{-# LANGUAGE CPP, DeriveFunctor, TypeFamilies, ScopedTypeVariables, TypeApplications, DerivingStrategies, GeneralizedNewtypeDeriving, ScopedTypeVariables #-} {-# OPTIONS_GHC -Wno-incomplete-record-updates -Wno-incomplete-uni-patterns #-} -- | Type definitions for the constraint solver module GHC.Tc.Solver.Monad ( -- The work list WorkList(..), isEmptyWorkList, emptyWorkList, extendWorkListNonEq, extendWorkListCt, extendWorkListCts, extendWorkListEq, appendWorkList, selectNextWorkItem, workListSize, getWorkList, updWorkListTcS, pushLevelNoWorkList, -- The TcS monad TcS, runTcS, runTcSDeriveds, runTcSWithEvBinds, runTcSInerts, failTcS, warnTcS, addErrTcS, wrapTcS, runTcSEqualities, nestTcS, nestImplicTcS, setEvBindsTcS, emitImplicationTcS, emitTvImplicationTcS, runTcPluginTcS, addUsedGRE, addUsedGREs, keepAlive, matchGlobalInst, TcM.ClsInstResult(..), QCInst(..), -- Tracing etc panicTcS, traceTcS, traceFireTcS, bumpStepCountTcS, csTraceTcS, wrapErrTcS, wrapWarnTcS, resetUnificationFlag, setUnificationFlag, -- Evidence creation and transformation MaybeNew(..), freshGoals, isFresh, getEvExpr, newTcEvBinds, newNoTcEvBinds, newWantedEq, newWantedEq_SI, emitNewWantedEq, newWanted, newWanted_SI, newWantedEvVar, newWantedNC, newWantedEvVarNC, newDerivedNC, newBoundEvVarId, unifyTyVar, reportUnifications, setEvBind, setWantedEq, setWantedEvTerm, setEvBindIfWanted, newEvVar, newGivenEvVar, newGivenEvVars, emitNewDeriveds, emitNewDerivedEq, checkReductionDepth, getSolvedDicts, setSolvedDicts, getInstEnvs, getFamInstEnvs, -- Getting the environments getTopEnv, getGblEnv, getLclEnv, getTcEvBindsVar, getTcLevel, getTcEvTyCoVars, getTcEvBindsMap, setTcEvBindsMap, tcLookupClass, tcLookupId, -- Inerts InertSet(..), InertCans(..), emptyInert, updInertTcS, updInertCans, updInertDicts, updInertIrreds, getHasGivenEqs, setInertCans, getInertEqs, getInertCans, getInertGivens, getInertInsols, getInnermostGivenEqLevel, getTcSInerts, setTcSInerts, matchableGivens, prohibitedSuperClassSolve, mightEqualLater, getUnsolvedInerts, removeInertCts, getPendingGivenScs, addInertCan, insertFunEq, addInertForAll, emitWorkNC, emitWork, isImprovable, -- The Model kickOutAfterUnification, -- Inert Safe Haskell safe-overlap failures addInertSafehask, insertSafeOverlapFailureTcS, updInertSafehask, getSafeOverlapFailures, -- Inert CDictCans DictMap, emptyDictMap, lookupInertDict, findDictsByClass, addDict, addDictsByClass, delDict, foldDicts, filterDicts, findDict, -- Inert CEqCans EqualCtList(..), findTyEqs, foldTyEqs, findEq, -- Inert solved dictionaries addSolvedDict, lookupSolvedDict, -- Irreds foldIrreds, -- The family application cache lookupFamAppInert, lookupFamAppCache, extendFamAppCache, pprKicked, -- Inert function equalities findFunEq, findFunEqsByTyCon, instDFunType, -- Instantiation -- MetaTyVars newFlexiTcSTy, instFlexi, instFlexiX, cloneMetaTyVar, tcInstSkolTyVarsX, TcLevel, isFilledMetaTyVar_maybe, isFilledMetaTyVar, zonkTyCoVarsAndFV, zonkTcType, zonkTcTypes, zonkTcTyVar, zonkCo, zonkTyCoVarsAndFVList, zonkSimples, zonkWC, zonkTyCoVarKind, -- References newTcRef, readTcRef, writeTcRef, updTcRef, -- Misc getDefaultInfo, getDynFlags, getGlobalRdrEnvTcS, matchFam, matchFamTcM, checkWellStagedDFun, pprEq, -- Smaller utils, re-exported from TcM -- TODO (DV): these are only really used in the -- instance matcher in GHC.Tc.Solver. I am wondering -- if the whole instance matcher simply belongs -- here breakTyVarCycle, rewriterView ) where #include "HsVersions.h" import GHC.Prelude import GHC.Driver.Env import qualified GHC.Tc.Utils.Instantiate as TcM import GHC.Core.InstEnv import GHC.Tc.Instance.Family as FamInst import GHC.Core.FamInstEnv import qualified GHC.Tc.Utils.Monad as TcM import qualified GHC.Tc.Utils.TcMType as TcM import qualified GHC.Tc.Instance.Class as TcM( matchGlobalInst, ClsInstResult(..) ) import qualified GHC.Tc.Utils.Env as TcM ( checkWellStaged, tcGetDefaultTys, tcLookupClass, tcLookupId, topIdLvl ) import GHC.Tc.Instance.Class( InstanceWhat(..), safeOverlap, instanceReturnsDictCon ) import GHC.Tc.Utils.TcType import GHC.Driver.Session import GHC.Core.Type import qualified GHC.Core.TyCo.Rep as Rep -- this needs to be used only very locally import GHC.Core.Coercion import GHC.Core.Unify import GHC.Tc.Types.Evidence import GHC.Core.Class import GHC.Core.TyCon import GHC.Tc.Errors ( solverDepthErrorTcS ) import GHC.Types.Name import GHC.Types.TyThing import GHC.Unit.Module ( HasModule, getModule ) import GHC.Types.Name.Reader ( GlobalRdrEnv, GlobalRdrElt ) import qualified GHC.Rename.Env as TcM import GHC.Types.Var import GHC.Types.Var.Env import GHC.Types.Var.Set import GHC.Utils.Outputable import GHC.Utils.Panic import GHC.Utils.Logger import GHC.Data.Bag as Bag import GHC.Types.Unique.Supply import GHC.Utils.Misc import GHC.Tc.Types import GHC.Tc.Types.Origin import GHC.Tc.Types.Constraint import GHC.Core.Predicate import GHC.Types.Unique.Set import GHC.Core.TyCon.Env import GHC.Data.Maybe import GHC.Core.Map.Type import GHC.Data.TrieMap import Control.Monad import GHC.Utils.Monad import Data.IORef import GHC.Exts (oneShot) import Data.List ( partition, mapAccumL ) import Data.List.NonEmpty ( NonEmpty(..), cons, toList, nonEmpty ) import qualified Data.List.NonEmpty as NE import Control.Arrow ( first ) #if defined(DEBUG) import GHC.Data.Graph.Directed #endif {- ************************************************************************ * * * Worklists * * Canonical and non-canonical constraints that the simplifier has to * * work on. Including their simplification depths. * * * * * ************************************************************************ Note [WorkList priorities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ A WorkList contains canonical and non-canonical items (of all flavours). Notice that each Ct now has a simplification depth. We may consider using this depth for prioritization as well in the future. As a simple form of priority queue, our worklist separates out * equalities (wl_eqs); see Note [Prioritise equalities] * all the rest (wl_rest) Note [Prioritise equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~ It's very important to process equalities /first/: * (Efficiency) The general reason to do so is that if we process a class constraint first, we may end up putting it into the inert set and then kicking it out later. That's extra work compared to just doing the equality first. * (Avoiding fundep iteration) As #14723 showed, it's possible to get non-termination if we - Emit the Derived fundep equalities for a class constraint, generating some fresh unification variables. - That leads to some unification - Which kicks out the class constraint - Which isn't solved (because there are still some more Derived equalities in the work-list), but generates yet more fundeps Solution: prioritise derived equalities over class constraints * (Class equalities) We need to prioritise equalities even if they are hidden inside a class constraint; see Note [Prioritise class equalities] * (Kick-out) We want to apply this priority scheme to kicked-out constraints too (see the call to extendWorkListCt in kick_out_rewritable E.g. a CIrredCan can be a hetero-kinded (t1 ~ t2), which may become homo-kinded when kicked out, and hence we want to prioritise it. * (Derived equalities) Originally we tried to postpone processing Derived equalities, in the hope that we might never need to deal with them at all; but in fact we must process Derived equalities eagerly, partly for the (Efficiency) reason, and more importantly for (Avoiding fundep iteration). Note [Prioritise class equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We prioritise equalities in the solver (see selectWorkItem). But class constraints like (a ~ b) and (a ~~ b) are actually equalities too; see Note [The equality types story] in GHC.Builtin.Types.Prim. Failing to prioritise these is inefficient (more kick-outs etc). But, worse, it can prevent us spotting a "recursive knot" among Wanted constraints. See comment:10 of #12734 for a worked-out example. So we arrange to put these particular class constraints in the wl_eqs. NB: since we do not currently apply the substitution to the inert_solved_dicts, the knot-tying still seems a bit fragile. But this makes it better. -} -- See Note [WorkList priorities] data WorkList = WL { wl_eqs :: [Ct] -- CEqCan, CDictCan, CIrredCan -- Given, Wanted, and Derived -- Contains both equality constraints and their -- class-level variants (a~b) and (a~~b); -- See Note [Prioritise equalities] -- See Note [Prioritise class equalities] , wl_rest :: [Ct] , wl_implics :: Bag Implication -- See Note [Residual implications] } appendWorkList :: WorkList -> WorkList -> WorkList appendWorkList (WL { wl_eqs = eqs1, wl_rest = rest1 , wl_implics = implics1 }) (WL { wl_eqs = eqs2, wl_rest = rest2 , wl_implics = implics2 }) = WL { wl_eqs = eqs1 ++ eqs2 , wl_rest = rest1 ++ rest2 , wl_implics = implics1 `unionBags` implics2 } workListSize :: WorkList -> Int workListSize (WL { wl_eqs = eqs, wl_rest = rest }) = length eqs + length rest extendWorkListEq :: Ct -> WorkList -> WorkList extendWorkListEq ct wl = wl { wl_eqs = ct : wl_eqs wl } extendWorkListNonEq :: Ct -> WorkList -> WorkList -- Extension by non equality extendWorkListNonEq ct wl = wl { wl_rest = ct : wl_rest wl } extendWorkListDeriveds :: [CtEvidence] -> WorkList -> WorkList extendWorkListDeriveds evs wl = extendWorkListCts (map mkNonCanonical evs) wl extendWorkListImplic :: Implication -> WorkList -> WorkList extendWorkListImplic implic wl = wl { wl_implics = implic `consBag` wl_implics wl } extendWorkListCt :: Ct -> WorkList -> WorkList -- Agnostic extendWorkListCt ct wl = case classifyPredType (ctPred ct) of EqPred {} -> extendWorkListEq ct wl ClassPred cls _ -- See Note [Prioritise class equalities] | isEqPredClass cls -> extendWorkListEq ct wl _ -> extendWorkListNonEq ct wl extendWorkListCts :: [Ct] -> WorkList -> WorkList -- Agnostic extendWorkListCts cts wl = foldr extendWorkListCt wl cts isEmptyWorkList :: WorkList -> Bool isEmptyWorkList (WL { wl_eqs = eqs, wl_rest = rest, wl_implics = implics }) = null eqs && null rest && isEmptyBag implics emptyWorkList :: WorkList emptyWorkList = WL { wl_eqs = [], wl_rest = [], wl_implics = emptyBag } selectWorkItem :: WorkList -> Maybe (Ct, WorkList) -- See Note [Prioritise equalities] selectWorkItem wl@(WL { wl_eqs = eqs, wl_rest = rest }) | ct:cts <- eqs = Just (ct, wl { wl_eqs = cts }) | ct:cts <- rest = Just (ct, wl { wl_rest = cts }) | otherwise = Nothing getWorkList :: TcS WorkList getWorkList = do { wl_var <- getTcSWorkListRef ; wrapTcS (TcM.readTcRef wl_var) } selectNextWorkItem :: TcS (Maybe Ct) -- Pick which work item to do next -- See Note [Prioritise equalities] selectNextWorkItem = do { wl_var <- getTcSWorkListRef ; wl <- readTcRef wl_var ; case selectWorkItem wl of { Nothing -> return Nothing ; Just (ct, new_wl) -> do { -- checkReductionDepth (ctLoc ct) (ctPred ct) -- This is done by GHC.Tc.Solver.Interact.chooseInstance ; writeTcRef wl_var new_wl ; return (Just ct) } } } -- Pretty printing instance Outputable WorkList where ppr (WL { wl_eqs = eqs, wl_rest = rest, wl_implics = implics }) = text "WL" <+> (braces $ vcat [ ppUnless (null eqs) $ text "Eqs =" <+> vcat (map ppr eqs) , ppUnless (null rest) $ text "Non-eqs =" <+> vcat (map ppr rest) , ppUnless (isEmptyBag implics) $ ifPprDebug (text "Implics =" <+> vcat (map ppr (bagToList implics))) (text "(Implics omitted)") ]) {- ********************************************************************* * * InertSet: the inert set * * * * ********************************************************************* -} data InertSet = IS { inert_cans :: InertCans -- Canonical Given, Wanted, Derived -- Sometimes called "the inert set" , inert_cycle_breakers :: [(TcTyVar, TcType)] -- a list of CycleBreakerTv / original family applications -- used to undo the cycle-breaking needed to handle -- Note [Type variable cycles in Givens] in GHC.Tc.Solver.Canonical , inert_famapp_cache :: FunEqMap (TcCoercion, TcType) -- Just a hash-cons cache for use when reducing family applications -- only -- -- If F tys :-> (co, rhs, flav), -- then co :: rhs ~N F tys -- all evidence is from instances or Givens; no coercion holes here -- (We have no way of "kicking out" from the cache, so putting -- wanteds here means we can end up solving a Wanted with itself. Bad) , inert_solved_dicts :: DictMap CtEvidence -- All Wanteds, of form ev :: C t1 .. tn -- See Note [Solved dictionaries] -- and Note [Do not add superclasses of solved dictionaries] } instance Outputable InertSet where ppr (IS { inert_cans = ics , inert_solved_dicts = solved_dicts }) = vcat [ ppr ics , ppUnless (null dicts) $ text "Solved dicts =" <+> vcat (map ppr dicts) ] where dicts = bagToList (dictsToBag solved_dicts) emptyInertCans :: InertCans emptyInertCans = IC { inert_eqs = emptyDVarEnv , inert_given_eq_lvl = topTcLevel , inert_given_eqs = False , inert_dicts = emptyDicts , inert_safehask = emptyDicts , inert_funeqs = emptyFunEqs , inert_insts = [] , inert_irreds = emptyCts } emptyInert :: InertSet emptyInert = IS { inert_cans = emptyInertCans , inert_cycle_breakers = [] , inert_famapp_cache = emptyFunEqs , inert_solved_dicts = emptyDictMap } {- Note [Solved dictionaries] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ When we apply a top-level instance declaration, we add the "solved" dictionary to the inert_solved_dicts. In general, we use it to avoid creating a new EvVar when we have a new goal that we have solved in the past. But in particular, we can use it to create *recursive* dictionaries. The simplest, degenerate case is instance C [a] => C [a] where ... If we have [W] d1 :: C [x] then we can apply the instance to get d1 = $dfCList d [W] d2 :: C [x] Now 'd1' goes in inert_solved_dicts, and we can solve d2 directly from d1. d1 = $dfCList d d2 = d1 See Note [Example of recursive dictionaries] VERY IMPORTANT INVARIANT: (Solved Dictionary Invariant) Every member of the inert_solved_dicts is the result of applying an instance declaration that "takes a step" An instance "takes a step" if it has the form dfunDList d1 d2 = MkD (...) (...) (...) That is, the dfun is lazy in its arguments, and guarantees to immediately return a dictionary constructor. NB: all dictionary data constructors are lazy in their arguments. This property is crucial to ensure that all dictionaries are non-bottom, which in turn ensures that the whole "recursive dictionary" idea works at all, even if we get something like rec { d = dfunDList d dx } See Note [Recursive superclasses] in GHC.Tc.TyCl.Instance. Reason: - All instances, except two exceptions listed below, "take a step" in the above sense - Exception 1: local quantified constraints have no such guarantee; indeed, adding a "solved dictionary" when appling a quantified constraint led to the ability to define unsafeCoerce in #17267. - Exception 2: the magic built-in instance for (~) has no such guarantee. It behaves as if we had class (a ~# b) => (a ~ b) where {} instance (a ~# b) => (a ~ b) where {} The "dfun" for the instance is strict in the coercion. Anyway there's no point in recording a "solved dict" for (t1 ~ t2); it's not going to allow a recursive dictionary to be constructed. Ditto (~~) and Coercible. THEREFORE we only add a "solved dictionary" - when applying an instance declaration - subject to Exceptions 1 and 2 above In implementation terms - GHC.Tc.Solver.Monad.addSolvedDict adds a new solved dictionary, conditional on the kind of instance - It is only called when applying an instance decl, in GHC.Tc.Solver.Interact.doTopReactDict - ClsInst.InstanceWhat says what kind of instance was used to solve the constraint. In particular * LocalInstance identifies quantified constraints * BuiltinEqInstance identifies the strange built-in instances for equality. - ClsInst.instanceReturnsDictCon says which kind of instance guarantees