -- Deriveds don't yield more superclasses because we will -- add them transitively in the case of wanteds. | isGiven flavor = do { let sc_theta = immSuperClasses cls xis xev_decomp x = zipWith (\_ i -> EvSuperClass x i) sc_theta [0..] xev = XEvTerm { ev_preds = sc_theta , ev_comp = panic "Can't compose for given!" , ev_decomp = xev_decomp } ; ctevs <- xCtEvidence flavor xev ; emitWorkNC ctevs } | isEmptyVarSet (tyVarsOfTypes xis) = return () -- Wanteds with no variables yield no deriveds. -- See Note [Improvement from Ground Wanteds] | otherwise -- Wanted case, just add those SC that can lead to improvement. = do { let sc_rec_theta = transSuperClasses cls xis impr_theta = filter is_improvement_pty sc_rec_theta loc = ctev_loc flavor ; traceTcS "newSCWork/Derived" $ text "impr_theta =" <+> ppr impr_theta ; mb_der_evs <- mapM (newDerived loc) impr_theta ; emitWorkNC (catMaybes mb_der_evs) } is_improvement_pty :: PredType -> Bool -- Either it's an equality, or has some functional dependency is_improvement_pty ty = go (classifyPredType ty) where go (EqPred {}) = True go (ClassPred cls _tys) = not $ null fundeps where (_,fundeps) = classTvsFds cls go (TuplePred ts) = any is_improvement_pty ts go (IrredPred {}) = True -- Might have equalities after reduction?\end{code} %************************************************************************ %* * %* Irreducibles canonicalization %* * %************************************************************************ \begin{code}

canIrred :: CtEvidence -> TcS StopOrContinue -- Precondition: ty not a tuple and no other evidence form canIrred old_ev = do { let old_ty = ctEvPred old_ev ; traceTcS "can_pred" (text "IrredPred = " <+> ppr old_ty) ; (xi,co) <- flatten FMFullFlatten old_ev old_ty -- co :: xi ~ old_ty ; mb <- rewriteEvidence old_ev xi co ; case mb of { Nothing -> return Stop ; Just 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 TuplePred tys -> canTuple new_ev tys EqPred ty1 ty2 -> canEqNC new_ev ty1 ty2 _ -> continueWith $ CIrredEvCan { cc_ev = new_ev } } } } canHole :: CtEvidence -> OccName -> TcS StopOrContinue canHole ev occ = do { let ty = ctEvPred ev ; (xi,co) <- flatten FMFullFlatten ev ty -- co :: xi ~ ty ; mb <- rewriteEvidence ev xi co ; case mb of Just new_ev -> emitInsoluble (CHoleCan { cc_ev = new_ev, cc_occ = occ }) Nothing -> return () -- Found a cached copy; won't happen ; return Stop }\end{code} %************************************************************************ %* * %* Flattening (eliminating all function symbols) * %* * %************************************************************************ 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. \begin{code}

data FlattenMode = FMSubstOnly | FMFullFlatten -- See Note [Flattening under a forall] -- Flatten a bunch of types all at once. flattenMany :: FlattenMode -> CtEvidence -> [Type] -> TcS ([Xi], [TcCoercion]) -- Coercions :: Xi ~ Type -- Returns True iff (no flattening happened) -- NB: The EvVar inside the 'ctxt :: CtEvidence' is unused, -- we merely want (a) Given/Solved/Derived/Wanted info -- (b) the GivenLoc/WantedLoc for when we create new evidence flattenMany f ctxt tys = -- pprTrace "flattenMany" empty $ go tys where go [] = return ([],[]) go (ty:tys) = do { (xi,co) <- flatten f ctxt ty ; (xis,cos) <- go tys ; return (xi:xis,co:cos) } flatten :: FlattenMode -> CtEvidence -> 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 flatten _ _ xi@(LitTy {}) = return (xi, mkTcNomReflCo xi) flatten f ctxt (TyVarTy tv) = flattenTyVar f ctxt tv flatten f ctxt (AppTy ty1 ty2) = do { (xi1,co1) <- flatten f ctxt ty1 ; (xi2,co2) <- flatten f ctxt ty2 ; traceTcS "flatten/appty" (ppr ty1 $$ ppr ty2 $$ ppr xi1 $$ ppr co1 $$ ppr xi2 $$ ppr co2) ; return (mkAppTy xi1 xi2, mkTcAppCo co1 co2) } flatten f ctxt (FunTy ty1 ty2) = do { (xi1,co1) <- flatten f ctxt ty1 ; (xi2,co2) <- flatten f ctxt ty2 ; return (mkFunTy xi1 xi2, mkTcFunCo Nominal co1 co2) } flatten f ctxt (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 , any isSynFamilyTyCon (tyConsOfType rhs) = flatten f ctxt (mkAppTys (substTy (mkTopTvSubst tenv) rhs) tys') -- For * a normal data type application -- * data family application -- * type synonym application whose RHS does not mention