to return a dictionary constructor Other notes about solved dictionaries * See also Note [Do not add superclasses of solved dictionaries] * The inert_solved_dicts field is not rewritten by equalities, so it may get out of date. * The inert_solved_dicts are all Wanteds, never givens * We only cache dictionaries from top-level instances, not from local quantified constraints. Reason: if we cached the latter we'd need to purge the cache when bringing new quantified constraints into scope, because quantified constraints "shadow" top-level instances. Note [Do not add superclasses of solved dictionaries] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Every member of inert_solved_dicts is the result of applying a dictionary function, NOT of applying superclass selection to anything. Consider class Ord a => C a where instance Ord [a] => C [a] where ... Suppose we are trying to solve [G] d1 : Ord a [W] d2 : C [a] Then we'll use the instance decl to give [G] d1 : Ord a Solved: d2 : C [a] = $dfCList d3 [W] d3 : Ord [a] We must not add d4 : Ord [a] to the 'solved' set (by taking the superclass of d2), otherwise we'll use it to solve d3, without ever using d1, which would be a catastrophe. Solution: when extending the solved dictionaries, do not add superclasses. That's why each element of the inert_solved_dicts is the result of applying a dictionary function. Note [Example of recursive dictionaries] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ --- Example 1 data D r = ZeroD | SuccD (r (D r)); instance (Eq (r (D r))) => Eq (D r) where ZeroD == ZeroD = True (SuccD a) == (SuccD b) = a == b _ == _ = False; equalDC :: D [] -> D [] -> Bool; equalDC = (==); We need to prove (Eq (D [])). Here's how we go: [W] d1 : Eq (D []) By instance decl of Eq (D r): [W] d2 : Eq [D []] where d1 = dfEqD d2 By instance decl of Eq [a]: [W] d3 : Eq (D []) where d2 = dfEqList d3 d1 = dfEqD d2 Now this wanted can interact with our "solved" d1 to get: d3 = d1 -- Example 2: This code arises in the context of "Scrap Your Boilerplate with Class" class Sat a class Data ctx a instance Sat (ctx Char) => Data ctx Char -- dfunData1 instance (Sat (ctx [a]), Data ctx a) => Data ctx [a] -- dfunData2 class Data Maybe a => Foo a instance Foo t => Sat (Maybe t) -- dfunSat instance Data Maybe a => Foo a -- dfunFoo1 instance Foo a => Foo [a] -- dfunFoo2 instance Foo [Char] -- dfunFoo3 Consider generating the superclasses of the instance declaration instance Foo a => Foo [a] So our problem is this [G] d0 : Foo t [W] d1 : Data Maybe [t] -- Desired superclass We may add the given in the inert set, along with its superclasses Inert: [G] d0 : Foo t [G] d01 : Data Maybe t -- Superclass of d0 WorkList [W] d1 : Data Maybe [t] Solve d1 using instance dfunData2; d1 := dfunData2 d2 d3 Inert: [G] d0 : Foo t [G] d01 : Data Maybe t -- Superclass of d0 Solved: d1 : Data Maybe [t] WorkList: [W] d2 : Sat (Maybe [t]) [W] d3 : Data Maybe t Now, we may simplify d2 using dfunSat; d2 := dfunSat d4 Inert: [G] d0 : Foo t [G] d01 : Data Maybe t -- Superclass of d0 Solved: d1 : Data Maybe [t] d2 : Sat (Maybe [t]) WorkList: [W] d3 : Data Maybe t [W] d4 : Foo [t] Now, we can just solve d3 from d01; d3 := d01 Inert [G] d0 : Foo t [G] d01 : Data Maybe t -- Superclass of d0 Solved: d1 : Data Maybe [t] d2 : Sat (Maybe [t]) WorkList [W] d4 : Foo [t] Now, solve d4 using dfunFoo2; d4 := dfunFoo2 d5 Inert [G] d0 : Foo t [G] d01 : Data Maybe t -- Superclass of d0 Solved: d1 : Data Maybe [t] d2 : Sat (Maybe [t]) d4 : Foo [t] WorkList: [W] d5 : Foo t Now, d5 can be solved! d5 := d0 Result d1 := dfunData2 d2 d3 d2 := dfunSat d4 d3 := d01 d4 := dfunFoo2 d5 d5 := d0 -} {- ********************************************************************* * * InertCans: the canonical inerts * * * * ********************************************************************* -} data InertCans -- See Note [Detailed InertCans Invariants] for more = IC { inert_eqs :: InertEqs -- See Note [inert_eqs: the inert equalities] -- All CEqCans with a TyVarLHS; index is the LHS tyvar -- Domain = skolems and untouchables; a touchable would be unified , inert_funeqs :: FunEqMap EqualCtList -- All CEqCans with a TyFamLHS; index is the whole family head type. -- LHS is fully rewritten (modulo eqCanRewrite constraints) -- wrt inert_eqs -- Can include all flavours, [G], [W], [WD], [D] , inert_dicts :: DictMap Ct -- Dictionaries only -- All fully rewritten (modulo flavour constraints) -- wrt inert_eqs , inert_insts :: [QCInst] , inert_safehask :: DictMap Ct -- Failed dictionary resolution due to Safe Haskell overlapping -- instances restriction. We keep this separate from inert_dicts -- as it doesn't cause compilation failure, just safe inference -- failure. -- -- ^ See Note [Safe Haskell Overlapping Instances Implementation] -- in "GHC.Tc.Solver" , inert_irreds :: Cts -- Irreducible predicates that cannot be made canonical, -- and which don't interact with others (e.g. (c a)) -- and insoluble predicates (e.g. Int ~ Bool, or a ~ [a]) , inert_given_eq_lvl :: TcLevel -- The TcLevel of the innermost implication that has a Given -- equality of the sort that make a unification variable untouchable -- (see Note [Unification preconditions] in GHC.Tc.Utils.Unify). -- See Note [Tracking Given equalities] below , inert_given_eqs :: Bool -- True <=> The inert Givens *at this level* (tcl_tclvl) -- could includes at least one equality /other than/ a -- let-bound skolem equality. -- Reason: report these givens when reporting a failed equality -- See Note [Tracking Given equalities] } type InertEqs = DTyVarEnv EqualCtList newtype EqualCtList = EqualCtList (NonEmpty Ct) deriving newtype Outputable -- See Note [EqualCtList invariants] unitEqualCtList :: Ct -> EqualCtList unitEqualCtList ct = EqualCtList (ct :| []) addToEqualCtList :: Ct -> EqualCtList -> EqualCtList -- NB: This function maintains the "derived-before-wanted" invariant of EqualCtList, -- but not the others. See Note [EqualCtList invariants] addToEqualCtList ct (EqualCtList old_eqs) | isWantedCt ct , eq1 :| eqs <- old_eqs = EqualCtList (eq1 :| ct : eqs) | otherwise = EqualCtList (ct `cons` old_eqs) filterEqualCtList :: (Ct -> Bool) -> EqualCtList -> Maybe EqualCtList filterEqualCtList pred (EqualCtList cts) = fmap EqualCtList (nonEmpty $ NE.filter pred cts) equalCtListToList :: EqualCtList -> [Ct] equalCtListToList (EqualCtList cts) = toList cts listToEqualCtList :: [Ct] -> Maybe EqualCtList -- NB: This does not maintain invariants other than having the EqualCtList be -- non-empty listToEqualCtList cts = EqualCtList <$> nonEmpty cts {- Note [Tracking Given equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ For reasons described in (UNTOUCHABLE) in GHC.Tc.Utils.Unify Note [Unification preconditions], we can't unify alpha[2] ~ Int under a level-4 implication if there are any Given equalities bound by the implications at level 3 of 4. To that end, the InertCans tracks inert_given_eq_lvl :: TcLevel -- The TcLevel of the innermost implication that has a Given -- equality of the sort that make a unification variable untouchable -- (see Note [Unification preconditions] in GHC.Tc.Utils.Unify). We update inert_given_eq_lvl whenever we add a Given to the inert set, in updateGivenEqs. Then a unification variable alpha[n] is untouchable iff n < inert_given_eq_lvl that is, if the unification variable was born outside an enclosing Given equality. Exactly which constraints should trigger (UNTOUCHABLE), and hence should update inert_given_eq_lvl? * We do /not/ need to worry about let-bound skolems, such ast forall[2] a. a ~ [b] => blah See Note [Let-bound skolems] * Consider an implication forall[2]. beta[1] => alpha[1] ~ Int where beta is a unification variable that has already been unified to () in an outer scope. Then alpha[1] is perfectly touchable and we can unify alpha := Int. So when deciding whether the givens contain an equality, we should canonicalise first, rather than just looking at the /original/ givens (#8644). * However, we must take account of *potential* equalities. Consider the same example again, but this time we have /not/ yet unified beta: forall[2] beta[1] => ...blah... Because beta might turn into an equality, updateGivenEqs conservatively treats it as a potential equality, and updates inert_give_eq_lvl * What about something like forall[2] a b. a ~ F b => [W] alpha[1] ~ X y z? That Given cannot affect the Wanted, because the Given is entirely *local*: it mentions only skolems bound in the very same implication. Such equalities need not make alpha untouchable. (Test case typecheck/should_compile/LocalGivenEqs has a real-life motivating example, with some detailed commentary.) Hence the 'mentionsOuterVar' test in updateGivenEqs. However, solely to support better error messages (see Note [HasGivenEqs] in GHC.Tc.Types.Constraint) we also track these "local" equalities in the boolean inert_given_eqs field. This field is used only to set the ic_given_eqs field to LocalGivenEqs; see the function getHasGivenEqs. Here is a simpler case that triggers this behaviour: data T where MkT :: F a ~ G b => a -> b -> T f (MkT _ _) = True Because of this behaviour around local equality givens, we can infer the type of f. This is typecheck/should_compile/LocalGivenEqs2. * We need not look at the equality relation involved (nominal vs representational), because representational equalities can still imply nominal ones. For example, if (G a ~R G b) and G's argument's role is nominal, then we can deduce a ~N b. Note [Let-bound skolems] ~~~~~~~~~~~~~~~~~~~~~~~~ If * the inert set contains a canonical Given CEqCan (a ~ ty) and * 'a' is a skolem bound in this very implication, then: a) The Given is pretty much a let-binding, like f :: (a ~ b->c) => a -> a Here the equality constraint is like saying let a = b->c in ... It is not adding any new, local equality information, and hence can be ignored by has_given_eqs b) 'a' will have been completely substituted out in the inert set, so we can safely discard it. For an example, see #9211. See also GHC.Tc.Utils.Unify Note [Deeper level on the left] for how we ensure that the right variable is on the left of the equality when both are tyvars. You might wonder whether the skolem really needs to be bound "in the very same implication" as the equuality constraint. Consider this (c.f. #15009): data S a where MkS :: (a ~ Int) => S a g :: forall a. S a -> a -> blah g x y = let h = \z. ( z :: Int , case x of MkS -> [y,z]) in ... From the type signature for `g`, we get `y::a` . Then when we encounter the `\z`, we'll assign `z :: alpha[1]`, say. Next, from the body of the lambda we'll get [W] alpha[1] ~ Int -- From z::Int [W] forall[2]. (a ~ Int) => [W] alpha[1] ~ a -- From [y,z] Now, unify alpha := a. Now we are stuck with an unsolved alpha~Int! So we must treat alpha as untouchable under the forall[2] implication. Note [Detailed InertCans Invariants] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The InertCans represents a collection of constraints with the following properties: * All canonical * No two dictionaries with the same head * No two CIrreds with the same type * Family equations inert wrt top-level family axioms * Dictionaries have no matching top-level instance * Given family or dictionary constraints don't mention touchable unification variables * Non-CEqCan constraints are fully rewritten with respect to the CEqCan equalities (modulo eqCanRewrite of course; eg a wanted cannot rewrite a given) * CEqCan equalities: see Note [inert_eqs: the inert equalities] Also see documentation in Constraint.Ct for a list of invariants Note [EqualCtList invariants] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ * All are equalities * All these equalities have the same LHS * The list is never empty * No element of the list can rewrite any other * Derived before Wanted From the fourth invariant it follows that the list is - A single [G], or - Zero or one [D] or [WD], followed by any number of [W] The Wanteds can't rewrite anything which is why we put them last Note [inert_eqs: 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 (See Note [Why R2?].) Lemma (L0). 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 (lhs -f-> t), where lhs is a type variable or an exactly-saturated type family application (that is, lhs is a CanEqLHS) t is a type f is a flavour such that (WF1) if (lhs1 -f1-> t1) in S (lhs2 -f2-> t2) in S then (f1 >= f2) implies that lhs1 does not appear within lhs2 (WF2) if (lhs -f-> t) is in S, then t /= lhs Definition [Applying a generalised substitution] If S is a generalised substitution S(f,t0) = t, if (t0 -fs-> t) in S, and fs >= f = apply S to components of t0, otherwise See also Note [Flavours with roles]. Theorem: S(f,t0) is well defined as a function. Proof: Suppose (lhs -f1-> t1) and (lhs -f2-> t2) are both in S, and f1 >= f and f2 >= f Then by (R2) f1 >= f2 or f2 >= f1, which contradicts (WF1) Notation: repeated application. S^0(f,t) = t S^(n+1)(f,t) = S(f, S^n(t)) Definition: terminating generalised substitution A generalised substitution S is *terminating* iff (IG1) there is an n such that for every f,t, S^n(f,t) = S^(n+1)(f,t) By (IG1) we define S*(f,t) to be the result of exahaustively applying S(f,_) to t. ----------------------------------------------------------------------------- Our main invariant: the CEqCans in inert_eqs should be a terminating generalised substitution ----------------------------------------------------------------------------- Note that termination is not the same as idempotence. To apply S to a type, you may have to apply it recursively. But termination does guarantee that this recursive use will terminate. Note [Why R2?] ~~~~~~~~~~~~~~ R2 states that, if we have f1 >= f and f2 >= f, then either f1 >= f2 or f2 >= f1. If we do not have R2, we will easily fall into a loop. To see why, imagine we have f1 >= f, f2 >= f, and that's it. Then, let our inert set S = {a -f1-> b, b -f2-> a}. Computing S(f,a) does not terminate. And yet, we have a hard time noticing an occurs-check problem when building S, as the two equalities cannot rewrite one another. R2 actually restricts our ability to accept user-written programs. See Note [Deriveds do rewrite Deriveds] in GHC.Tc.Types.Constraint for an example. Note [Rewritable] ~~~~~~~~~~~~~~~~~ This Note defines what it means for a type variable or type family application (that is, a CanEqLHS) to be rewritable in a type. This definition is used by the anyRewritableXXX family of functions and is meant to model the actual behaviour in GHC.Tc.Solver.Rewrite. Ignoring roles (for now): A CanEqLHS lhs is *rewritable* in a type t if the lhs tree appears as a subtree within t without traversing any of the following components of t: * coercions (whether they appear in casts CastTy or as arguments CoercionTy) * kinds of variable occurrences The check for rewritability *does* look in kinds of the bound variable of a ForAllTy. Goal: If lhs is not rewritable in t, then t is a fixpoint of the generalised substitution containing only {lhs -f*-> t'}, where f* is a flavour such that f* >= f for all f. The reason for this definition is that the rewriter does not rewrite in coercions or variables' kinds. In turn, the rewriter does not need to rewrite there because those places are never used for controlling the behaviour of the solver: these places are not used in matching instances or in decomposing equalities. There is one exception to the claim that non-rewritable parts of the tree do not affect the solver: we sometimes do an occurs-check to decide e.g. how to orient an equality. (See the comments on GHC.Tc.Solver.Canonical.canEqTyVarFunEq.) Accordingly, the presence of a variable in a kind or coercion just might influence the solver. Here is an example: type family Const x y where Const x y = x AxConst :: forall x y. Const x y ~# x alpha :: Const Type Nat [W] alpha ~ Int |> (sym (AxConst Type alpha) ;; AxConst Type alpha ;; sym (AxConst Type Nat)) The cast is clearly ludicrous (it ties together a cast and its symmetric version), but we can't quite rule it out. (See (EQ1) from Note [Respecting definitional equality] in GHC.Core.TyCo.Rep to see why we need the Const Type Nat bit.) And yet this cast will (quite rightly) prevent alpha from unifying with the RHS. I (Richard E) don't have an example of where this problem can arise from a Haskell program, but we don't have an air-tight argument for why the definition of *rewritable* given here is correct. Taking roles into account: we must consider a rewrite at a given role. That is, a rewrite arises from some equality, and that equality has a role associated with it. As we traverse a type, we track what role we are allowed to rewrite with. For example, suppose we have an inert [G] b ~R# Int. Then b is rewritable in Maybe b but not in F b, where F is a type function. This role-aware logic is present in both the anyRewritableXXX functions and in the rewriter. See also Note [anyRewritableTyVar must be role-aware] in GHC.Tc.Utils.TcType. Note [Extending the inert equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Main Theorem [Stability under extension] Suppose we have a "work item" lhs -fw-> t and a terminating generalised substitution S, THEN the extended substitution T = S+(lhs -fw-> t) is a terminating generalised substitution PROVIDED (T1) S(fw,lhs) = lhs -- 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) lhs not in t -- No occurs check in the work item -- If lhs is a type family application, we require only that -- lhs is not *rewritable* in t. See Note [Rewritable] and -- Note [CEqCan occurs check] in GHC.Tc.Types.Constraint. AND, for every (lhs1 -fs-> s) in S: (K0) not (fw >= fs) Reason: suppose we kick out (lhs1 -fs-> s), and add (lhs -fw-> t) to the inert set. The latter can't rewrite the former, so the kick-out achieved nothing -- From here, we can assume fw >= fs OR (K4) lhs1 is a tyvar AND fs >= fw OR { (K1) lhs is not rewritable in lhs1. See Note [Rewritable]. Reason: if fw >= fs, WF1 says we can't have both lhs0 -fw-> t and F lhs0 -fs-> s AND (K2): guarantees termination of the new substitution { (K2a) not (fs >= fs) OR (K2b) lhs not in s } AND (K3) See Note [K3: completeness of solving] { (K3a) If the role of fs is nominal: s /= lhs (K3b) If the role of fs is representational: s is not of form (lhs t1 .. tn) } } Conditions (T1-T3) are established by the canonicaliser Conditions (K1-K3) are established by GHC.Tc.Solver.Monad.kickOutRewritable The idea is that * T1 and T2 are guaranteed by exhaustively rewriting the work-item with S(fw,_). * T3 is guaranteed by an occurs-check on the work item. This is done during canonicalisation, in canEqOK and checkTypeEq; invariant (TyEq:OC) of CEqCan. See also Note [CEqCan occurs check] in GHC.Tc.Types.Constraint. * (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 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. See also (K2a) below. * Lemma (L2): if not (fw >= fw), then K0 holds and we kick out nothing Proof: using Definition [Can-rewrite relation], fw can't rewrite anything and so K0 holds. Intuitively, since fw can't rewrite anything (Lemma (L0)), adding it cannot cause any loops This is a common case, because Wanteds cannot rewrite Wanteds. It's used to avoid even looking for constraint to kick out. * Lemma (L1): The conditions of the Main Theorem imply that there is no (lhs -fs-> t) in S, s.t. (fs >= fw). Proof. Suppose the contrary (fs >= fw). Then because of (T1), S(fw,lhs)=lhs. But since fs>=fw, S(fw,lhs) = t, hence t=lhs. But now we have (lhs -fs-> lhs) in S, which contradicts (WF2). * The extended substitution satisfies (WF1) and (WF2) - (K1) plus (L1) guarantee that the extended substitution satisfies (WF1). - (T3) guarantees (WF2). * (K2) and (K4) are about termination. Intuitively, any infinite chain S^0(f,t), S^1(f,t), S^2(f,t).... must pass through the new work item infinitely often, since the substitution without the work item is terminating; and must pass through at least one of the triples in S infinitely often. - (K2a): if not(fs>=fs) then there is no f that fs can rewrite (fs>=f) (this is Lemma (L0)), and hence this triple never plays a role in application S(f,t). 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 lhs not in s, we have no further opportunity to apply the work item - (K4): See Note [K4] * Lemma (L3). Suppose we have f* such that, for all f, f* >= f. Then if we are adding lhs -fw-> t (where T1, T2, and T3 hold), we will keep a -f*-> s. Proof. K4 holds; thus, we keep. 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] Note [K4] ~~~~~~~~~ K4 is a "keep" condition of Note [Extending the inert equalities]. Here is the scenario: * We are considering adding (lhs -fw-> t) to the inert set S. * S already has (lhs1 -fs-> s). * We know S(fw, lhs) = lhs, S(fw, t) = t, and lhs is not rewritable in t. See Note [Rewritable]. These are (T1), (T2), and (T3). * We further know fw >= fs. (If not, then we short-circuit via (K0).) K4 says that we may keep lhs1 -fs-> s in S if: lhs1 is a tyvar AND fs >= fw Why K4 guarantees termination: * If fs >= fw, we know a is not rewritable in t, because of (T2). * We further know lhs /= a, because of (T1). * Accordingly, a use of the new inert item lhs -fw-> t cannot create the conditions for a use of a -fs-> s (precisely because t does not mention a), and hence, the extended substitution (with lhs -fw-> t in it) is a terminating generalised substitution. Recall that the termination generalised substitution includes only mappings that pass an occurs check. This is (T3). At one point, we worried that the argument here would fail if s mentioned a, but (T3) rules out this possibility. Put another way: the terminating generalised substitution considers only the inert_eqs, not other parts of the inert set (such as the irreds). Can we liberalise K4? No. Why we cannot drop the (fs >= fw) condition: * Suppose not (fs >= fw). It might be the case that t mentions a, and this can cause a loop. Example: Work: [G] b ~ a Inert: [D] a ~ b (where G >= G, G >= D, and D >= D) If we don't kick out the inert, then we get a loop on e.g. [D] a ~ Int. * Note that the above example is different if the inert is a Given G, because (T1) won't hold. Why we cannot drop the tyvar condition: * Presume fs >= fw. Thus, F tys is not rewritable in t, because of (T2). * Can the use of lhs -fw-> t create the conditions for a use of F tys -fs-> s? Yes! This can happen if t appears within tys. Here is an example: Work: [G] a ~ Int Inert: [G] F Int ~ F a Now, if we have [W] F a ~ Bool, we will rewrite ad infinitum on the left-hand side. The key reason why K2b works in the tyvar case is that tyvars are atomic: if the right-hand side of an equality does not mention a variable a, then it cannot allow an equality with an LHS of a to fire. This is not the case for type family applications. Bottom line: K4 can keep only inerts with tyvars on the left. Put differently, K4 will never prevent an inert with a type family on the left from being kicked out. Consequence: We never kick out a Given/Nominal equality with a tyvar on the left. This is Lemma (L3) of Note [Extending the inert equalities]. It is good because it means we can effectively model the mutable filling of metavariables with Given/Nominal equalities. That is: it should be the case that we could rewrite our solver never to fill in a metavariable; instead, it would "solve" a wanted like alpha ~ Int by turning it into a Given, allowing it to be used in rewriting. We would want the solver to behave the same whether it uses metavariables or Givens. And (L3) says that no Given/Nominals over tyvars are ever kicked out, just like we never unfill a metavariable. Nice. Getting this wrong (that is, allowing K4 to apply to situations with the type family on the left) led to #19042. (At that point, K4 was known as K2b.) Originally, this condition was part of K2, but #17672 suggests it should be a top-level K condition. Note [K3: completeness of solving] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ (K3) is not necessary for the extended substitution to be terminating. 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 terminating; 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 work-item c -G/N-> a inert-item a -W/R-> b c The work-item doesn't get rewritten by the inert, because (>=) doesn't hold. But we don't kick out the inert item because not (W/R >= W/R). So we just add the work item. But then, consider if we hit the following: work-item b -G/N-> Id inert-items a -W/R-> b c c -G/N-> a 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 lhs of a work item is "exposed", where exposed means being at the head of the top-level application chain (lhs t1 .. tn). See is_can_eq_lhs_head. This is encoded in (K3b). Beware: if we make this test succeed too often, we kick out too much, and the solver might loop. Consider (#14363) work item: [G] a ~R f b inert item: [G] b ~R f a In GHC 8.2 the completeness tests more aggressive, and kicked out the inert item; but no rewriting happened and there was an infinite loop. All we need is to have the tyvar at the head. Note [Flavours with roles] ~~~~~~~~~~~~~~~~~~~~~~~~~~ The system described in Note [inert_eqs: 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 [inert_eqs: the inert equalities], we must be careful to respect all components of a flavour. 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 substitution 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. -} instance Outputable InertCans where ppr (IC { inert_eqs = eqs , inert_funeqs = funeqs, inert_dicts = dicts , inert_safehask = safehask, inert_irreds = irreds , inert_given_eq_lvl = ge_lvl , inert_given_eqs = given_eqs , inert_insts = insts }) = braces $ vcat [ ppUnless (isEmptyDVarEnv eqs) $ text "Equalities:" <+> pprCts (foldDVarEnv folder emptyCts eqs) , ppUnless (isEmptyTcAppMap funeqs) $ text "Type-function equalities =" <+> pprCts (foldFunEqs folder funeqs emptyCts) , ppUnless (isEmptyTcAppMap dicts) $ text "Dictionaries =" <+> pprCts (dictsToBag dicts) , ppUnless (isEmptyTcAppMap safehask) $ text "Safe Haskell unsafe overlap =" <+> pprCts (dictsToBag safehask) , ppUnless (isEmptyCts irreds) $ text "Irreds =" <+> pprCts irreds , ppUnless (null insts) $ text "Given instances =" <+> vcat (map ppr insts) , text "Innermost given equalities =" <+> ppr ge_lvl , text "Given eqs at this level =" <+> ppr given_eqs ] where folder (EqualCtList eqs) rest = nonEmptyToBag eqs `andCts` rest {- ********************************************************************* * * Shadow constraints and improvement * * ************************************************************************ Note [The improvement story and derived shadows] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Because Wanteds cannot rewrite Wanteds (see Note [Wanteds do not rewrite Wanteds] in GHC.Tc.Types.Constraint), we may miss some opportunities for solving. Here's a classic example (indexed-types/should_fail/T4093a) Ambiguity check for f: (Foo e ~ Maybe e) => Foo e We get [G] Foo e ~ Maybe e (CEqCan) [W] Foo ee ~ Foo e (CEqCan) -- ee is a unification variable [W] Foo ee ~ Maybe ee (CEqCan) The first Wanted gets rewritten to [W] Foo ee ~ Maybe e But now we appear to be stuck, since we don't rewrite Wanteds with Wanteds. This is silly because we can see that ee := e is the only solution. The basic plan is * generate Derived constraints that shadow Wanted constraints * allow Derived to rewrite Derived * in order to cause some unifications to take place * that in turn solve the original Wanteds The ONLY reason for all these Derived equalities is to tell us how to unify a variable: that is, what Mark Jones calls "improvement". The same idea is sometimes also called "saturation"; find all the equalities that must hold in any solution. Or, equivalently, you can think of the derived shadows as implementing the "model": a non-idempotent but no-occurs-check substitution, reflecting *all* *Nominal* equalities (a ~N ty) that are not immediately soluble by unification. More specifically, here's how it works (Oct 16): * Wanted constraints are born as [WD]; this behaves like a [W] and a [D] paired together. * When we are about to add a [WD] to the inert set, if it can be rewritten by a [D] a ~ ty, then we split it into [W] and [D], putting the latter into the work list (see maybeEmitShadow). In the example above, we get to the point where we are stuck: [WD] Foo ee ~ Foo e [WD] Foo ee ~ Maybe ee But now when [WD] Foo ee ~ Maybe ee is about to be added, we'll split it into [W] and [D], since the inert [WD] Foo ee ~ Foo e can rewrite it. Then: work item: [D] Foo ee ~ Maybe ee inert: [W] Foo ee ~ Maybe ee [WD] Foo ee ~ Maybe e See Note [Splitting WD constraints]. Now the work item is rewritten by the [WD] and we soon get ee := e. Additional notes: * The derived shadow equalities live in inert_eqs, along with the Givens and Wanteds; see Note [EqualCtList invariants]. * We make Derived shadows only for Wanteds, not Givens. So we have only [G], not [GD] and [G] plus splitting. See Note [Add derived shadows only for Wanteds] * We also get Derived equalities from functional dependencies and type-function injectivity; see calls to unifyDerived. * It's worth having [WD] rather than just [W] and [D] because * efficiency: silly to process the same thing twice * inert_dicts is a finite map keyed by the type; it's inconvenient for it to map to TWO constraints Another example requiring Deriveds is in Note [Put touchable variables on the left] in GHC.Tc.Solver.Canonical. Note [Splitting WD constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We are about to add a [WD] constraint to the inert set; and we know that the inert set has fully rewritten it. Should we split it into [W] and [D], and put the [D] in the work list for further work? * CDictCan (C tys): Yes if the inert set could rewrite tys to make the class constraint, or type family, fire. That is, yes if the inert_eqs intersects with the free vars of tys. For this test we use (anyRewritableTyVar True) which ignores casts and coercions in tys, because rewriting the casts or coercions won't make the thing fire more often. * CEqCan (lhs ~ ty): Yes if the inert set could rewrite 'lhs' or 'ty'. We need to check both 'lhs' and 'ty' against the inert set: - Inert set contains [D] a ~ ty2 Then we want to put [D] a ~ ty in the worklist, so we'll get [D] ty ~ ty2 with consequent good things - Inert set contains [D] b ~ a, where b is in ty. We can't just add [WD] a ~ ty[b] to the inert set, because that breaks the inert-set invariants. If we tried to canonicalise another [D] constraint mentioning 'a', we'd get an infinite loop Moreover we must use (anyRewritableTyVar False) for the RHS, because even tyvars in the casts and coercions could give an infinite loop if we don't expose it * CIrredCan: Yes if the inert set can rewrite the constraint. We used to think splitting irreds was unnecessary, but see Note [Splitting Irred WD constraints] * Others: nothing is gained by splitting. Note [Splitting Irred WD constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Splitting Irred constraints can make a difference. Here is the scenario: a[sk] :: F v -- F is a type family beta :: alpha work item: [WD] a ~ beta This is heterogeneous, so we emit a kind equality and make the work item an inert Irred. work item: [D] F v ~ alpha inert: [WD] (a |> co) ~ beta (CIrredCan) Can't make progress on the work item. Add to inert set. This kicks out the old inert, because a [D] can rewrite a [WD]. work item: [WD] (a |> co) ~ beta inert: [D] F v ~ alpha (CEqCan) Can't make progress on this work item either (although GHC tries by decomposing the cast and rewriting... but that doesn't make a difference), which is still hetero. Emit a new kind equality and add to inert set. But, critically, we