type families -- See Note [Flattening synonyms] -- we just recursively flatten the arguments. | not (isSynFamilyTyCon tc) = do { (xis,cos) <- flattenMany f ctxt tys ; return (mkTyConApp tc xis, mkTcTyConAppCo Nominal tc cos) } -- 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. | otherwise = ASSERT( tyConArity tc <= length tys ) -- Type functions are saturated do { (xis, cos) <- flattenMany f ctxt tys ; let (xi_args, xi_rest) = splitAt (tyConArity tc) xis (cos_args, cos_rest) = splitAt (tyConArity tc) cos -- The type function might be *over* saturated -- in which case the remaining arguments should -- be dealt with by AppTys ; (rhs_xi, ret_co) <- flattenNestedFamApp f ctxt tc xi_args -- Emit the flat constraints ; return ( mkAppTys rhs_xi xi_rest -- NB mkAppTys: rhs_xi might not be a type variable -- cf Trac #5655 , mkTcAppCos (mkTcSymCo ret_co `mkTcTransCo` mkTcTyConAppCo Nominal tc cos_args) $ cos_rest ) } flatten _f ctxt 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 FMSubstOnly ctxt rho -- Substitute only under a forall -- See Note [Flattening under a forall] ; return (mkForAllTys tvs rho', foldr mkTcForAllCo co tvs) }\end{code} 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 functoin 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! \begin{code}

flattenNestedFamApp :: FlattenMode -> CtEvidence -> TyCon -> [TcType] -- Exactly-saturated type function application -> TcS (Xi, TcCoercion) flattenNestedFamApp FMSubstOnly _ tc xi_args = do { let fam_ty = mkTyConApp tc xi_args ; return (fam_ty, mkTcNomReflCo fam_ty) } flattenNestedFamApp FMFullFlatten ctxt tc xi_args -- Eactly saturated = do { let fam_ty = mkTyConApp tc xi_args ; mb_ct <- lookupFlatEqn tc xi_args ; case mb_ct of Just (ctev, rhs_ty) | ctev `canRewriteOrSame `ctxt -- Must allow [W]/[W] -> -- You may think that we can just return (cc_rhs ct) but not so. -- return (mkTcCoVarCo (ctId ct), cc_rhs ct, []) -- The cached constraint resides in the cache so we have to flatten -- the rhs to make sure we have applied any inert substitution to it. -- Alternatively we could be applying the inert substitution to the -- cache as well when we interact an equality with the inert. -- The design choice is: do we keep the flat cache rewritten or not? -- For now I say we don't keep it fully rewritten. do { (rhs_xi,co) <- flatten FMFullFlatten ctev rhs_ty ; let final_co = evTermCoercion (ctEvTerm ctev) `mkTcTransCo` mkTcSymCo co ; traceTcS "flatten/flat-cache hit" $ (ppr ctev $$ ppr rhs_xi $$ ppr final_co) ; return (rhs_xi, final_co) } _ -> do { (ctev, rhs_xi) <- newFlattenSkolem ctxt fam_ty ; extendFlatCache tc xi_args ctev rhs_xi -- The new constraint (F xi_args ~ rhs_xi) 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 = ctev , cc_fun = tc , cc_tyargs = xi_args , cc_rhs = rhs_xi } ; updWorkListTcS $ extendWorkListFunEq ct ; traceTcS "flatten/flat-cache miss" $ (ppr fam_ty $$ ppr rhs_xi $$ ppr ctev) ; return (rhs_xi, evTermCoercion (ctEvTerm ctev)) } }\end{code} \begin{code}

flattenTyVar :: FlattenMode -> CtEvidence -> TcTyVar -> TcS (Xi, TcCoercion) -- "Flattening" a type variable means to apply the substitution to it -- The substitution is actually the union of the substitution in the TyBinds -- for the unification variables that have been unified already with the inert -- equalities, see Note [Spontaneously solved in TyBinds] in TcInteract. -- -- Postcondition: co : xi ~ tv flattenTyVar f ctxt tv = do { mb_yes <- flattenTyVarOuter f ctxt tv ; case mb_yes of Left tv' -> -- Done return (ty, mkTcNomReflCo ty) where ty = mkTyVarTy tv' Right (ty1, co1) -> -- Recurse do { (ty2, co2) <- flatten f ctxt ty1 ; return (ty2, co2 `mkTcTransCo` co1) } } flattenTyVarOuter, flattenTyVarFinal :: FlattenMode -> CtEvidence -> 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 -- NB: in the latter case ty is not necessarily flattened flattenTyVarOuter f ctxt tv | not (isTcTyVar tv) -- Happens when flatten under a (forall a. ty) = flattenTyVarFinal f ctxt 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, mkTcNomReflCo ty)) } ; Nothing -> -- Try in ty_binds do { ty_binds <- getTcSTyBindsMap ; case lookupVarEnv ty_binds tv of { Just (_tv,ty) -> do { traceTcS "Following bound tyvar" (ppr tv <+> equals <+> ppr ty) ; return (Right (ty, mkTcNomReflCo ty)) } ; -- NB: ty_binds coercions are all ReflCo, Nothing -> -- Try in the inert equalities do { ieqs <- getInertEqs ; case lookupVarEnv ieqs tv of Just (ct:_) -- If the first doesn't work, | let ctev = ctEvidence ct -- the subsequent ones won't either rhs_ty = cc_rhs ct , ctev `canRewrite` ctxt -> do { traceTcS "Following inert tyvar" (ppr tv <+> equals <+> ppr rhs_ty $$ ppr ctev) ; return (Right (rhs_ty, mkTcSymCo (evTermCoercion (ctEvTerm ctev)))) } -- NB: even if ct is Derived we are not going to -- touch the actual coercion so we are fine. _other -> flattenTyVarFinal f ctxt tv } } } } } flattenTyVarFinal f ctxt tv = -- Done, but make sure the kind is zonked do { let knd = tyVarKind tv ; (new_knd, _kind_co) <- flatten f ctxt knd ; return (Left (setVarType tv new_knd)) }\end{code} Note [Non-idempotent inert substitution] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The inert substitution is not idempotent in the broad sense. It is only idempotent in that it cannot rewrite the RHS of other inert equalities any further. An example of such an inert substitution is: [G] g1 : ta8 ~ ta4 [W] g2 : ta4 ~ a5Fj Observe that the wanted cannot rewrite the solved goal, despite the fact that ta4 appears on an RHS of an equality. Now, imagine a constraint: [W] g3: ta8 ~ Int coming in. If we simply apply once the inert substitution we will get: [W] g3_1: ta4 ~ Int and because potentially ta4 is untouchable we will try to insert g3_1 in the inert set, getting a panic since the inert only allows ONE equation per LHS type variable (as it should). For this reason, when we reach to flatten a type variable, we flatten it recursively, so that we can make sure that the inert substitution /is/ fully applied. Insufficient (non-recursive) rewriting was the reason for #5668. %************************************************************************ %* * %* Equalities %* * %************************************************************************ \begin{code}

canEvVarsCreated :: [CtEvidence] -> TcS StopOrContinue canEvVarsCreated [] = return Stop -- Add all but one to the work list -- and return the first (if any) for futher processing canEvVarsCreated (ev : evs) = do { emitWorkNC evs; canEvNC ev } -- Note the "NC": these are fresh goals, not necessarily canonical emitWorkNC :: [CtEvidence] -> TcS () emitWorkNC evs | null evs = return () | otherwise = updWorkListTcS (extendWorkListCts (map mk_nc evs)) where mk_nc ev = mkNonCanonical ev ------------------------- canEqNC :: CtEvidence -> Type -> Type -> TcS StopOrContinue canEqNC ev ty1 ty2 = can_eq_nc ev ty1 ty1 ty2 ty2 can_eq_nc, can_eq_nc' :: CtEvidence -> Type -> Type -- LHS, after and before type-synonym expansion, resp -> Type -> Type -- RHS, after and before type-synonym expansion, resp -> TcS StopOrContinue can_eq_nc ev ty1 ps_ty1 ty2 ps_ty2 = do { traceTcS "can_eq_nc" $ vcat [ ppr ev, ppr ty1, ppr ps_ty1, ppr ty2, ppr ps_ty2 ] ; can_eq_nc' ev ty1 ps_ty1 ty2 ps_ty2 } -- Expand synonyms first; see Note [Type synonyms and canonicalization] can_eq_nc' ev ty1 ps_ty1 ty2 ps_ty2 | Just ty1' <- tcView ty1 = can_eq_nc ev ty1' ps_ty1 ty2 ps_ty2 | Just ty2' <- tcView ty2 = can_eq_nc ev ty1 ps_ty1 ty2' ps_ty2 -- Type family on LHS or RHS take priority can_eq_nc' ev (TyConApp fn tys) _ ty2 ps_ty2 | isSynFamilyTyCon fn = canEqLeafFun ev NotSwapped fn tys ty2 ps_ty2 can_eq_nc' ev ty1 ps_ty1 (TyConApp fn tys) _ | isSynFamilyTyCon fn = canEqLeafFun ev IsSwapped fn tys ty1 ps_ty1 -- Type variable on LHS or RHS are next can_eq_nc' ev (TyVarTy tv1) _ ty2 ps_ty2 = canEqTyVar ev NotSwapped tv1 ty2 ps_ty2 can_eq_nc' ev ty1 ps_ty1 (TyVarTy tv2) _ = canEqTyVar ev IsSwapped tv2 ty1 ps_ty1 ---------------------- -- Otherwise try to decompose ---------------------- -- Literals can_eq_nc' ev ty1@(LitTy l1) _ (LitTy l2) _ | l1 == l2 = do { when (isWanted ev) $ setEvBind (ctev_evar ev) (EvCoercion (mkTcNomReflCo ty1)) ; return Stop } -- Decomposable type constructor applications -- Synonyms and type functions (which are not decomposable) -- have already been dealt with can_eq_nc' ev (TyConApp tc1 tys1) _ (TyConApp tc2 tys2) _ | isDecomposableTyCon tc1 , isDecomposableTyCon tc2 = canDecomposableTyConApp ev tc1 tys1 tc2 tys2 can_eq_nc' ev (TyConApp tc1 _) ps_ty1 (FunTy {}) ps_ty2 | isDecomposableTyCon tc1 -- The guard is important -- e.g. (x -> y) ~ (F x y) where F has arity 1 -- should not fail, but get the app/app case = canEqFailure ev ps_ty1 ps_ty2 can_eq_nc' ev (FunTy s1 t1) _ (FunTy s2 t2) _ = canDecomposableTyConAppOK ev funTyCon [s1,t1] [s2,t2] can_eq_nc' ev (FunTy {}) ps_ty1 (TyConApp tc2 _) ps_ty2 | isDecomposableTyCon tc2 = canEqFailure ev ps_ty1 ps_ty2 can_eq_nc' ev s1@(ForAllTy {}) _ s2@(ForAllTy {}) _ | CtWanted { ctev_loc = loc, ctev_evar = orig_ev } <- ev = do { let (tvs1,body1) = tcSplitForAllTys s1 (tvs2,body2) = tcSplitForAllTys s2 ; if not (equalLength tvs1 tvs2) then canEqFailure ev s1 s2 else do { traceTcS "Creating implication for polytype equality" $ ppr ev ; ev_term <- deferTcSForAllEq Nominal loc (tvs1,body1) (tvs2,body2) ; setEvBind orig_ev ev_term ; return Stop } } | otherwise = do { traceTcS "Ommitting decomposition of given polytype equality" $ pprEq s1 s2 -- See Note [Do not decompose given polytype equalities] ; return Stop } can_eq_nc' ev (AppTy s1 t1) ps_ty1 ty2 ps_ty2 = can_eq_app ev NotSwapped s1 t1 ps_ty1 ty2 ps_ty2 can_eq_nc' ev ty1 ps_ty1 (AppTy s2 t2) ps_ty2 = can_eq_app ev IsSwapped s2 t2 ps_ty2 ty1 ps_ty1 -- Everything else is a definite type error, eg LitTy ~ TyConApp can_eq_nc' ev _ ps_ty1 _ ps_ty2 = canEqFailure ev ps_ty1 ps_ty2 ------------ can_eq_app, can_eq_flat_app :: CtEvidence -> SwapFlag -> Type -> Type -> Type -- LHS (s1 t2), after and before type-synonym expansion, resp -> Type -> Type -- RHS (ty2), after and before type-synonym expansion, resp -> TcS StopOrContinue -- See Note [Canonicalising type applications] can_eq_app ev swapped s1 t1 ps_ty1 ty2 ps_ty2 = do { traceTcS "can_eq_app 1" $ vcat [ ppr ev, ppr swapped, ppr s1, ppr t1, ppr ty2 ] ; (xi_s1, co_s1) <- flatten FMSubstOnly ev s1 ; traceTcS "can_eq_app 2" $ vcat [ ppr ev, ppr xi_s1 ] ; if s1 `tcEqType` xi_s1 then can_eq_flat_app ev swapped s1 t1 ps_ty1 ty2 ps_ty2 else do { (xi_t1, co_t1) <- flatten FMSubstOnly ev t1 -- We flatten t1 as well so that (xi_s1 xi_t1) is well-kinded -- If we form (xi_s1 t1) that might (appear) ill-kinded, -- and then crash in a call to typeKind ; let xi1 = mkAppTy xi_s1 xi_t1 co1 = mkTcAppCo co_s1 co_t1 ; traceTcS "can_eq_app 3" $ vcat [ ppr ev, ppr xi1, ppr co1 ] ; mb_ct <- rewriteEqEvidence ev swapped xi1 ps_ty2 co1 (mkTcNomReflCo ps_ty2) ; traceTcS "can_eq_app 4" $ vcat [ ppr ev, ppr xi1, ppr co1 ] ; case mb_ct of Nothing -> return Stop Just new_ev -> can_eq_nc new_ev xi1 xi1 ty2 ps_ty2 }} -- Preconditions: s1 is already flattened -- ty2 is not a type variable, so flattening -- can't turn it into an application if it -- doesn't look like one already -- See Note [Canonicalising type applications] can_eq_flat_app ev swapped s1 t1 ps_ty1 ty2 ps_ty2 | Just (s2,t2) <- tcSplitAppTy_maybe ty2 = unSwap swapped decompose_it (s1,t1) (s2,t2) | otherwise = unSwap swapped (canEqFailure ev) ps_ty1 ps_ty2 where decompose_it (s1,t1) (s2,t2) = do { let xevcomp [x,y] = EvCoercion (mkTcAppCo (evTermCoercion x) (evTermCoercion y)) xevcomp _ = error "canEqAppTy: can't happen" -- Can't happen xevdecomp x = let xco = evTermCoercion x in [ EvCoercion (mkTcLRCo CLeft xco) , EvCoercion (mkTcLRCo CRight xco)] ; ctevs <- xCtEvidence ev (XEvTerm [mkTcEqPred s1 s2, mkTcEqPred t1 t2] xevcomp xevdecomp) ; canEvVarsCreated ctevs } ------------------------ canDecomposableTyConApp :: CtEvidence -> TyCon -> [TcType] -> TyCon -> [TcType] -> TcS StopOrContinue canDecomposableTyConApp ev tc1 tys1 tc2 tys2 | tc1 /= tc2 || length tys1 /= length tys2 -- Fail straight away for better error messages = canEqFailure ev (mkTyConApp tc1 tys1) (mkTyConApp tc2 tys2) | otherwise = canDecomposableTyConAppOK ev tc1 tys1 tys2 canDecomposableTyConAppOK :: CtEvidence -> TyCon -> [TcType] -> [TcType] -> TcS StopOrContinue canDecomposableTyConAppOK ev tc1 tys1 tys2 = do { let xcomp xs = EvCoercion (mkTcTyConAppCo Nominal tc1 (map evTermCoercion xs)) xdecomp x = zipWith (\_ i -> EvCoercion $ mkTcNthCo i (evTermCoercion x)) tys1 [0..] xev = XEvTerm (zipWith mkTcEqPred tys1 tys2) xcomp xdecomp ; ctevs <- xCtEvidence ev xev ; canEvVarsCreated ctevs } canEqFailure :: CtEvidence -> TcType -> TcType -> TcS StopOrContinue -- See Note [Make sure that insolubles are fully rewritten] canEqFailure ev ty1 ty2 = do { (s1, co1) <- flatten FMSubstOnly ev ty1 ; (s2, co2) <- flatten FMSubstOnly ev ty2 ; mb_ct <- rewriteEqEvidence ev NotSwapped s1 s2 co1 co2 ; case mb_ct of Just new_ev -> emitInsoluble (mkNonCanonical new_ev) Nothing -> pprPanic "canEqFailure" (ppr ev $$ ppr ty1 $$ ppr ty2) ; return Stop }\end{code} 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_app flattens s1. If flattening does something, it rewrites, and goes round can_eq_nc again. If flattening does nothing, then (at least with our present state of knowledge) we can only decompose, and that is what can_eq_flat_app attempts to do. 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 [Kick out insolubles] in TcInteract. And if we don't do this there is a bad danger that TcSimplify.applyTyVarDefaulting will find a variable that has in fact been substituted. Note [Do not decompose Given polytype equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider [G] (forall a. t1 ~ forall a. t2). Can we decompose this? No -- what would the evidence look like? So instead we simply discard this given evidence. Note [Combining insoluble constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ As this point we have an insoluble constraint, like Int~Bool. * If it is Wanted, delete it from the cache, so that subsequent Int~Bool constraints give rise to separate error messages * But if it is Derived, DO NOT delete from cache. A class constraint may get kicked out of the inert set, and then have its functional dependency Derived constraints generated a second time. In that case we don't want to get two (or more) error messages by generating two (or more) insoluble fundep constraints from the same class constraint. Note [Canonical ordering for equality constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Implemented as (<+=) below: - Type function applications always come before anything else. - Variables always come before non-variables (other than type function applications). Note that we don't need to unfold type synonyms on the RHS to check the ordering; that is, in the rules above it's OK to consider only whether something is *syntactically* a type function application or not. To illustrate why this is OK, suppose we have an equality of the form 'tv ~ S a b c', where S is a type synonym which expands to a top-level application of the type function F, something like type S a b c = F d e Then to canonicalize 'tv ~ S a b c' we flatten the RHS, and since S's expansion contains type function applications the flattener will do the expansion and then generate a skolem variable for the type function application, so we end up with something like this: tv ~ x F d e ~ x where x is the skolem variable. This is one extra equation than absolutely necessary (we could have gotten away with just 'F d e ~ tv' if we had noticed that S expanded to a top-level type function application and flipped it around in the first place) but this way keeps the code simpler. Unlike the OutsideIn(X) draft of May 7, 2010, we do not care about the ordering of tv ~ tv constraints. There are several reasons why we might: (1) In order to be able to extract a substitution that doesn't mention untouchable variables after we are done solving, we might prefer to put touchable variables on the left. However, in and of itself this isn't necessary; we can always re-orient equality constraints at the end if necessary when extracting a substitution. (2) To ensure termination we might think it necessary to put variables in lexicographic order. However, this isn't actually necessary as outlined below. While building up an inert set of canonical constraints, we maintain the invariant that the equality constraints in the inert set form an acyclic rewrite system when viewed as L-R rewrite rules. Moreover, the given constraints form an idempotent substitution (i.e. none of the variables on the LHS occur in any of the RHS's, and type functions never show up in the RHS at all), the wanted constraints also form an idempotent substitution, and finally the LHS of a given constraint never shows up on the RHS of a wanted constraint. There may, however, be