split the Irred. work list: [D] F v ~ alpha (CEqCan) [D] (a |> co) ~ beta (CIrred) -- this one was split off inert: [W] (a |> co) ~ beta [D] F v ~ alpha We quickly solve the first work item, as it's the same as an inert. work item: [D] (a |> co) ~ beta inert: [W] (a |> co) ~ beta [D] F v ~ alpha We decompose the cast, yielding [D] a ~ beta We then rewrite the kinds. The lhs kind is F v, which flattens to alpha. co' :: F v ~ alpha [D] (a |> co') ~ beta Now this equality is homo-kinded. So we swizzle it around to [D] beta ~ (a |> co') and set beta := a |> co', and go home happy. If we don't split the Irreds, we loop. This is all dangerously subtle. This is triggered by test case typecheck/should_compile/SplitWD. Note [Add derived shadows only for Wanteds] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We only add shadows for Wanted constraints. That is, we have [WD] but not [GD]; and maybeEmitShaodw looks only at [WD] constraints. It does just possibly make sense ot add a derived shadow for a Given. If we created a Derived shadow of a Given, it could be rewritten by other Deriveds, and that could, conceivably, lead to a useful unification. But (a) I have been unable to come up with an example of this happening (b) see #12660 for how adding the derived shadows of a Given led to an infinite loop. (c) It's unlikely that rewriting derived Givens will lead to a unification because Givens don't mention touchable unification variables For (b) there may be other ways to solve the loop, but simply reraining from adding derived shadows of Givens is particularly simple. And it's more efficient too! Still, here's one possible reason for adding derived shadows for Givens. Consider work-item [G] a ~ [b], inerts has [D] b ~ a. If we added the derived shadow (into the work list) [D] a ~ [b] When we process it, we'll rewrite to a ~ [a] and get an occurs check. Without it we'll miss the occurs check (reporting inaccessible code); but that's probably OK. Note [Keep CDictCan shadows as CDictCan] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Suppose we have class C a => D a b and [G] D a b, [G] C a in the inert set. Now we insert [D] b ~ c. We want to kick out a derived shadow for [D] D a b, so we can rewrite it with the new constraint, and perhaps get instance reduction or other consequences. BUT we do not want to kick out a *non-canonical* (D a b). If we did, we would do this: - rewrite it to [D] D a c, with pend_sc = True - use expandSuperClasses to add C a - go round again, which solves C a from the givens This loop goes on for ever and triggers the simpl_loop limit. Solution: kick out the CDictCan which will have pend_sc = False, because we've already added its superclasses. So we won't re-add them. If we forget the pend_sc flag, our cunning scheme for avoiding generating superclasses repeatedly will fail. See #11379 for a case of this. Note [Do not do improvement for WOnly] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We do improvement between two constraints (e.g. for injectivity or functional dependencies) only if both are "improvable". And we improve a constraint wrt the top-level instances only if it is improvable. Improvable: [G] [WD] [D} Not improvable: [W] Reasons: * It's less work: fewer pairs to compare * Every [W] has a shadow [D] so nothing is lost * Consider [WD] C Int b, where 'b' is a skolem, and class C a b | a -> b instance C Int Bool We'll do a fundep on it and emit [D] b ~ Bool That will kick out constraint [WD] C Int b Then we'll split it to [W] C Int b (keep in inert) and [D] C Int b (in work list) When processing the latter we'll rewrite it to [D] C Int Bool At that point it would be /stupid/ to interact it with the inert [W] C Int b in the inert set; after all, it's the very constraint from which the [D] C Int Bool was split! We can avoid this by not doing improvement on [W] constraints. This came up in #12860. -} maybeEmitShadow :: InertCans -> Ct -> TcS Ct -- See Note [The improvement story and derived shadows] maybeEmitShadow ics ct | let ev = ctEvidence ct , CtWanted { ctev_pred = pred, ctev_loc = loc , ctev_nosh = WDeriv } <- ev , shouldSplitWD (inert_eqs ics) (inert_funeqs ics) ct = do { traceTcS "Emit derived shadow" (ppr ct) ; let derived_ev = CtDerived { ctev_pred = pred , ctev_loc = loc } shadow_ct = ct { cc_ev = derived_ev } -- Te shadow constraint keeps the canonical shape. -- This just saves work, but is sometimes important; -- see Note [Keep CDictCan shadows as CDictCan] ; emitWork [shadow_ct] ; let ev' = ev { ctev_nosh = WOnly } ct' = ct { cc_ev = ev' } -- Record that it now has a shadow -- This is /the/ place we set the flag to WOnly ; return ct' } | otherwise = return ct shouldSplitWD :: InertEqs -> FunEqMap EqualCtList -> Ct -> Bool -- Precondition: 'ct' is [WD], and is inert -- True <=> we should split ct ito [W] and [D] because -- the inert_eqs can make progress on the [D] -- See Note [Splitting WD constraints] shouldSplitWD inert_eqs fun_eqs (CDictCan { cc_tyargs = tys }) = should_split_match_args inert_eqs fun_eqs tys -- NB True: ignore coercions -- See Note [Splitting WD constraints] shouldSplitWD inert_eqs fun_eqs (CEqCan { cc_lhs = TyVarLHS tv, cc_rhs = ty , cc_eq_rel = eq_rel }) = tv `elemDVarEnv` inert_eqs || anyRewritableCanEqLHS eq_rel (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs) ty -- NB False: do not ignore casts and coercions -- See Note [Splitting WD constraints] shouldSplitWD inert_eqs fun_eqs (CEqCan { cc_ev = ev, cc_eq_rel = eq_rel }) = anyRewritableCanEqLHS eq_rel (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs) (ctEvPred ev) shouldSplitWD inert_eqs fun_eqs (CIrredCan { cc_ev = ev }) = anyRewritableCanEqLHS (ctEvEqRel ev) (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs) (ctEvPred ev) shouldSplitWD _ _ _ = False -- No point in splitting otherwise should_split_match_args :: InertEqs -> FunEqMap EqualCtList -> [TcType] -> Bool -- True if the inert_eqs can rewrite anything in the argument types should_split_match_args inert_eqs fun_eqs tys = any (anyRewritableCanEqLHS NomEq (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs)) tys -- See Note [Splitting WD constraints] canRewriteTv :: InertEqs -> EqRel -> TyVar -> Bool canRewriteTv inert_eqs eq_rel tv | Just (EqualCtList (ct :| _)) <- lookupDVarEnv inert_eqs tv , CEqCan { cc_eq_rel = eq_rel1 } <- ct = eq_rel1 `eqCanRewrite` eq_rel | otherwise = False canRewriteTyFam :: FunEqMap EqualCtList -> EqRel -> TyCon -> [Type] -> Bool canRewriteTyFam fun_eqs eq_rel tf args | Just (EqualCtList (ct :| _)) <- findFunEq fun_eqs tf args , CEqCan { cc_eq_rel = eq_rel1 } <- ct = eq_rel1 `eqCanRewrite` eq_rel | otherwise = False isImprovable :: CtEvidence -> Bool -- See Note [Do not do improvement for WOnly] isImprovable (CtWanted { ctev_nosh = WOnly }) = False isImprovable _ = True {- ********************************************************************* * * Inert equalities * * ********************************************************************* -} addTyEq :: InertEqs -> TcTyVar -> Ct -> InertEqs addTyEq old_eqs tv ct = extendDVarEnv_C add_eq old_eqs tv (unitEqualCtList ct) where add_eq old_eqs _ = addToEqualCtList ct old_eqs addCanFunEq :: FunEqMap EqualCtList -> TyCon -> [TcType] -> Ct -> FunEqMap EqualCtList addCanFunEq old_eqs fun_tc fun_args ct = alterTcApp old_eqs fun_tc fun_args upd where upd (Just old_equal_ct_list) = Just $ addToEqualCtList ct old_equal_ct_list upd Nothing = Just $ unitEqualCtList ct foldTyEqs :: (Ct -> b -> b) -> InertEqs -> b -> b foldTyEqs k eqs z = foldDVarEnv (\(EqualCtList cts) z -> foldr k z cts) z eqs findTyEqs :: InertCans -> TyVar -> [Ct] findTyEqs icans tv = maybe [] id (fmap @Maybe equalCtListToList $ lookupDVarEnv (inert_eqs icans) tv) delEq :: InertCans -> CanEqLHS -> TcType -> InertCans delEq ic lhs rhs = case lhs of TyVarLHS tv -> ic { inert_eqs = alterDVarEnv upd (inert_eqs ic) tv } TyFamLHS tf args -> ic { inert_funeqs = alterTcApp (inert_funeqs ic) tf args upd } where isThisOne :: Ct -> Bool isThisOne (CEqCan { cc_rhs = t1 }) = tcEqTypeNoKindCheck rhs t1 isThisOne other = pprPanic "delEq" (ppr lhs $$ ppr ic $$ ppr other) upd :: Maybe EqualCtList -> Maybe EqualCtList upd (Just eq_ct_list) = filterEqualCtList (not . isThisOne) eq_ct_list upd Nothing = Nothing findEq :: InertCans -> CanEqLHS -> [Ct] findEq icans (TyVarLHS tv) = findTyEqs icans tv findEq icans (TyFamLHS fun_tc fun_args) = maybe [] id (fmap @Maybe equalCtListToList $ findFunEq (inert_funeqs icans) fun_tc fun_args) {- ********************************************************************* * * Inert instances: inert_insts * * ********************************************************************* -} addInertForAll :: QCInst -> TcS () -- Add a local Given instance, typically arising from a type signature addInertForAll new_qci = do { ics <- getInertCans ; ics1 <- add_qci ics -- Update given equalities. C.f updateGivenEqs ; tclvl <- getTcLevel ; let pred = qci_pred new_qci not_equality = isClassPred pred && not (isEqPred pred) -- True <=> definitely not an equality -- A qci_pred like (f a) might be an equality ics2 | not_equality = ics1 | otherwise = ics1 { inert_given_eq_lvl = tclvl , inert_given_eqs = True } ; setInertCans ics2 } where add_qci :: InertCans -> TcS InertCans -- See Note [Do not add duplicate quantified instances] add_qci ics@(IC { inert_insts = qcis }) | any same_qci qcis = do { traceTcS "skipping duplicate quantified instance" (ppr new_qci) ; return ics } | otherwise = do { traceTcS "adding new inert quantified instance" (ppr new_qci) ; return (ics { inert_insts = new_qci : qcis }) } same_qci old_qci = tcEqType (ctEvPred (qci_ev old_qci)) (ctEvPred (qci_ev new_qci)) {- Note [Do not add duplicate quantified instances] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider this (#15244): f :: (C g, D g) => .... class S g => C g where ... class S g => D g where ... class (forall a. Eq a => Eq (g a)) => S g where ... Then in f's RHS there are two identical quantified constraints available, one via the superclasses of C and one via the superclasses of D. The two are identical, and it seems wrong to reject the program because of that. But without doing duplicate-elimination we will have two matching QCInsts when we try to solve constraints arising from f's RHS. The simplest thing is simply to eliminate duplicates, which we do here. -} {- ********************************************************************* * * Adding an inert * * ************************************************************************ Note [Adding an equality to the InertCans] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ When adding an equality to the inerts: * Split [WD] into [W] and [D] if the inerts can rewrite the latter; done by maybeEmitShadow. * Kick out any constraints that can be rewritten by the thing we are adding. Done by kickOutRewritable. * Note that unifying a:=ty, is like adding [G] a~ty; just use kickOutRewritable with Nominal, Given. See kickOutAfterUnification. -} addInertCan :: Ct -> TcS () -- Precondition: item /is/ canonical -- See Note [Adding an equality to the InertCans] addInertCan ct = do { traceTcS "addInertCan {" $ text "Trying to insert new inert item:" <+> ppr ct ; ics <- getInertCans ; ct <- maybeEmitShadow ics ct ; ics <- maybeKickOut ics ct ; tclvl <- getTcLevel ; setInertCans (add_item tclvl ics ct) ; traceTcS "addInertCan }" $ empty } maybeKickOut :: InertCans -> Ct -> TcS InertCans -- For a CEqCan, kick out any inert that can be rewritten by the CEqCan maybeKickOut ics ct | CEqCan { cc_lhs = lhs, cc_ev = ev, cc_eq_rel = eq_rel } <- ct = do { (_, ics') <- kickOutRewritable (ctEvFlavour ev, eq_rel) lhs ics ; return ics' } | otherwise = return ics add_item :: TcLevel -> InertCans -> Ct -> InertCans add_item tc_lvl ics@(IC { inert_funeqs = funeqs, inert_eqs = eqs }) item@(CEqCan { cc_lhs = lhs }) = updateGivenEqs tc_lvl item $ case lhs of TyFamLHS tc tys -> ics { inert_funeqs = addCanFunEq funeqs tc tys item } TyVarLHS tv -> ics { inert_eqs = addTyEq eqs tv item } add_item tc_lvl ics@(IC { inert_irreds = irreds }) item@(CIrredCan {}) = updateGivenEqs tc_lvl item $ -- An Irred might turn out to be an -- equality, so we play safe ics { inert_irreds = irreds `Bag.snocBag` item } add_item _ ics item@(CDictCan { cc_class = cls, cc_tyargs = tys }) = ics { inert_dicts = addDictCt (inert_dicts ics) cls tys item } add_item _ _ item = pprPanic "upd_inert set: can't happen! Inserting " $ ppr item -- Can't be CNonCanonical because they only land in inert_irreds updateGivenEqs :: TcLevel -> Ct -> InertCans -> InertCans -- Set the inert_given_eq_level to the current level (tclvl) -- if the constraint is a given equality that should prevent -- filling in an outer unification variable. -- See See Note [Tracking Given equalities] updateGivenEqs tclvl ct inerts@(IC { inert_given_eq_lvl = ge_lvl }) | not (isGivenCt ct) = inerts | not_equality ct = inerts -- See Note [Let-bound skolems] | otherwise = inerts { inert_given_eq_lvl = ge_lvl' , inert_given_eqs = True } where ge_lvl' | mentionsOuterVar tclvl (ctEvidence ct) -- Includes things like (c a), which *might* be an equality = tclvl | otherwise = ge_lvl not_equality :: Ct -> Bool -- True <=> definitely not an equality of any kind -- except for a let-bound skolem, which doesn't count -- See Note [Let-bound skolems] -- NB: no need to spot the boxed CDictCan (a ~ b) because its -- superclass (a ~# b) will be a CEqCan not_equality (CEqCan { cc_lhs = TyVarLHS tv }) = not (isOuterTyVar tclvl tv) not_equality (CDictCan {}) = True not_equality _ = False ----------------------------------------- kickOutRewritable :: CtFlavourRole -- Flavour/role of the equality that -- is being added to the inert set -> CanEqLHS -- The new equality is lhs ~ ty -> InertCans -> TcS (Int, InertCans) kickOutRewritable new_fr new_lhs ics = do { let (kicked_out, ics') = kick_out_rewritable new_fr new_lhs ics n_kicked = workListSize kicked_out ; unless (n_kicked == 0) $ do { updWorkListTcS (appendWorkList kicked_out) -- The famapp-cache contains Given evidence from the inert set. -- If we're kicking out Givens, we need to remove this evidence -- from the cache, too. ; let kicked_given_ev_vars = [ ev_var | ct <- wl_eqs kicked_out , CtGiven { ctev_evar = ev_var } <- [ctEvidence ct] ] ; when (new_fr `eqCanRewriteFR` (Given, NomEq) && -- if this isn't true, no use looking through the constraints not (null kicked_given_ev_vars)) $ do { traceTcS "Given(s) have been kicked out; drop from famapp-cache" (ppr kicked_given_ev_vars) ; dropFromFamAppCache (mkVarSet kicked_given_ev_vars) } ; csTraceTcS $ hang (text "Kick out, lhs =" <+> ppr new_lhs) 2 (vcat [ text "n-kicked =" <+> int n_kicked , text "kicked_out =" <+> ppr kicked_out , text "Residual inerts =" <+> ppr ics' ]) } ; return (n_kicked, ics') } kick_out_rewritable :: CtFlavourRole -- Flavour/role of the equality that -- is being added to the inert set -> CanEqLHS -- The new equality is lhs ~ ty -> InertCans -> (WorkList, InertCans) -- See Note [kickOutRewritable] kick_out_rewritable new_fr new_lhs ics@(IC { inert_eqs = tv_eqs , inert_dicts = dictmap , inert_funeqs = funeqmap , inert_irreds = irreds , inert_insts = old_insts }) | not (new_fr `eqMayRewriteFR` new_fr) = (emptyWorkList, ics) -- If new_fr can't rewrite itself, it can't rewrite -- anything else, so no need to kick out anything. -- (This is a common case: wanteds can't rewrite wanteds) -- Lemma (L2) in Note [Extending the inert equalities] | otherwise = (kicked_out, inert_cans_in) where -- inert_safehask stays unchanged; is that right? inert_cans_in = ics { inert_eqs = tv_eqs_in , inert_dicts = dicts_in , inert_funeqs = feqs_in , inert_irreds = irs_in , inert_insts = insts_in } kicked_out :: WorkList -- NB: use extendWorkList to ensure that kicked-out equalities get priority -- See Note [Prioritise equalities] (Kick-out). -- The irreds may include non-canonical (hetero-kinded) equality -- constraints, which perhaps may have become soluble after new_lhs -- is substituted; ditto the dictionaries, which may include (a~b) -- or (a~~b) constraints. kicked_out = foldr extendWorkListCt (emptyWorkList { wl_eqs = tv_eqs_out ++ feqs_out }) ((dicts_out `andCts` irs_out) `extendCtsList` insts_out) (tv_eqs_out, tv_eqs_in) = foldDVarEnv (kick_out_eqs extend_tv_eqs) ([], emptyDVarEnv) tv_eqs (feqs_out, feqs_in) = foldFunEqs (kick_out_eqs extend_fun_eqs) funeqmap ([], emptyFunEqs) (dicts_out, dicts_in) = partitionDicts kick_out_ct dictmap (irs_out, irs_in) = partitionBag kick_out_ct irreds -- Kick out even insolubles: See Note [Rewrite insolubles] -- Of course we must kick out irreducibles like (c a), in case -- we can rewrite 'c' to something more useful -- Kick-out for inert instances -- See Note [Quantified constraints] in GHC.Tc.Solver.Canonical insts_out :: [Ct] insts_in :: [QCInst] (insts_out, insts_in) | fr_may_rewrite (Given, NomEq) -- All the insts are Givens = partitionWith kick_out_qci old_insts | otherwise = ([], old_insts) kick_out_qci qci | let ev = qci_ev qci , fr_can_rewrite_ty NomEq (ctEvPred (qci_ev qci)) = Left (mkNonCanonical ev) | otherwise = Right qci (_, new_role) = new_fr fr_tv_can_rewrite_ty :: TyVar -> EqRel -> Type -> Bool fr_tv_can_rewrite_ty new_tv role ty = anyRewritableTyVar True role can_rewrite ty -- True: ignore casts and coercions where can_rewrite :: EqRel -> TyVar -> Bool can_rewrite old_role tv = new_role `eqCanRewrite` old_role && tv == new_tv fr_tf_can_rewrite_ty :: TyCon -> [TcType] -> EqRel -> Type -> Bool fr_tf_can_rewrite_ty new_tf new_tf_args role ty = anyRewritableTyFamApp role can_rewrite ty where can_rewrite :: EqRel -> TyCon -> [TcType] -> Bool can_rewrite old_role old_tf old_tf_args = new_role `eqCanRewrite` old_role && tcEqTyConApps new_tf new_tf_args old_tf old_tf_args -- it's possible for old_tf_args to have too many. This is fine; -- we'll only check what we need to. {-# INLINE fr_can_rewrite_ty #-} -- perform the check here only once fr_can_rewrite_ty :: EqRel -> Type -> Bool fr_can_rewrite_ty = case new_lhs of TyVarLHS new_tv -> fr_tv_can_rewrite_ty new_tv TyFamLHS new_tf new_tf_args -> fr_tf_can_rewrite_ty new_tf new_tf_args fr_may_rewrite :: CtFlavourRole -> Bool fr_may_rewrite fs = new_fr `eqMayRewriteFR` fs -- Can the new item rewrite the inert item? {-# INLINE kick_out_ct #-} -- perform case on new_lhs here only once kick_out_ct :: Ct -> Bool -- Kick it out if the new CEqCan can rewrite the inert one -- See Note [kickOutRewritable] kick_out_ct = case new_lhs of TyVarLHS new_tv -> \ct -> let fs@(_,role) = ctFlavourRole ct in fr_may_rewrite fs && fr_tv_can_rewrite_ty new_tv role (ctPred ct) TyFamLHS new_tf new_tf_args -> \ct -> let fs@(_, role) = ctFlavourRole ct in fr_may_rewrite fs && fr_tf_can_rewrite_ty new_tf new_tf_args role (ctPred ct) extend_tv_eqs :: InertEqs -> CanEqLHS -> EqualCtList -> InertEqs extend_tv_eqs eqs (TyVarLHS tv) cts = extendDVarEnv eqs tv cts extend_tv_eqs eqs other _cts = pprPanic "extend_tv_eqs" (ppr eqs $$ ppr other) extend_fun_eqs :: FunEqMap EqualCtList -> CanEqLHS -> EqualCtList -> FunEqMap EqualCtList extend_fun_eqs eqs (TyFamLHS fam_tc fam_args) cts = insertTcApp eqs fam_tc fam_args cts extend_fun_eqs eqs other _cts = pprPanic "extend_fun_eqs" (ppr eqs $$ ppr other) kick_out_eqs :: (container -> CanEqLHS -> EqualCtList -> container) -> EqualCtList -> ([Ct], container) -> ([Ct], container) kick_out_eqs extend eqs (acc_out, acc_in) = (eqs_out `chkAppend` acc_out, case listToEqualCtList eqs_in of Nothing -> acc_in Just eqs_in_ecl@(EqualCtList (eq1 :| _)) -> extend acc_in (cc_lhs eq1) eqs_in_ecl) where (eqs_out, eqs_in) = partition kick_out_eq (equalCtListToList eqs) -- Implements criteria K1-K3 in Note [Extending the inert equalities] kick_out_eq (CEqCan { cc_lhs = lhs, cc_rhs = rhs_ty , cc_ev = ev, cc_eq_rel = eq_rel }) | not (fr_may_rewrite fs) = False -- (K0) Keep it in the inert set if the new thing can't rewrite it -- Below here (fr_may_rewrite fs) is True | TyVarLHS _ <- lhs , fs `eqMayRewriteFR` new_fr = False -- (K4) Keep it in the inert set if the LHS is a tyvar and -- it can rewrite the work item. See Note [K4] | fr_can_rewrite_ty eq_rel (canEqLHSType lhs) = True -- (K1) -- The above check redundantly checks the role & flavour, -- but it's very convenient | kick_out_for_inertness = True -- (K2) | kick_out_for_completeness = True -- (K3) | otherwise = False where fs = (ctEvFlavour ev, eq_rel) kick_out_for_inertness = (fs `eqMayRewriteFR` fs) -- (K2a) && fr_can_rewrite_ty eq_rel rhs_ty -- (K2b) kick_out_for_completeness -- (K3) and Note [K3: completeness of solving] = case eq_rel of NomEq -> rhs_ty `eqType` canEqLHSType new_lhs -- (K3a) ReprEq -> is_can_eq_lhs_head new_lhs rhs_ty -- (K3b) kick_out_eq ct = pprPanic "keep_eq" (ppr ct) is_can_eq_lhs_head (TyVarLHS tv) = go where go (Rep.TyVarTy tv') = tv == tv' go (Rep.AppTy fun _) = go fun go (Rep.CastTy ty _) = go ty go (Rep.TyConApp {}) = False go (Rep.LitTy {}) = False go (Rep.ForAllTy {}) = False go (Rep.FunTy {}) = False go (Rep.CoercionTy {}) = False is_can_eq_lhs_head (TyFamLHS fun_tc fun_args) = go where go (Rep.TyVarTy {}) = False go (Rep.AppTy {}) = False -- no TyConApp to the left of an AppTy go (Rep.CastTy ty _) = go ty go (Rep.TyConApp tc args) = tcEqTyConApps fun_tc fun_args tc args go (Rep.LitTy {}) = False go (Rep.ForAllTy {}) = False go (Rep.FunTy {}) = False go (Rep.CoercionTy {}) = False kickOutAfterUnification :: TcTyVar -> TcS Int kickOutAfterUnification new_tv = do { ics <- getInertCans ; (n_kicked, ics2) <- kickOutRewritable (Given,NomEq) (TyVarLHS new_tv) ics -- Given because the tv := xi is given; NomEq because -- only nominal equalities are solved by unification ; setInertCans ics2 ; return n_kicked } -- See Wrinkle (2) in Note [Equalities with incompatible kinds] in GHC.Tc.Solver.Canonical kickOutAfterFillingCoercionHole :: CoercionHole -> Coercion -> TcS () kickOutAfterFillingCoercionHole hole filled_co = do { ics <- getInertCans ; let (kicked_out, ics') = kick_out ics n_kicked = workListSize kicked_out ; unless (n_kicked == 0) $ do { updWorkListTcS (appendWorkList kicked_out) ; csTraceTcS $ hang (text "Kick out, hole =" <+> ppr hole) 2 (vcat [ text "n-kicked =" <+> int n_kicked , text "kicked_out =" <+> ppr kicked_out , text "Residual inerts =" <+> ppr ics' ]) } ; setInertCans ics' } where holes_of_co = coercionHolesOfCo filled_co kick_out :: InertCans -> (WorkList, InertCans) kick_out ics@(IC { inert_irreds = irreds }) = let (to_kick, to_keep) = partitionBagWith kick_ct irreds kicked_out = extendWorkListCts (bagToList to_kick) emptyWorkList ics' = ics { inert_irreds = to_keep } in (kicked_out, ics') kick_ct :: Ct -> Either Ct Ct -- Left: kick out; Right: keep. But even if we keep, we may need -- to update the set of blocking holes kick_ct ct@(CIrredCan { cc_status = BlockedCIS holes }) | hole `elementOfUniqSet` holes = let new_holes = holes `delOneFromUniqSet` hole `unionUniqSets` holes_of_co updated_ct = ct { cc_status = BlockedCIS new_holes } in if isEmptyUniqSet new_holes then Left updated_ct else Right updated_ct kick_ct other = Right other {- Note [kickOutRewritable] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ See also Note [inert_eqs: the inert equalities]. When we add a new inert equality (lhs ~N ty) to the inert set, we must kick out any inert items that could be rewritten by the new equality, to maintain the inert-set invariants. - We want to kick out an existing inert constraint if a) the new constraint can rewrite the inert one b) 'lhs' is free in the inert constraint (so that it *will*) rewrite it if we kick it out. For (b) we use anyRewritableCanLHS, which examines the types /and kinds/ that are directly visible in the type. Hence we will have exposed all the rewriting we care about to make the most precise kinds visible for matching classes etc. No need to kick out constraints that mention type variables whose kinds contain this LHS! - A Derived equality can kick out [D] constraints in inert_eqs, inert_dicts, inert_irreds etc. - We don't kick out constraints from inert_solved_dicts, and inert_solved_funeqs optimistically. But when we lookup we have to take the substitution into account NB: we could in principle avoid kick-out: a) When unifying a meta-tyvar from an outer level, because then the entire implication will be iterated; see Note [The Unification Level Flag] b) For Givens, after a unification. By (GivenInv) in GHC.Tc.Utils.TcType Note [TcLevel invariants], a Given can't include a meta-tyvar from its own level, so it falls under (a). Of course, we must still kick out Givens when adding a new non-unification Given. But kicking out more vigorously may lead to earlier unification and fewer iterations, so we don't take advantage of these possibilities. Note [Rewrite insolubles] ~~~~~~~~~~~~~~~~~~~~~~~~~ Suppose we have an insoluble alpha ~ [alpha], which is insoluble because an occurs check. And then we unify alpha := [Int]. Then we really want to rewrite the insoluble to [Int] ~ [[Int]]. Now it can be decomposed. Otherwise we end up with a "Can't match [Int] ~ [[Int]]" which is true, but a bit confusing because the outer type constructors match. Hence: * In the main simplifier loops in GHC.Tc.Solver (solveWanteds, simpl_loop), we feed the insolubles in solveSimpleWanteds, so that they get rewritten (albeit not solved). * We kick insolubles out of the inert set, if they can be rewritten (see GHC.Tc.Solver.Monad.kick_out_rewritable) * We rewrite those insolubles in GHC.Tc.Solver.Canonical. See Note [Make sure that insolubles are fully rewritten] -} -------------- addInertSafehask :: InertCans -> Ct -> InertCans addInertSafehask ics item@(CDictCan { cc_class = cls, cc_tyargs = tys }) = ics { inert_safehask = addDictCt (inert_dicts ics) cls tys item } addInertSafehask _ item = pprPanic "addInertSafehask: can't happen! Inserting " $ ppr item insertSafeOverlapFailureTcS :: InstanceWhat -> Ct -> TcS () -- See Note [Safe Haskell Overlapping Instances Implementation] in GHC.Tc.Solver insertSafeOverlapFailureTcS what item | safeOverlap what = return () | otherwise = updInertCans (\ics -> addInertSafehask ics item) getSafeOverlapFailures :: TcS Cts -- See Note [Safe Haskell Overlapping Instances Implementation] in GHC.Tc.Solver getSafeOverlapFailures = do { IC { inert_safehask = safehask } <- getInertCans ; return $ foldDicts consCts safehask emptyCts } -------------- addSolvedDict :: InstanceWhat -> CtEvidence -> Class -> [Type] -> TcS () -- Conditionally add a new item in the solved set of the monad -- See Note [Solved dictionaries] addSolvedDict what item cls tys | isWanted item , instanceReturnsDictCon what = do { traceTcS "updSolvedSetTcs:" $ ppr item ; updInertTcS $ \ ics -> ics { inert_solved_dicts = addDict (inert_solved_dicts ics) cls tys item } } | otherwise = return () getSolvedDicts :: TcS (DictMap CtEvidence) getSolvedDicts = do { ics <- getTcSInerts; return (inert_solved_dicts ics) } setSolvedDicts :: DictMap CtEvidence -> TcS () setSolvedDicts solved_dicts = updInertTcS $ \ ics -> ics { inert_solved_dicts = solved_dicts } {- ********************************************************************* * * Other inert-set operations * * ********************************************************************* -} updInertTcS :: (InertSet -> InertSet) -> TcS () -- Modify the inert set with the supplied function updInertTcS upd_fn = do { is_var <- getTcSInertsRef ; wrapTcS (do { curr_inert <- TcM.readTcRef is_var ; TcM.writeTcRef is_var (upd_fn curr_inert) }) } getInertCans :: TcS InertCans getInertCans = do { inerts <- getTcSInerts; return (inert_cans inerts) } setInertCans :: InertCans -> TcS () setInertCans ics = updInertTcS $ \ inerts -> inerts { inert_cans = ics } updRetInertCans :: (InertCans -> (a, InertCans)) -> TcS a -- Modify the inert set with the supplied function updRetInertCans upd_fn = do { is_var <- getTcSInertsRef ; wrapTcS (do { inerts <- TcM.readTcRef is_var ; let (res, cans') = upd_fn (inert_cans inerts) ; TcM.writeTcRef is_var (inerts { inert_cans = cans' }) ; return res }) } updInertCans :: (InertCans -> InertCans) -> TcS () -- Modify the inert set with the supplied function updInertCans upd_fn = updInertTcS $ \ inerts -> inerts { inert_cans = upd_fn (inert_cans inerts) } updInertDicts :: (DictMap Ct -> DictMap Ct) -> TcS () -- Modify the inert set with the supplied function updInertDicts upd_fn = updInertCans $ \ ics -> ics { inert_dicts = upd_fn (inert_dicts ics) } updInertSafehask :: (DictMap Ct -> DictMap Ct) -> TcS () -- Modify the inert set with the supplied function updInertSafehask upd_fn = updInertCans $ \ ics -> ics { inert_safehask = upd_fn (inert_safehask ics) } updInertIrreds :: (Cts -> Cts) -> TcS () -- Modify the inert set with the supplied function updInertIrreds upd_fn = updInertCans $ \ ics -> ics { inert_irreds = upd_fn (inert_irreds ics) } getInertEqs :: TcS (DTyVarEnv EqualCtList) getInertEqs = do { inert <- getInertCans; return (inert_eqs inert) } getInnermostGivenEqLevel :: TcS TcLevel getInnermostGivenEqLevel = do { inert <- getInertCans ; return (inert_given_eq_lvl inert) } getInertInsols :: TcS Cts -- Returns insoluble equality constraints -- specifically including Givens getInertInsols = do { inert <- getInertCans ; return (filterBag insolubleEqCt (inert_irreds inert)) } getInertGivens :: TcS [Ct] -- Returns the Given constraints in the inert set getInertGivens = do { inerts <- getInertCans ; let all_cts = foldDicts (:) (inert_dicts inerts) $ foldFunEqs (\ecl out -> equalCtListToList ecl ++ out) (inert_funeqs inerts) $ concatMap equalCtListToList (dVarEnvElts (inert_eqs inerts)) ; return (filter isGivenCt all_cts) } getPendingGivenScs :: TcS [Ct] -- Find all inert Given dictionaries, or quantified constraints, -- whose cc_pend_sc flag is True -- and that belong to the current level -- Set their cc_pend_sc flag to False in the inert set, and return that Ct getPendingGivenScs = do { lvl <- getTcLevel ; updRetInertCans (get_sc_pending lvl) } get_sc_pending :: TcLevel -> InertCans -> ([Ct], InertCans) get_sc_pending this_lvl ic@(IC { inert_dicts = dicts, inert_insts = insts }) = ASSERT2( all isGivenCt sc_pending, ppr sc_pending ) -- When getPendingScDics is called, -- there are never any Wanteds in the inert set (sc_pending, ic { inert_dicts = dicts', inert_insts = insts' }) where sc_pending = sc_pend_insts ++ sc_pend_dicts sc_pend_dicts = foldDicts get_pending dicts [] dicts' = foldr add dicts sc_pend_dicts (sc_pend_insts, insts') = mapAccumL get_pending_inst [] insts get_pending :: Ct -> [Ct] -> [Ct] -- Get dicts with cc_pend_sc = True -- but flipping the flag get_pending dict dicts | Just dict' <- isPendingScDict dict , belongs_to_this_level (ctEvidence dict) = dict' : dicts | otherwise = dicts add :: Ct -> DictMap Ct -> DictMap Ct add ct@(CDictCan { cc_class = cls, cc_tyargs = tys }) dicts = addDictCt dicts cls tys ct add ct _ = pprPanic "getPendingScDicts" (ppr ct) get_pending_inst :: [Ct] -> QCInst -> ([Ct], QCInst) get_pending_inst cts qci@(QCI { qci_ev = ev }) | Just qci' <- isPendingScInst qci , belongs_to_this_level ev = (CQuantCan qci' : cts, qci') | otherwise = (cts, qci) belongs_to_this_level ev = ctLocLevel (ctEvLoc ev) == this_lvl -- We only want Givens from this level; see (3a) in -- Note [The superclass story] in GHC.Tc.Solver.Canonical getUnsolvedInerts :: TcS ( Bag Implication , Cts ) -- All simple constraints -- Return all the unsolved [Wanted] or [Derived] constraints -- -- Post-condition: the returned simple constraints are all fully zonked -- (because they come from the inert set) -- the unsolved implics may not be getUnsolvedInerts = do { IC { inert_eqs = tv_eqs , inert_funeqs = fun_eqs , inert_irreds = irreds , inert_dicts = idicts } <- getInertCans ; let unsolved_tv_eqs = foldTyEqs add_if_unsolved tv_eqs emptyCts unsolved_fun_eqs = foldFunEqs add_if_unsolveds fun_eqs emptyCts