a wanted LHS that shows up in a given RHS, since we do not rewrite given constraints with wanted constraints. Suppose we have an inert constraint set tg_1 ~ xig_1 -- givens tg_2 ~ xig_2 ... tw_1 ~ xiw_1 -- wanteds tw_2 ~ xiw_2 ... where each t_i can be either a type variable or a type function application. Now suppose we take a new canonical equality constraint, t' ~ xi' (note among other things this means t' does not occur in xi') and try to react it with the existing inert set. We show by induction on the number of t_i which occur in t' ~ xi' that this process will terminate. There are several ways t' ~ xi' could react with an existing constraint: TODO: finish this proof. The below was for the case where the entire inert set is an idempotent subustitution... (b) We could have t' = t_j for some j. Then we obtain the new equality xi_j ~ xi'; note that neither xi_j or xi' contain t_j. We now canonicalize the new equality, which may involve decomposing it into several canonical equalities, and recurse on these. However, none of the new equalities will contain t_j, so they have fewer occurrences of the t_i than the original equation. (a) We could have t_j occurring in xi' for some j, with t' /= t_j. Then we substitute xi_j for t_j in xi' and continue. However, since none of the t_i occur in xi_j, we have decreased the number of t_i that occur in xi', since we eliminated t_j and did not introduce any new ones. \begin{code}

canEqLeafFun :: CtEvidence -> SwapFlag -> TyCon -> [TcType] -- LHS -> TcType -> TcType -- RHS -> TcS StopOrContinue canEqLeafFun ev swapped fn tys1 ty2 ps_ty2 | length tys1 > tyConArity fn = -- Over-saturated type function on LHS: -- flatten LHS, leaving an AppTy, and go around again do { (xi1, co1) <- flatten FMFullFlatten ev (mkTyConApp fn tys1) ; mb <- rewriteEqEvidence ev swapped xi1 ps_ty2 co1 (mkTcNomReflCo ps_ty2) ; case mb of Nothing -> return Stop Just new_ev -> can_eq_nc new_ev xi1 xi1 ty2 ps_ty2 } | otherwise = -- ev :: F tys1 ~ ty2, if not swapped -- ev :: ty2 ~ F tys1, if swapped ASSERT( length tys1 == tyConArity fn ) -- Type functions are never under-saturated -- Previous equation checks for over-saturation do { traceTcS "canEqLeafFun" $ pprEq (mkTyConApp fn tys1) ps_ty2 -- Flatten type function arguments -- cos1 :: xis1 ~ tys1 -- co2 :: xi2 ~ ty2 ; (xis1,cos1) <- flattenMany FMFullFlatten ev tys1 ; (xi2, co2) <- flatten FMFullFlatten ev ps_ty2 ; let fam_head = mkTyConApp fn xis1 co1 = mkTcTyConAppCo Nominal fn cos1 ; mb <- rewriteEqEvidence ev swapped fam_head xi2 co1 co2 ; let k1 = typeKind fam_head k2 = typeKind xi2 ; case mb of Nothing -> return Stop Just new_ev | k1 `isSubKind` k2 -- Establish CFunEqCan kind invariant -> continueWith (CFunEqCan { cc_ev = new_ev, cc_fun = fn , cc_tyargs = xis1, cc_rhs = xi2 }) | otherwise -> checkKind new_ev fam_head k1 xi2 k2 } --------------------- canEqTyVar :: CtEvidence -> SwapFlag -> TcTyVar -> TcType -> TcType -> TcS StopOrContinue -- A TyVar on LHS, but so far un-zonked canEqTyVar ev swapped tv1 ty2 ps_ty2 -- ev :: tv ~ s2 = do { traceTcS "canEqTyVar" (ppr tv1 $$ ppr ty2 $$ ppr swapped) ; mb_yes <- flattenTyVarOuter FMFullFlatten ev tv1 ; case mb_yes of Right (ty1, co1) -> -- co1 :: ty1 ~ tv1 do { mb <- rewriteEqEvidence ev swapped ty1 ps_ty2 co1 (mkTcNomReflCo ps_ty2) ; traceTcS "canEqTyVar2" (vcat [ppr tv1, ppr ty2, ppr swapped, ppr ty1, ppUnless (isDerived ev) (ppr co1)]) ; case mb of Nothing -> return Stop Just new_ev -> can_eq_nc new_ev ty1 ty1 ty2 ps_ty2 } Left tv1' -> do { (xi2, co2) <- flatten FMFullFlatten ev ps_ty2 -- co2 :: xi2 ~ ps_ty2 -- Use ps_ty2 to preserve type synonyms if poss ; dflags <- getDynFlags ; canEqTyVar2 dflags ev swapped tv1' xi2 co2 } } canEqTyVar2 :: DynFlags -> CtEvidence -- olhs ~ orhs (or, if swapped, orhs ~ olhs) -> SwapFlag -> TcTyVar -- olhs -> TcType -- nrhs -> TcCoercion -- nrhs ~ orhs -> TcS StopOrContinue -- LHS is an inert type variable, -- and RHS is fully rewritten, but with type synonyms -- preserved as must as possible canEqTyVar2 dflags ev swapped tv1 xi2 co2 | Just tv2 <- getTyVar_maybe xi2 = canEqTyVarTyVar ev swapped tv1 tv2 co2 | OC_OK xi2' <- occurCheckExpand dflags tv1 xi2 -- No occurs check = do { mb <- rewriteEqEvidence ev swapped xi1 xi2' co1 co2 -- Ensure that the new goal has enough type synonyms -- expanded by the occurCheckExpand; hence using xi2' here -- See Note [occurCheckExpand] ; let k1 = tyVarKind tv1 k2 = typeKind xi2' ; case mb of Nothing -> return Stop Just new_ev | k2 `isSubKind` k1 -- Establish CTyEqCan kind invariant -- Reorientation has done its best, but the kinds might -- simply be incompatible -> continueWith (CTyEqCan { cc_ev = new_ev , cc_tyvar = tv1, cc_rhs = xi2' }) | otherwise -> checkKind new_ev xi1 k1 xi2' k2 } | otherwise -- Occurs check error = do { mb <- rewriteEqEvidence ev swapped xi1 xi2 co1 co2 ; case mb of Nothing -> return () Just new_ev -> emitInsoluble (mkNonCanonical new_ev) ; return Stop } where xi1 = mkTyVarTy tv1 co1 = mkTcNomReflCo xi1 canEqTyVarTyVar :: CtEvidence -- tv1 ~ orhs (or orhs ~ tv1, if swapped) -> SwapFlag -> TyVar -> TyVar -- tv2, tv2 -> TcCoercion -- tv2 ~ orhs -> TcS StopOrContinue -- Both LHS and RHS rewrote to a type variable, canEqTyVarTyVar ev swapped tv1 tv2 co2 | tv1 == tv2 = do { when (isWanted ev) $ ASSERT( tcCoercionRole co2 == Nominal ) setEvBind (ctev_evar ev) (EvCoercion (maybeSym swapped co2)) ; return Stop } | reorient_me -- See note [Canonical ordering for equality constraints]. -- True => the kinds are compatible, -- so no need for further sub-kind check -- If swapped = NotSwapped, then -- rw_orhs = tv1, rw_olhs = orhs -- rw_nlhs = tv2, rw_nrhs = xi1 = do { mb <- rewriteEqEvidence ev (flipSwap swapped) xi2 xi1 co2 (mkTcNomReflCo xi1) ; case mb of Nothing -> return Stop Just new_ev -> continueWith (CTyEqCan { cc_ev = new_ev , cc_tyvar = tv2, cc_rhs = xi1 }) } | otherwise = do { mb <- rewriteEqEvidence ev swapped xi1 xi2 (mkTcNomReflCo xi1) co2 ; case mb of Nothing -> return Stop Just new_ev | k2 `isSubKind` k1 -> continueWith (CTyEqCan { cc_ev = new_ev , cc_tyvar = tv1, cc_rhs = xi2 }) | otherwise -> checkKind new_ev xi1 k1 xi2 k2 } where reorient_me | k1 `tcEqKind` k2 = tv2 `better_than` tv1 | k1 `isSubKind` k2 = True -- Note [Kind orientation for CTyEqCan] | otherwise = False -- in TcRnTypes xi1 = mkTyVarTy tv1 xi2 = mkTyVarTy tv2 k1 = tyVarKind tv1 k2 = tyVarKind tv2 tv2 `better_than` tv1 | isMetaTyVar tv1 = False -- Never swap a meta-tyvar | isFlatSkolTyVar tv1 = isMetaTyVar tv2 | otherwise = isMetaTyVar tv2 || isFlatSkolTyVar tv2 -- Note [Eliminate flat-skols] checkKind :: CtEvidence -- t1~t2 -> TcType -> TcKind -> TcType -> TcKind -- s1~s2, flattened and zonked -> TcS StopOrContinue -- LHS and RHS have incompatible kinds, so emit an "irreducible" constraint -- CIrredEvCan (NOT CTyEqCan or CFunEqCan) -- for the type equality; and continue with the kind equality constraint. -- When the latter is solved, it'll kick out the irreducible equality for -- a second attempt at solving checkKind new_ev s1 k1 s2 k2 -- See Note [Equalities with incompatible kinds] = ASSERT( isKind k1 && isKind k2 ) do { traceTcS "canEqLeaf: incompatible kinds" (vcat [ppr k1, ppr k2]) -- Create a derived kind-equality, and solve it ; mw <- newDerived kind_co_loc (mkEqPred k1 k2) ; case mw of Nothing -> return () Just kev -> emitWorkNC [kev] -- Put the not-currently-soluble thing into the inert set ; continueWith (CIrredEvCan { cc_ev = new_ev }) } where loc = ctev_loc new_ev kind_co_loc = setCtLocOrigin loc (KindEqOrigin s1 s2 (ctLocOrigin loc))\end{code} Note [Eliminate flat-skols] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ Suppose we have [G] Num (F [a]) then we flatten to [G] Num fsk [G] F [a] ~ fsk where fsk is a flatten-skolem (FlatSkol). Suppose we have type instance F [a] = a then we'll reduce the second constraint to [G] a ~ fsk and then replace all uses of 'a' with fsk. That's bad because in error messages intead of saying 'a' we'll say (F [a]). In all places, including those where the programmer wrote 'a' in the first place. Very confusing! See Trac #7862. Solution: re-orient a~fsk to fsk~a, so that we preferentially eliminate the fsk. Note [Equalities with incompatible kinds] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ canEqLeaf is about to make a CTyEqCan or CFunEqCan; but both have the invariant that LHS and RHS satisfy the kind invariants for CTyEqCan, CFunEqCan. What if we try to unify two things with incompatible kinds? eg a ~ b where a::*, b::*->* or a ~ b where a::*, b::k, k is a kind variable The CTyEqCan compatKind invariant is important. If we make a CTyEqCan for a~b, then we might well *substitute* 'b' for 'a', and that might make a well-kinded type ill-kinded; and that is bad (eg typeKind can crash, see Trac #7696). So instead for these ill-kinded equalities we generate a CIrredCan, and put it in the inert set, which keeps it out of the way until a subsequent substitution (on kind variables, say) re-activates it. NB: it is important that the types s1,s2 are flattened and zonked so that their kinds k1, k2 are inert wrt the substitution. That means that they can only become the same if we change the inert set, which in turn will kick out the irreducible equality E.g. it is WRONG to make an irred (a:k1)~(b:k2) if we already have a substitution k1:=k2 NB: it's important that the new CIrredCan goes in the inert set rather than back into the work list. We used to do the latter, but that led to an infinite loop when we encountered it again, and put it back it the work list again. See also Note [Kind orientation for CTyEqCan] and Note [Kind orientation for CFunEqCan] in TcRnTypes Note [Type synonyms and canonicalization] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We treat type synonym applications as xi types, that is, they do not count as type function applications. However, we do need to be a bit careful with type synonyms: like type functions they may not be generative or injective. However, unlike type functions, they are parametric, so there is no problem in expanding them whenever we see them, since we do not need to know anything about their arguments in order to expand them; this is what justifies not having to treat them as specially as type function applications. The thing that causes some subtleties is that we prefer to leave type synonym applications *unexpanded* whenever possible, in order to generate better error messages. If we encounter an equality constraint with type synonym applications on both sides, or a type synonym application on one side and some sort of type application on the other, we simply must expand out the type synonyms in order to continue decomposing the equality constraint into primitive equality constraints. For example, suppose we have type F a = [Int] and we encounter the equality F a ~ [b] In order to continue we must expand F a into [Int], giving us the equality [Int] ~ [b] which we can then decompose into the more primitive equality constraint Int ~ b. However, if we encounter an equality constraint with a type synonym application on one side and a variable on the other side, we should NOT (necessarily) expand the type synonym, since for the purpose of good error messages we want to leave type synonyms unexpanded as much as possible. Hence the ps_ty1, ps_ty2 argument passed to canEqTyVar. Note [occurCheckExpand] ~~~~~~~~~~~~~~~~~~~~~~~ There is a subtle point with type synonyms and the occurs check that takes place for equality constraints of the form tv ~ xi. As an example, suppose we have type F a = Int and we come across the equality constraint a ~ F a This should not actually fail the occurs check, since expanding out the type synonym results in the legitimate equality constraint a ~ Int. We must actually do this expansion, because unifying a with F a will lead the type checker into infinite loops later. Put another way, canonical equality constraints should never *syntactically* contain the LHS variable in the RHS type. However, we don't always need to expand type synonyms when doing an occurs check; for example, the constraint a ~ F b is obviously fine no matter what F expands to. And in this case we would rather unify a with F b (rather than F b's expansion) in order to get better error messages later. So, when doing an occurs check with a type synonym application on the RHS, we use some heuristics to find an expansion of the RHS which does not contain the variable from the LHS. In particular, given a ~ F t1 ... tn we first try expanding each of the ti to types which no longer contain a. If this turns out to be impossible, we next try expanding F itself, and so on. See Note [Occurs check expansion] in TcType