unsolved_irreds = Bag.filterBag is_unsolved irreds unsolved_dicts = foldDicts add_if_unsolved idicts emptyCts unsolved_others = unsolved_irreds `unionBags` unsolved_dicts ; implics <- getWorkListImplics ; traceTcS "getUnsolvedInerts" $ vcat [ text " tv eqs =" <+> ppr unsolved_tv_eqs , text "fun eqs =" <+> ppr unsolved_fun_eqs , text "others =" <+> ppr unsolved_others , text "implics =" <+> ppr implics ] ; return ( implics, unsolved_tv_eqs `unionBags` unsolved_fun_eqs `unionBags` unsolved_others) } where add_if_unsolved :: Ct -> Cts -> Cts add_if_unsolved ct cts | is_unsolved ct = ct `consCts` cts | otherwise = cts add_if_unsolveds :: EqualCtList -> Cts -> Cts add_if_unsolveds new_cts old_cts = foldr add_if_unsolved old_cts (equalCtListToList new_cts) is_unsolved ct = not (isGivenCt ct) -- Wanted or Derived getHasGivenEqs :: TcLevel -- TcLevel of this implication -> TcS ( HasGivenEqs -- are there Given equalities? , Cts ) -- Insoluble equalities arising from givens -- See Note [Tracking Given equalities] getHasGivenEqs tclvl = do { inerts@(IC { inert_irreds = irreds , inert_given_eqs = given_eqs , inert_given_eq_lvl = ge_lvl }) <- getInertCans ; let insols = filterBag insolubleEqCt irreds -- Specifically includes ones that originated in some -- outer context but were refined to an insoluble by -- a local equality; so do /not/ add ct_given_here. -- See Note [HasGivenEqs] in GHC.Tc.Types.Constraint, and -- Note [Tracking Given equalities] in this module has_ge | ge_lvl == tclvl = MaybeGivenEqs | given_eqs = LocalGivenEqs | otherwise = NoGivenEqs ; traceTcS "getHasGivenEqs" $ vcat [ text "given_eqs:" <+> ppr given_eqs , text "ge_lvl:" <+> ppr ge_lvl , text "ambient level:" <+> ppr tclvl , text "Inerts:" <+> ppr inerts , text "Insols:" <+> ppr insols] ; return (has_ge, insols) } mentionsOuterVar :: TcLevel -> CtEvidence -> Bool mentionsOuterVar tclvl ev = anyFreeVarsOfType (isOuterTyVar tclvl) $ ctEvPred ev isOuterTyVar :: TcLevel -> TyCoVar -> Bool -- True of a type variable that comes from a -- shallower level than the ambient level (tclvl) isOuterTyVar tclvl tv | isTyVar tv = ASSERT2( not (isTouchableMetaTyVar tclvl tv), ppr tv <+> ppr tclvl ) tclvl `strictlyDeeperThan` tcTyVarLevel tv -- ASSERT: we are dealing with Givens here, and invariant (GivenInv) from -- Note Note [TcLevel invariants] in GHC.Tc.Utils.TcType ensures that there can't -- be a touchable meta tyvar. If this wasn't true, you might worry that, -- at level 3, a meta-tv alpha[3] gets unified with skolem b[2], and thereby -- becomes "outer" even though its level numbers says it isn't. | otherwise = False -- Coercion variables; doesn't much matter -- | Returns Given constraints that might, -- potentially, match the given pred. This is used when checking to see if a -- Given might overlap with an instance. See Note [Instance and Given overlap] -- in "GHC.Tc.Solver.Interact" matchableGivens :: CtLoc -> PredType -> InertSet -> Cts matchableGivens loc_w pred_w inerts@(IS { inert_cans = inert_cans }) = filterBag matchable_given all_relevant_givens where -- just look in class constraints and irreds. matchableGivens does get called -- for ~R constraints, but we don't need to look through equalities, because -- canonical equalities are used for rewriting. We'll only get caught by -- non-canonical -- that is, irreducible -- equalities. all_relevant_givens :: Cts all_relevant_givens | Just (clas, _) <- getClassPredTys_maybe pred_w = findDictsByClass (inert_dicts inert_cans) clas `unionBags` inert_irreds inert_cans | otherwise = inert_irreds inert_cans matchable_given :: Ct -> Bool matchable_given ct | CtGiven { ctev_loc = loc_g, ctev_pred = pred_g } <- ctEvidence ct = mightEqualLater inerts pred_g loc_g pred_w loc_w | otherwise = False mightEqualLater :: InertSet -> TcPredType -> CtLoc -> TcPredType -> CtLoc -> Bool -- See Note [What might equal later?] -- Used to implement logic in Note [Instance and Given overlap] in GHC.Tc.Solver.Interact mightEqualLater (IS { inert_cycle_breakers = cbvs }) given_pred given_loc wanted_pred wanted_loc | prohibitedSuperClassSolve given_loc wanted_loc = False | otherwise = case tcUnifyTysFG bind_fun [flattened_given] [flattened_wanted] of SurelyApart -> False -- types that are surely apart do not equal later MaybeApart MARInfinite _ -> False -- see Example 7 in the Note. _ -> True where in_scope = mkInScopeSet $ tyCoVarsOfTypes [given_pred, wanted_pred] -- NB: flatten both at the same time, so that we can share mappings -- from type family applications to variables, and also to guarantee -- that the fresh variables are really fresh between the given and -- the wanted. Flattening both at the same time is needed to get -- Example 10 from the Note. ([flattened_given, flattened_wanted], var_mapping) = flattenTysX in_scope [given_pred, wanted_pred] bind_fun :: BindFun bind_fun tv rhs_ty | isMetaTyVar tv , can_unify tv (metaTyVarInfo tv) rhs_ty -- this checks for CycleBreakerTvs and TyVarTvs; forgetting -- the latter was #19106. = BindMe -- See Examples 4, 5, and 6 from the Note | Just (_fam_tc, fam_args) <- lookupVarEnv var_mapping tv , anyFreeVarsOfTypes mentions_meta_ty_var fam_args = BindMe | otherwise = Apart -- True for TauTv and TyVarTv (and RuntimeUnkTv) meta-tyvars -- (as they can be unified) -- and also for CycleBreakerTvs that mentions meta-tyvars mentions_meta_ty_var :: TyVar -> Bool mentions_meta_ty_var tv | isMetaTyVar tv = case metaTyVarInfo tv of -- See Examples 8 and 9 in the Note CycleBreakerTv | Just tyfam_app <- lookup tv cbvs -> anyFreeVarsOfType mentions_meta_ty_var tyfam_app | otherwise -> pprPanic "mightEqualLater finds an unbound cbv" (ppr tv $$ ppr cbvs) _ -> True | otherwise = False -- like canSolveByUnification, but allows cbv variables to unify can_unify :: TcTyVar -> MetaInfo -> Type -> Bool can_unify _lhs_tv TyVarTv rhs_ty -- see Example 3 from the Note | Just rhs_tv <- tcGetTyVar_maybe rhs_ty = case tcTyVarDetails rhs_tv of MetaTv { mtv_info = TyVarTv } -> True MetaTv {} -> False -- could unify with anything SkolemTv {} -> True RuntimeUnk -> True | otherwise -- not a var on the RHS = False can_unify lhs_tv _other _rhs_ty = mentions_meta_ty_var lhs_tv prohibitedSuperClassSolve :: CtLoc -> CtLoc -> Bool -- See Note [Solving superclass constraints] in GHC.Tc.TyCl.Instance prohibitedSuperClassSolve from_loc solve_loc | GivenOrigin (InstSC given_size) <- ctLocOrigin from_loc , ScOrigin wanted_size <- ctLocOrigin solve_loc = given_size >= wanted_size | otherwise = False {- Note [Unsolved Derived equalities] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In getUnsolvedInerts, we return a derived equality from the inert_eqs because it is a candidate for floating out of this implication. We only float equalities with a meta-tyvar on the left, so we only pull those out here. Note [What might equal later?] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We must determine whether a Given might later equal a Wanted. We definitely need to account for the possibility that any metavariable might be arbitrarily instantiated. Yet we do *not* want to allow skolems in to be instantiated, as we've already rewritten with respect to any Givens. (We're solving a Wanted here, and so all Givens have already been processed.) This is best understood by example. 1. C alpha ~? C Int That given certainly might match later. 2. C a ~? C Int No. No new givens are going to arise that will get the `a` to rewrite to Int. 3. C alpha[tv] ~? C Int That alpha[tv] is a TyVarTv, unifiable only with other type variables. It cannot equal later. 4. C (F alpha) ~? C Int Sure -- that can equal later, if we learn something useful about alpha. 5. C (F alpha[tv]) ~? C Int This, too, might equal later. Perhaps we have [G] F b ~ Int elsewhere. Or maybe we have C (F alpha[tv] beta[tv]), these unify with each other, and F x x = Int. Remember: returning True doesn't commit ourselves to anything. 6. C (F a) ~? C a No, this won't match later. If we could rewrite (F a) or a, we would have by now. 7. C (Maybe alpha) ~? C alpha We say this cannot equal later, because it would require alpha := Maybe (Maybe (Maybe ...)). While such a type can be contrived, we choose not to worry about it. See Note [Infinitary substitution in lookup] in GHC.Core.InstEnv. Getting this wrong let to #19107, tested in typecheck/should_compile/T19107. 8. C cbv ~? C Int where cbv = F a The cbv is a cycle-breaker var which stands for F a. See Note [Type variable cycles in Givens] in GHC.Tc.Solver.Canonical. This is just like case 6, and we say "no". Saying "no" here is essential in getting the parser to type-check, with its use of DisambECP. 9. C cbv ~? C Int where cbv = F alpha Here, we might indeed equal later. Distinguishing between this case and Example 8 is why we need the InertSet in mightEqualLater. 10. C (F alpha, Int) ~? C (Bool, F alpha) This cannot equal later, because F a would have to equal both Bool and Int. To deal with type family applications, we use the Core flattener. See Note [Flattening type-family applications when matching instances] in GHC.Core.Unify. The Core flattener replaces all type family applications with fresh variables. The next question: should we allow these fresh variables in the domain of a unifying substitution? A type family application that mentions only skolems (example 6) is settled: any skolems would have been rewritten w.r.t. Givens by now. These type family applications match only themselves. A type family application that mentions metavariables, on the other hand, can match anything. So, if the original type family application contains a metavariable, we use BindMe to tell the unifier to allow it in the substitution. On the other hand, a type family application with only skolems is considered rigid. This treatment fixes #18910 and is tested in typecheck/should_compile/InstanceGivenOverlap{,2} -} removeInertCts :: [Ct] -> InertCans -> InertCans -- ^ Remove inert constraints from the 'InertCans', for use when a -- typechecker plugin wishes to discard a given. removeInertCts cts icans = foldl' removeInertCt icans cts removeInertCt :: InertCans -> Ct -> InertCans removeInertCt is ct = case ct of CDictCan { cc_class = cl, cc_tyargs = tys } -> is { inert_dicts = delDict (inert_dicts is) cl tys } CEqCan { cc_lhs = lhs, cc_rhs = rhs } -> delEq is lhs rhs CQuantCan {} -> panic "removeInertCt: CQuantCan" CIrredCan {} -> panic "removeInertCt: CIrredEvCan" CNonCanonical {} -> panic "removeInertCt: CNonCanonical" -- | Looks up a family application in the inerts; returned coercion -- is oriented input ~ output lookupFamAppInert :: TyCon -> [Type] -> TcS (Maybe (TcCoercion, TcType, CtFlavourRole)) lookupFamAppInert fam_tc tys = do { IS { inert_cans = IC { inert_funeqs = inert_funeqs } } <- getTcSInerts ; return (lookup_inerts inert_funeqs) } where lookup_inerts inert_funeqs | Just (EqualCtList (CEqCan { cc_ev = ctev, cc_rhs = rhs } :| _)) <- findFunEq inert_funeqs fam_tc tys = Just (ctEvCoercion ctev, rhs, ctEvFlavourRole ctev) | otherwise = Nothing lookupInInerts :: CtLoc -> TcPredType -> TcS (Maybe CtEvidence) -- Is this exact predicate type cached in the solved or canonicals of the InertSet? lookupInInerts loc pty | ClassPred cls tys <- classifyPredType pty = do { inerts <- getTcSInerts ; return (lookupSolvedDict inerts loc cls tys `mplus` fmap ctEvidence (lookupInertDict (inert_cans inerts) loc cls tys)) } | otherwise -- NB: No caching for equalities, IPs, holes, or errors = return Nothing -- | Look up a dictionary inert. lookupInertDict :: InertCans -> CtLoc -> Class -> [Type] -> Maybe Ct lookupInertDict (IC { inert_dicts = dicts }) loc cls tys = case findDict dicts loc cls tys of Just ct -> Just ct _ -> Nothing -- | Look up a solved inert. lookupSolvedDict :: InertSet -> CtLoc -> Class -> [Type] -> Maybe CtEvidence -- Returns just if exactly this predicate type exists in the solved. lookupSolvedDict (IS { inert_solved_dicts = solved }) loc cls tys = case findDict solved loc cls tys of Just ev -> Just ev _ -> Nothing --------------------------- lookupFamAppCache :: TyCon -> [Type] -> TcS (Maybe (TcCoercion, TcType)) lookupFamAppCache fam_tc tys = do { IS { inert_famapp_cache = famapp_cache } <- getTcSInerts ; case findFunEq famapp_cache fam_tc tys of result@(Just (co, ty)) -> do { traceTcS "famapp_cache hit" (vcat [ ppr (mkTyConApp fam_tc tys) , ppr ty , ppr co ]) ; return result } Nothing -> return Nothing } extendFamAppCache :: TyCon -> [Type] -> (TcCoercion, TcType) -> TcS () -- NB: co :: rhs ~ F tys, to match expectations of rewriter extendFamAppCache tc xi_args stuff@(_, ty) = do { dflags <- getDynFlags ; when (gopt Opt_FamAppCache dflags) $ do { traceTcS "extendFamAppCache" (vcat [ ppr tc <+> ppr xi_args , ppr ty ]) -- 'co' can be bottom, in the case of derived items ; updInertTcS $ \ is@(IS { inert_famapp_cache = fc }) -> is { inert_famapp_cache = insertFunEq fc tc xi_args stuff } } } -- Remove entries from the cache whose evidence mentions variables in the -- supplied set dropFromFamAppCache :: VarSet -> TcS () dropFromFamAppCache varset = do { inerts@(IS { inert_famapp_cache = famapp_cache }) <- getTcSInerts ; let filtered = filterTcAppMap check famapp_cache ; setTcSInerts $ inerts { inert_famapp_cache = filtered } } where check :: (TcCoercion, TcType) -> Bool check (co, _) = not (anyFreeVarsOfCo (`elemVarSet` varset) co) {- ********************************************************************* * * Irreds * * ********************************************************************* -} foldIrreds :: (Ct -> b -> b) -> Cts -> b -> b foldIrreds k irreds z = foldr k z irreds {- ********************************************************************* * * TcAppMap * * ************************************************************************ Note [Use loose types in inert set] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Whenever we are looking up an inert dictionary (CDictCan) or function equality (CEqCan), we use a TcAppMap, which uses the Unique of the class/type family tycon and then a trie which maps the arguments. This trie does *not* need to match the kinds of the arguments; this Note explains why. Consider the types ty0 = (T ty1 ty2 ty3 ty4) and ty0' = (T ty1' ty2' ty3' ty4'), where ty4 and ty4' have different kinds. Let's further assume that both types ty0 and ty0' are well-typed. Because the kind of T is closed, it must be that one of the ty1..ty3 does not match ty1'..ty3' (and that the kind of the fourth argument to T is dependent on whichever one changed). Since we are matching all arguments, during the inert-set lookup, we know that ty1..ty3 do indeed match ty1'..ty3'. Therefore, the kind of ty4 and ty4' must match, too -- without ever looking at it. Accordingly, we use LooseTypeMap, which skips the kind check when looking up a type. I (Richard E) believe this is just an optimization, and that looking at kinds would be harmless. -} type TcAppMap a = DTyConEnv (ListMap LooseTypeMap a) -- Indexed by tycon then the arg types, using "loose" matching, where -- we don't require kind equality. This allows, for example, (a |> co) -- to match (a). -- See Note [Use loose types in inert set] -- Used for types and classes; hence UniqDFM -- See Note [foldTM determinism] in GHC.Data.TrieMap for why we use DTyConEnv here isEmptyTcAppMap :: TcAppMap a -> Bool isEmptyTcAppMap m = isEmptyDTyConEnv m emptyTcAppMap :: TcAppMap a emptyTcAppMap = emptyDTyConEnv findTcApp :: TcAppMap a -> TyCon -> [Type] -> Maybe a findTcApp m tc tys = do { tys_map <- lookupDTyConEnv m tc ; lookupTM tys tys_map } delTcApp :: TcAppMap a -> TyCon -> [Type] -> TcAppMap a delTcApp m tc tys = adjustDTyConEnv (deleteTM tys) m tc insertTcApp :: TcAppMap a -> TyCon -> [Type] -> a -> TcAppMap a insertTcApp m tc tys ct = alterDTyConEnv alter_tm m tc where alter_tm mb_tm = Just (insertTM tys ct (mb_tm `orElse` emptyTM)) alterTcApp :: forall a. TcAppMap a -> TyCon -> [Type] -> XT a -> TcAppMap a alterTcApp m tc tys upd = alterDTyConEnv alter_tm m tc where alter_tm :: Maybe (ListMap LooseTypeMap a) -> Maybe (ListMap LooseTypeMap a) alter_tm m_elt = Just (alterTM tys upd (m_elt `orElse` emptyTM)) filterTcAppMap :: forall a. (a -> Bool) -> TcAppMap a -> TcAppMap a filterTcAppMap f m = mapMaybeDTyConEnv one_tycon m where one_tycon :: ListMap LooseTypeMap a -> Maybe (ListMap LooseTypeMap a) one_tycon tm | isEmptyTM filtered_tm = Nothing | otherwise = Just filtered_tm where filtered_tm = filterTM f tm tcAppMapToBag :: TcAppMap a -> Bag a tcAppMapToBag m = foldTcAppMap consBag m emptyBag foldTcAppMap :: (a -> b -> b) -> TcAppMap a -> b -> b foldTcAppMap k m z = foldDTyConEnv (foldTM k) z m {- ********************************************************************* * * DictMap * * ********************************************************************* -} {- Note [Tuples hiding implicit parameters] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider f,g :: (?x::Int, C a) => a -> a f v = let ?x = 4 in g v The call to 'g' gives rise to a Wanted constraint (?x::Int, C a). We must /not/ solve this from the Given (?x::Int, C a), because of the intervening binding for (?x::Int). #14218. We deal with this by arranging that we always fail when looking up a tuple constraint that hides an implicit parameter. Not that this applies * both to the inert_dicts (lookupInertDict) * and to the solved_dicts (looukpSolvedDict) An alternative would be not to extend these sets with such tuple constraints, but it seemed more direct to deal with the lookup. Note [Solving CallStack constraints] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Suppose f :: HasCallStack => blah. Then * Each call to 'f' gives rise to [W] s1 :: IP "callStack" CallStack -- CtOrigin = OccurrenceOf f with a CtOrigin that says "OccurrenceOf f". Remember that HasCallStack is just shorthand for IP "callStack CallStack See Note [Overview of implicit CallStacks] in GHC.Tc.Types.Evidence * We cannonicalise such constraints, in GHC.Tc.Solver.Canonical.canClassNC, by pushing the call-site info on the stack, and changing the CtOrigin to record that has been done. Bind: s1 = pushCallStack <site-info> s2 [W] s2 :: IP "callStack" CallStack -- CtOrigin = IPOccOrigin * Then, and only then, we can solve the constraint from an enclosing Given. So we must be careful /not/ to solve 's1' from the Givens. Again, we ensure this by arranging that findDict always misses when looking up souch constraints. -} type DictMap a = TcAppMap a emptyDictMap :: DictMap a emptyDictMap = emptyTcAppMap findDict :: DictMap a -> CtLoc -> Class -> [Type] -> Maybe a findDict m loc cls tys | hasIPSuperClasses cls tys -- See Note [Tuples hiding implicit parameters] = Nothing | Just {} <- isCallStackPred cls tys , OccurrenceOf {} <- ctLocOrigin loc = Nothing -- See Note [Solving CallStack constraints] | otherwise = findTcApp m (classTyCon cls) tys findDictsByClass :: DictMap a -> Class -> Bag a findDictsByClass m cls | Just tm <- lookupDTyConEnv m (classTyCon cls) = foldTM consBag tm emptyBag | otherwise = emptyBag delDict :: DictMap a -> Class -> [Type] -> DictMap a delDict m cls tys = delTcApp m (classTyCon cls) tys addDict :: DictMap a -> Class -> [Type] -> a -> DictMap a addDict m cls tys item = insertTcApp m (classTyCon cls) tys item addDictCt :: DictMap Ct -> Class -> [Type] -> Ct -> DictMap Ct -- Like addDict, but combines [W] and [D] to [WD] -- See Note [KeepBoth] in GHC.Tc.Solver.Interact addDictCt m cls tys new_ct = alterTcApp m (classTyCon cls) tys xt_ct where new_ct_ev = ctEvidence new_ct xt_ct :: Maybe Ct -> Maybe Ct xt_ct (Just old_ct) | CtWanted { ctev_nosh = WOnly } <- old_ct_ev , CtDerived {} <- new_ct_ev = Just (old_ct { cc_ev = old_ct_ev { ctev_nosh = WDeriv }}) | CtDerived {} <- old_ct_ev , CtWanted { ctev_nosh = WOnly } <- new_ct_ev = Just (new_ct { cc_ev = new_ct_ev { ctev_nosh = WDeriv }}) where old_ct_ev = ctEvidence old_ct xt_ct _ = Just new_ct addDictsByClass :: DictMap Ct -> Class -> Bag Ct -> DictMap Ct addDictsByClass m cls items = extendDTyConEnv m (classTyCon cls) (foldr add emptyTM items) where add ct@(CDictCan { cc_tyargs = tys }) tm = insertTM tys ct tm add ct _ = pprPanic "addDictsByClass" (ppr ct) filterDicts :: (Ct -> Bool) -> DictMap Ct -> DictMap Ct filterDicts f m = filterTcAppMap f m partitionDicts :: (Ct -> Bool) -> DictMap Ct -> (Bag Ct, DictMap Ct) partitionDicts f m = foldTcAppMap k m (emptyBag, emptyDicts) where k ct (yeses, noes) | f ct = (ct `consBag` yeses, noes) | otherwise = (yeses, add ct noes) add ct@(CDictCan { cc_class = cls, cc_tyargs = tys }) m = addDict m cls tys ct add ct _ = pprPanic "partitionDicts" (ppr ct) dictsToBag :: DictMap a -> Bag a dictsToBag = tcAppMapToBag foldDicts :: (a -> b -> b) -> DictMap a -> b -> b foldDicts = foldTcAppMap emptyDicts :: DictMap a emptyDicts = emptyTcAppMap {- ********************************************************************* * * FunEqMap * * ********************************************************************* -} type FunEqMap a = TcAppMap a -- A map whose key is a (TyCon, [Type]) pair emptyFunEqs :: TcAppMap a emptyFunEqs = emptyTcAppMap findFunEq :: FunEqMap a -> TyCon -> [Type] -> Maybe a findFunEq m tc tys = findTcApp m tc tys findFunEqsByTyCon :: FunEqMap a -> TyCon -> [a] -- Get inert function equation constraints that have the given tycon -- in their head. Not that the constraints remain in the inert set. -- We use this to check for derived interactions with built-in type-function -- constructors. findFunEqsByTyCon m tc | Just tm <- lookupDTyConEnv m tc = foldTM (:) tm [] | otherwise = [] foldFunEqs :: (a -> b -> b) -> FunEqMap a -> b -> b foldFunEqs = foldTcAppMap insertFunEq :: FunEqMap a -> TyCon -> [Type] -> a -> FunEqMap a insertFunEq m tc tys val = insertTcApp m tc tys val {- ************************************************************************ * * * The TcS solver monad * * * ************************************************************************ Note [The TcS monad] ~~~~~~~~~~~~~~~~~~~~ The TcS monad is a weak form of the main Tc monad All you can do is * fail * allocate new variables * fill in evidence variables Filling in a dictionary evidence variable means to create a binding for it, so TcS carries a mutable location where the binding can be added. This is initialised from the innermost implication constraint. -} data TcSEnv = TcSEnv { tcs_ev_binds :: EvBindsVar, tcs_unified :: IORef Int, -- The number of unification variables we have filled -- The important thing is whether it is non-zero tcs_unif_lvl :: IORef (Maybe TcLevel), -- The Unification Level Flag -- Outermost level at which we have unified a meta tyvar -- Starts at Nothing, then (Just i), then (Just j) where j<i -- See Note [The Unification Level Flag] tcs_count :: IORef Int, -- Global step count tcs_inerts :: IORef InertSet, -- Current inert set -- See Note [WorkList priorities] and tcs_worklist :: IORef WorkList -- Current worklist } --------------- newtype TcS a = TcS { unTcS :: TcSEnv -> TcM a } deriving (Functor) -- | Smart constructor for 'TcS', as describe in Note [The one-shot state -- monad trick] in "GHC.Utils.Monad". mkTcS :: (TcSEnv -> TcM a) -> TcS a mkTcS f = TcS (oneShot f) instance Applicative TcS where pure x = mkTcS $ \_ -> return x (<*>) = ap instance Monad TcS where m >>= k = mkTcS $ \ebs -> do unTcS m ebs >>= (\r -> unTcS (k r) ebs) instance MonadFail TcS where fail err = mkTcS $ \_ -> fail err instance MonadUnique TcS where getUniqueSupplyM = wrapTcS getUniqueSupplyM instance HasModule TcS where getModule = wrapTcS getModule instance MonadThings TcS where lookupThing n = wrapTcS (lookupThing n) -- Basic functionality -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ wrapTcS :: TcM a -> TcS a -- Do not export wrapTcS, because it promotes an arbitrary TcM to TcS, -- and TcS is supposed to have limited functionality wrapTcS action = mkTcS $ \_env -> action -- a TcM action will not use the TcEvBinds wrapErrTcS :: TcM a -> TcS a -- The thing wrapped should just fail -- There's no static check; it's up to the user -- Having a variant for each error message is too painful wrapErrTcS = wrapTcS wrapWarnTcS :: TcM a -> TcS a -- The thing wrapped should just add a warning, or no-op -- There's no static check; it's up to the user wrapWarnTcS = wrapTcS failTcS, panicTcS :: SDoc -> TcS a warnTcS :: WarningFlag -> SDoc -> TcS () addErrTcS :: SDoc -> TcS () failTcS = wrapTcS . TcM.failWith warnTcS flag = wrapTcS . TcM.addWarn (Reason flag) addErrTcS = wrapTcS . TcM.addErr panicTcS doc = pprPanic "GHC.Tc.Solver.Canonical" doc traceTcS :: String -> SDoc -> TcS () traceTcS herald doc = wrapTcS (TcM.traceTc herald doc) {-# INLINE traceTcS #-} -- see Note [INLINE conditional tracing utilities] runTcPluginTcS :: TcPluginM a -> TcS a runTcPluginTcS m = wrapTcS . runTcPluginM m =<< getTcEvBindsVar instance HasDynFlags TcS where getDynFlags = wrapTcS getDynFlags getGlobalRdrEnvTcS :: TcS GlobalRdrEnv getGlobalRdrEnvTcS = wrapTcS TcM.getGlobalRdrEnv bumpStepCountTcS :: TcS () bumpStepCountTcS = mkTcS $ \env -> do { let ref = tcs_count env ; n <- TcM.readTcRef ref ; TcM.writeTcRef ref (n+1) } csTraceTcS :: SDoc -> TcS () csTraceTcS doc = wrapTcS $ csTraceTcM (return doc) {-# INLINE csTraceTcS #-} -- see Note [INLINE conditional tracing utilities] traceFireTcS :: CtEvidence -> SDoc -> TcS () -- Dump a rule-firing trace traceFireTcS ev doc = mkTcS $ \env -> csTraceTcM $ do { n <- TcM.readTcRef (tcs_count env) ; tclvl <- TcM.getTcLevel ; return (hang (text "Step" <+> int n <> brackets (text "l:" <> ppr tclvl <> comma <> text "d:" <> ppr (ctLocDepth (ctEvLoc ev))) <+> doc <> colon) 4 (ppr ev)) } {-# INLINE traceFireTcS #-} -- see Note [INLINE conditional tracing utilities] csTraceTcM :: TcM SDoc -> TcM () -- Constraint-solver tracing, -ddump-cs-trace csTraceTcM mk_doc = do { dflags <- getDynFlags ; when ( dopt Opt_D_dump_cs_trace dflags || dopt Opt_D_dump_tc_trace dflags ) ( do { msg <- mk_doc ; TcM.dumpTcRn False Opt_D_dump_cs_trace "" FormatText msg }) } {-# INLINE csTraceTcM #-} -- see Note [INLINE conditional tracing utilities] runTcS :: TcS a -- What to run -> TcM (a, EvBindMap) runTcS tcs = do { ev_binds_var <- TcM.newTcEvBinds ; res <- runTcSWithEvBinds ev_binds_var tcs ; ev_binds <- TcM.getTcEvBindsMap ev_binds_var ; return (res, ev_binds) } -- | This variant of 'runTcS' will keep solving, even when only Deriveds -- are left around. It also doesn't return any evidence, as callers won't -- need it. runTcSDeriveds :: TcS a -> TcM a runTcSDeriveds tcs = do { ev_binds_var <- TcM.newTcEvBinds ; runTcSWithEvBinds ev_binds_var tcs } -- | This can deal only with equality constraints. runTcSEqualities :: TcS a -> TcM a runTcSEqualities thing_inside = do { ev_binds_var <- TcM.newNoTcEvBinds ; runTcSWithEvBinds ev_binds_var thing_inside } -- | A variant of 'runTcS' that takes and returns an 'InertSet' for -- later resumption of the 'TcS' session. runTcSInerts :: InertSet -> TcS a -> TcM (a, InertSet) runTcSInerts inerts tcs = do ev_binds_var <- TcM.newTcEvBinds runTcSWithEvBinds' False ev_binds_var $ do setTcSInerts inerts a <- tcs new_inerts <- getTcSInerts return (a, new_inerts) runTcSWithEvBinds :: EvBindsVar -> TcS a -> TcM a runTcSWithEvBinds = runTcSWithEvBinds' True runTcSWithEvBinds' :: Bool -- ^ Restore type variable cycles afterwards? -- Don't if you want to reuse the InertSet. -- See also Note [Type variable cycles in Givens] -- in GHC.Tc.Solver.Canonical -> EvBindsVar -> TcS a -> TcM a runTcSWithEvBinds' restore_cycles ev_binds_var tcs = do { unified_var <- TcM.newTcRef 0 ; step_count <- TcM.newTcRef 0 ; inert_var <- TcM.newTcRef emptyInert ; wl_var <- TcM.newTcRef emptyWorkList ; unif_lvl_var <- TcM.newTcRef Nothing ; let env = TcSEnv { tcs_ev_binds = ev_binds_var , tcs_unified = unified_var , tcs_unif_lvl = unif_lvl_var , tcs_count = step_count , tcs_inerts = inert_var , tcs_worklist = wl_var } -- Run the computation ; res <- unTcS tcs env ; count <- TcM.readTcRef step_count ; when (count > 0) $ csTraceTcM $ return (text "Constraint solver steps =" <+> int count) ; when restore_cycles $ do { inert_set <- TcM.readTcRef inert_var ; restoreTyVarCycles inert_set } #if defined(DEBUG) ; ev_binds <- TcM.getTcEvBindsMap ev_binds_var ; checkForCyclicBinds ev_binds #endif ; return res } ---------------------------- #if defined(DEBUG) checkForCyclicBinds :: EvBindMap -> TcM () checkForCyclicBinds ev_binds_map | null cycles = return () | null coercion_cycles = TcM.traceTc "Cycle in evidence binds" $ ppr cycles | otherwise = pprPanic "Cycle in coercion bindings" $ ppr coercion_cycles where ev_binds = evBindMapBinds ev_binds_map cycles :: [[EvBind]] cycles = [c | CyclicSCC c <- stronglyConnCompFromEdgedVerticesUniq edges] coercion_cycles = [c | c <- cycles, any is_co_bind c] is_co_bind (EvBind { eb_lhs = b }) = isEqPrimPred (varType b) edges :: [ Node EvVar EvBind ] edges = [ DigraphNode bind bndr (nonDetEltsUniqSet (evVarsOfTerm rhs)) | bind@(EvBind { eb_lhs = bndr, eb_rhs = rhs}) <- bagToList ev_binds ] -- It's OK to use nonDetEltsUFM here as -- stronglyConnCompFromEdgedVertices is still deterministic even -- if the edges are in nondeterministic order as explained in -- Note [Deterministic SCC] in GHC.Data.Graph.Directed. #endif ---------------------------- setEvBindsTcS :: EvBindsVar -> TcS a -> TcS a setEvBindsTcS ref (TcS thing_inside) = TcS $ \ env -> thing_inside (env { tcs_ev_binds = ref }) nestImplicTcS :: EvBindsVar -> TcLevel -> TcS a -> TcS a nestImplicTcS ref inner_tclvl (TcS thing_inside) = TcS $ \ TcSEnv { tcs_unified = unified_var , tcs_inerts = old_inert_var , tcs_count = count , tcs_unif_lvl = unif_lvl } -> do { inerts <- TcM.readTcRef old_inert_var ; let nest_inert = inerts { inert_cycle_breakers = [] , inert_cans = (inert_cans inerts) { inert_given_eqs = False } } -- All other InertSet fields are inherited ; new_inert_var <- TcM.newTcRef nest_inert ; new_wl_var <- TcM.newTcRef emptyWorkList ; let nest_env = TcSEnv { tcs_count = count -- Inherited , tcs_unif_lvl = unif_lvl -- Inherited , tcs_ev_binds = ref , tcs_unified = unified_var , tcs_inerts = new_inert_var , tcs_worklist = new_wl_var } ; res <- TcM.setTcLevel inner_tclvl $ thing_inside nest_env ; out_inert_set <- TcM.readTcRef new_inert_var ; restoreTyVarCycles out_inert_set #if defined(DEBUG) -- Perform a check that the thing_inside did not cause cycles ; ev_binds <- TcM.getTcEvBindsMap ref ; checkForCyclicBinds ev_binds #endif ; return res } nestTcS :: TcS a -> TcS a -- Use the current untouchables, augmenting the current -- evidence bindings, and solved dictionaries -- But have no effect on the InertCans, or on the inert_famapp_cache -- (we want to inherit the latter from processing the Givens) nestTcS (TcS thing_inside) = TcS $ \ env@(TcSEnv { tcs_inerts = inerts_var }) -> do { inerts <- TcM.readTcRef inerts_var ; new_inert_var <- TcM.newTcRef inerts ; new_wl_var <- TcM.newTcRef emptyWorkList ; let nest_env = env { tcs_inerts = new_inert_var , tcs_worklist = new_wl_var } ; res <- thing_inside nest_env ; new_inerts <- TcM.readTcRef new_inert_var -- we want to propagate the safe haskell failures ; let old_ic = inert_cans inerts new_ic = inert_cans new_inerts nxt_ic = old_ic { inert_safehask = inert_safehask new_ic } ; TcM.writeTcRef inerts_var -- See Note [Propagate the solved dictionaries] (inerts { inert_solved_dicts = inert_solved_dicts new_inerts , inert_cans = nxt_ic }) ; return res } emitImplicationTcS :: TcLevel -> SkolemInfo -> [TcTyVar] -- Skolems -> [EvVar] -- Givens -> Cts -- Wanteds -> TcS TcEvBinds -- Add an implication to the TcS monad work-list emitImplicationTcS new_tclvl skol_info skol_tvs givens wanteds = do { let wc = emptyWC { wc_simple = wanteds } ; imp <- wrapTcS $ do { ev_binds_var <- TcM.newTcEvBinds ; imp <- TcM.newImplication ; return (imp { ic_tclvl = new_tclvl , ic_skols = skol_tvs , ic_given = givens , ic_wanted = wc , ic_binds = ev_binds_var , ic_info = skol_info }) } ; emitImplication imp ; return (TcEvBinds (ic_binds imp)) } emitTvImplicationTcS :: TcLevel -> SkolemInfo -> [TcTyVar] -- Skolems -> Cts -- Wanteds -> TcS () -- Just like emitImplicationTcS but no givens and no bindings emitTvImplicationTcS new_tclvl skol_info skol_tvs wanteds = do { let wc = emptyWC { wc_simple = wanteds } ; imp <- wrapTcS $ do { ev_binds_var <- TcM.newNoTcEvBinds ; imp <- TcM.newImplication ; return (imp { ic_tclvl = new_tclvl , ic_skols = skol_tvs , ic_wanted = wc , ic_binds = ev_binds_var , ic_info = skol_info }) } ; emitImplication imp } {- Note [Propagate the solved dictionaries] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ It's really quite important that nestTcS does not discard the solved dictionaries from the thing_inside. Consider Eq [a] forall b. empty => Eq [a] We solve the simple (Eq [a]), under nestTcS, and then turn our attention to the implications. It's definitely fine to use the solved dictionaries on the inner implications, and it can make a significant performance difference if you do so. -} -- Getters and setters of GHC.Tc.Utils.Env fields -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ -- Getter of inerts and worklist getTcSInertsRef :: TcS (IORef InertSet) getTcSInertsRef = TcS (return . tcs_inerts) getTcSWorkListRef :: TcS (IORef WorkList) getTcSWorkListRef = TcS (return . tcs_worklist) getTcSInerts :: TcS InertSet getTcSInerts = getTcSInertsRef >>= readTcRef setTcSInerts :: InertSet -> TcS () setTcSInerts ics = do { r <- getTcSInertsRef; writeTcRef r ics } getWorkListImplics :: TcS (Bag Implication) getWorkListImplics = do { wl_var <- getTcSWorkListRef ; wl_curr <- readTcRef wl_var ; return (wl_implics wl_curr) } pushLevelNoWorkList :: SDoc -> TcS a -> TcS (TcLevel, a) -- Push the level and run thing_inside -- However, thing_inside should not generate any work items #if defined(DEBUG) pushLevelNoWorkList err_doc (TcS thing_inside) = TcS (\env -> TcM.pushTcLevelM $ thing_inside (env { tcs_worklist = wl_panic }) ) where wl_panic = pprPanic "GHC.Tc.Solver.Monad.buildImplication" err_doc -- This panic checks that the thing-inside -- does not emit any work-list constraints #else pushLevelNoWorkList _ (TcS thing_inside) = TcS (\env -> TcM.pushTcLevelM (thing_inside env)) -- Don't check #endif updWorkListTcS :: (WorkList -> WorkList) -> TcS () updWorkListTcS f = do { wl_var <- getTcSWorkListRef ; updTcRef wl_var f } emitWorkNC :: [CtEvidence] -> TcS () emitWorkNC evs | null evs = return () | otherwise = emitWork (map mkNonCanonical evs) emitWork :: [Ct] -> TcS () emitWork [] = return () -- avoid printing, among other work emitWork cts = do { traceTcS "Emitting fresh work" (vcat (map ppr cts)) ; updWorkListTcS (extendWorkListCts cts) } emitImplication :: Implication -> TcS () emitImplication implic = updWorkListTcS (extendWorkListImplic implic) newTcRef :: a -> TcS (TcRef a) newTcRef x = wrapTcS (TcM.newTcRef x) readTcRef :: TcRef a -> TcS a readTcRef ref = wrapTcS (TcM.readTcRef ref) writeTcRef :: TcRef a -> a -> TcS () writeTcRef ref val = wrapTcS (TcM.writeTcRef ref val) updTcRef :: TcRef a -> (a->a) -> TcS () updTcRef ref upd_fn = wrapTcS (TcM.updTcRef ref upd_fn) getTcEvBindsVar :: TcS EvBindsVar getTcEvBindsVar = TcS (return . tcs_ev_binds) getTcLevel :: TcS TcLevel getTcLevel = wrapTcS TcM.getTcLevel getTcEvTyCoVars :: EvBindsVar -> TcS TyCoVarSet getTcEvTyCoVars ev_binds_var = wrapTcS $ TcM.getTcEvTyCoVars ev_binds_var getTcEvBindsMap :: EvBindsVar -> TcS EvBindMap getTcEvBindsMap ev_binds_var = wrapTcS $ TcM.getTcEvBindsMap ev_binds_var setTcEvBindsMap :: EvBindsVar -> EvBindMap -> TcS () setTcEvBindsMap ev_binds_var binds = wrapTcS $ TcM.setTcEvBindsMap ev_binds_var binds unifyTyVar :: TcTyVar -> TcType -> TcS () -- Unify a meta-tyvar with a type -- We keep track of how many unifications have happened in tcs_unified, -- -- We should never unify the same variable twice! unifyTyVar tv ty = ASSERT2( isMetaTyVar tv, ppr tv ) TcS $ \ env -> do { TcM.traceTc "unifyTyVar" (ppr tv <+> text ":=" <+> ppr ty) ; TcM.writeMetaTyVar tv ty ; TcM.updTcRef (tcs_unified env) (+1) } reportUnifications :: TcS a -> TcS (Int, a) reportUnifications (TcS thing_inside) = TcS $ \ env -> do { inner_unified <- TcM.newTcRef 0 ; res <- thing_inside (env { tcs_unified = inner_unified }) ; n_unifs <- TcM.readTcRef inner_unified ; TcM.updTcRef (tcs_unified env) (+ n_unifs) ; return (n_unifs, res) } getDefaultInfo :: TcS ([Type], (Bool, Bool)) getDefaultInfo = wrapTcS TcM.tcGetDefaultTys -- Just get some environments needed for instance looking up and matching -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ getInstEnvs :: TcS InstEnvs getInstEnvs = wrapTcS $ TcM.tcGetInstEnvs getFamInstEnvs :: TcS (FamInstEnv, FamInstEnv) getFamInstEnvs = wrapTcS $ FamInst.tcGetFamInstEnvs getTopEnv :: TcS HscEnv getTopEnv = wrapTcS $ TcM.getTopEnv getGblEnv :: TcS TcGblEnv getGblEnv = wrapTcS $ TcM.getGblEnv getLclEnv :: TcS TcLclEnv getLclEnv = wrapTcS $ TcM.getLclEnv tcLookupClass :: Name -> TcS Class tcLookupClass c = wrapTcS $ TcM.tcLookupClass c tcLookupId :: Name -> TcS Id tcLookupId n = wrapTcS $ TcM.tcLookupId n -- Setting names as used (used in the deriving of Coercible evidence) -- Too hackish to expose it to TcS? In that case somehow extract the used -- constructors from the result of solveInteract addUsedGREs :: [GlobalRdrElt] -> TcS () addUsedGREs gres = wrapTcS $ TcM.addUsedGREs gres addUsedGRE :: Bool -> GlobalRdrElt -> TcS () addUsedGRE warn_if_deprec gre = wrapTcS $ TcM.addUsedGRE warn_if_deprec gre keepAlive :: Name -> TcS () keepAlive = wrapTcS . TcM.keepAlive -- Various smaller utilities [TODO, maybe will be absorbed in the instance matcher] -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ checkWellStagedDFun :: CtLoc -> InstanceWhat -> PredType -> TcS () -- Check that we do not try to use an instance before it is available. E.g. -- instance Eq T where ... -- f x = $( ... (\(p::T) -> p == p)... ) -- Here we can't use the equality function from the instance in the splice checkWellStagedDFun loc what pred | TopLevInstance { iw_dfun_id = dfun_id } <- what , let bind_lvl = TcM.topIdLvl dfun_id , bind_lvl > impLevel = wrapTcS $ TcM.setCtLocM loc $ do { use_stage <- TcM.getStage ; TcM.checkWellStaged pp_thing bind_lvl (thLevel use_stage) } | otherwise = return () -- Fast path for common case where pp_thing = text "instance for" <+> quotes (ppr pred) pprEq :: TcType -> TcType -> SDoc pprEq ty1 ty2 = pprParendType ty1 <+> char '~' <+> pprParendType ty2 isFilledMetaTyVar_maybe :: TcTyVar -> TcS (Maybe Type) isFilledMetaTyVar_maybe tv = wrapTcS (TcM.isFilledMetaTyVar_maybe tv) isFilledMetaTyVar :: TcTyVar -> TcS Bool isFilledMetaTyVar tv = wrapTcS (TcM.isFilledMetaTyVar tv) zonkTyCoVarsAndFV :: TcTyCoVarSet -> TcS TcTyCoVarSet zonkTyCoVarsAndFV tvs = wrapTcS (TcM.zonkTyCoVarsAndFV tvs) zonkTyCoVarsAndFVList :: [TcTyCoVar] -> TcS [TcTyCoVar] zonkTyCoVarsAndFVList tvs = wrapTcS (TcM.zonkTyCoVarsAndFVList tvs) zonkCo :: Coercion -> TcS Coercion zonkCo = wrapTcS . TcM.zonkCo zonkTcType :: TcType -> TcS TcType zonkTcType ty = wrapTcS (TcM.zonkTcType ty) zonkTcTypes :: [TcType] -> TcS [TcType] zonkTcTypes tys = wrapTcS (TcM.zonkTcTypes tys) zonkTcTyVar :: TcTyVar -> TcS TcType zonkTcTyVar tv = wrapTcS (TcM.zonkTcTyVar tv) zonkSimples :: Cts -> TcS Cts zonkSimples cts = wrapTcS (TcM.zonkSimples cts) zonkWC :: WantedConstraints -> TcS WantedConstraints zonkWC wc = wrapTcS (TcM.zonkWC wc) zonkTyCoVarKind :: TcTyCoVar -> TcS TcTyCoVar zonkTyCoVarKind tv = wrapTcS (TcM.zonkTyCoVarKind tv) ---------------------------- pprKicked :: Int -> SDoc pprKicked 0 = empty pprKicked n = parens (int n <+> text "kicked out") {- ********************************************************************* * * * The Unification Level Flag * * * ********************************************************************* -} {- Note [The Unification Level Flag] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider a deep tree of implication constraints forall[1] a. -- Outer-implic C alpha[1] -- Simple forall[2] c. ....(C alpha[1]).... -- Implic-1 forall[2] b. ....(alpha[1] ~ Int).... -- Implic-2 The (C alpha) is insoluble until we know alpha. We solve alpha by unifying alpha:=Int somewhere deep inside Implic-2. But then we must try to solve the Outer-implic all over again. This time we can solve (C alpha) both in Outer-implic, and nested inside Implic-1. When should we iterate solving a level-n implication? Answer: if any unification of a tyvar at level n takes place in the ic_implics of that implication. * What if a unification takes place at level n-1? Then don't iterate level n, because we'll iterate level n-1, and that will in turn iterate level n. * What if a unification takes place at level n, in the ic_simples of level n? No need to track this, because the kick-out mechanism deals with it. (We can't drop kick-out in favour of iteration, because kick-out works for skolem-equalities, not just unifications.) So the monad-global Unification Level Flag, kept in tcs_unif_lvl keeps track of - Whether any unifications at all have taken place (Nothing => no unifications) - If so, what is the outermost level that has seen a unification (Just lvl) The iteration done in the simplify_loop/maybe_simplify_again loop in GHC.Tc.Solver. It helpful not to iterate unless there is a chance of progress. #8474 is an example: * There's a deeply-nested chain of implication constraints. ?x:alpha => ?y1:beta1 => ... ?yn:betan => [W] ?x:Int * From the innermost one we get a [D] alpha[1] ~ Int, so we can unify. * It's better not to iterate the inner implications, but go all the way out to level 1 before iterating -- because iterating level 1 will iterate the inner levels anyway. (In the olden days when we "floated" thse Derived constraints, this was much, much more important -- we got exponential behaviour, as each iteration produced the same Derived constraint.) -} resetUnificationFlag :: TcS Bool -- We are at ambient level i -- If the unification flag = Just i, reset it to Nothing and return True -- Otherwise leave it unchanged and return False resetUnificationFlag = TcS $ \env -> do { let ref = tcs_unif_lvl env ; ambient_lvl <- TcM.getTcLevel ; mb_lvl <- TcM.readTcRef ref ; TcM.traceTc "resetUnificationFlag" $ vcat [ text "ambient:" <+> ppr ambient_lvl , text "unif_lvl:" <+> ppr mb_lvl ] ; case mb_lvl of Nothing -> return False Just unif_lvl | ambient_lvl `strictlyDeeperThan` unif_lvl -> return False | otherwise -> do { TcM.writeTcRef ref Nothing ; return True } } setUnificationFlag :: TcLevel -> TcS () -- (setUnificationFlag i) sets the unification level to (Just i) -- unless it already is (Just j) where j <= i setUnificationFlag lvl = TcS $ \env -> do { let ref = tcs_unif_lvl env ; mb_lvl <- TcM.readTcRef ref ; case mb_lvl of Just unif_lvl | lvl `deeperThanOrSame` unif_lvl -> return () _ -> TcM.writeTcRef ref (Just lvl) } {- ********************************************************************* * * * Instantiation etc. * * ********************************************************************* -} -- Instantiations -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ instDFunType :: DFunId -> [DFunInstType] -> TcS ([TcType], TcThetaType) instDFunType dfun_id inst_tys = wrapTcS $ TcM.instDFunType dfun_id inst_tys newFlexiTcSTy :: Kind -> TcS TcType newFlexiTcSTy knd = wrapTcS (TcM.newFlexiTyVarTy knd) cloneMetaTyVar :: TcTyVar -> TcS TcTyVar cloneMetaTyVar tv = wrapTcS (TcM.cloneMetaTyVar tv) instFlexi :: [TKVar] -> TcS TCvSubst instFlexi = instFlexiX emptyTCvSubst instFlexiX :: TCvSubst -> [TKVar] -> TcS TCvSubst instFlexiX subst tvs = wrapTcS (foldlM instFlexiHelper subst tvs) instFlexiHelper :: TCvSubst -> TKVar -> TcM TCvSubst instFlexiHelper subst tv = do { uniq <- TcM.newUnique ; details <- TcM.newMetaDetails TauTv ; let name = setNameUnique (tyVarName tv) uniq kind = substTyUnchecked subst (tyVarKind tv) ty' = mkTyVarTy (mkTcTyVar name kind details) ; TcM.traceTc "instFlexi" (ppr ty') ; return (extendTvSubst subst tv ty') } matchGlobalInst :: DynFlags -> Bool -- True <=> caller is the short-cut solver -- See Note [Shortcut solving: overlap] -> Class -> [Type] -> TcS TcM.ClsInstResult matchGlobalInst dflags short_cut cls tys = wrapTcS (TcM.matchGlobalInst dflags short_cut cls tys) tcInstSkolTyVarsX :: TCvSubst -> [TyVar] -> TcS (TCvSubst, [TcTyVar]) tcInstSkolTyVarsX subst tvs = wrapTcS $ TcM.tcInstSkolTyVarsX subst tvs -- Creating and setting evidence variables and CtFlavors -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ data MaybeNew = Fresh CtEvidence | Cached EvExpr isFresh :: MaybeNew -> Bool isFresh (Fresh {}) = True isFresh (Cached {}) = False freshGoals :: [MaybeNew] -> [CtEvidence] freshGoals mns = [ ctev | Fresh ctev <- mns ] getEvExpr :: MaybeNew -> EvExpr getEvExpr (Fresh ctev) = ctEvExpr ctev getEvExpr (Cached evt) = evt setEvBind :: EvBind -> TcS () setEvBind ev_bind = do { evb <- getTcEvBindsVar ; wrapTcS $ TcM.addTcEvBind evb ev_bind } -- | Mark variables as used filling a coercion hole useVars :: CoVarSet -> TcS () useVars co_vars = do { ev_binds_var <- getTcEvBindsVar ; let ref = ebv_tcvs ev_binds_var ; wrapTcS $ do { tcvs <- TcM.readTcRef ref ; let tcvs' = tcvs `unionVarSet` co_vars ; TcM.writeTcRef ref tcvs' } } -- | Equalities only setWantedEq :: TcEvDest -> Coercion -> TcS () setWantedEq (HoleDest hole) co = do { useVars (coVarsOfCo co) ; fillCoercionHole hole co } setWantedEq (EvVarDest ev) _ = pprPanic "setWantedEq" (ppr ev) -- | Good for both equalities and non-equalities setWantedEvTerm :: TcEvDest -> EvTerm -> TcS () setWantedEvTerm (HoleDest hole) tm | Just co <- evTermCoercion_maybe tm = do { useVars (coVarsOfCo co) ; fillCoercionHole hole co } | otherwise = -- See Note [Yukky eq_sel for a HoleDest] do { let co_var = coHoleCoVar hole ; setEvBind (mkWantedEvBind co_var tm) ; fillCoercionHole hole (mkTcCoVarCo co_var) } setWantedEvTerm (EvVarDest ev_id) tm = setEvBind (mkWantedEvBind ev_id tm) {- Note [Yukky eq_sel for a HoleDest] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ How can it be that a Wanted with HoleDest gets evidence that isn't just a coercion? i.e. evTermCoercion_maybe returns Nothing. Consider [G] forall a. blah => a ~ T [W] S ~# T Then doTopReactEqPred carefully looks up the (boxed) constraint (S ~ T) in the quantified constraints, and wraps the (boxed) evidence it gets back in an eq_sel to extract the unboxed (S ~# T). We can't put that term into a coercion, so we add a value binding h = eq_sel (...) and the coercion variable h to fill the coercion hole. We even re-use the CoHole's Id for this binding! Yuk! -} fillCoercionHole :: CoercionHole -> Coercion -> TcS () fillCoercionHole hole co = do { wrapTcS $ TcM.fillCoercionHole hole co ; kickOutAfterFillingCoercionHole hole co } setEvBindIfWanted :: CtEvidence -> EvTerm -> TcS () setEvBindIfWanted ev tm = case ev of CtWanted { ctev_dest = dest } -> setWantedEvTerm dest tm _ -> return () newTcEvBinds :: TcS EvBindsVar newTcEvBinds = wrapTcS TcM.newTcEvBinds newNoTcEvBinds :: TcS EvBindsVar newNoTcEvBinds = wrapTcS TcM.newNoTcEvBinds newEvVar :: TcPredType -> TcS EvVar newEvVar pred = wrapTcS (TcM.newEvVar pred) newGivenEvVar :: CtLoc -> (TcPredType, EvTerm) -> TcS CtEvidence -- Make a new variable of the given PredType, -- immediately bind it to the given term -- and return its CtEvidence -- See Note [Bind new Givens immediately] in GHC.Tc.Types.Constraint newGivenEvVar loc (pred, rhs) = do { new_ev <- newBoundEvVarId pred rhs ; return (CtGiven { ctev_pred = pred, ctev_evar = new_ev, ctev_loc = loc }) } -- | Make a new 'Id' of the given type, bound (in the monad's EvBinds) to the -- given term newBoundEvVarId :: TcPredType -> EvTerm -> TcS EvVar newBoundEvVarId pred rhs = do { new_ev <- newEvVar pred ; setEvBind (mkGivenEvBind new_ev rhs) ; return new_ev } newGivenEvVars :: CtLoc -> [(TcPredType, EvTerm)] -> TcS [CtEvidence] newGivenEvVars loc pts = mapM (newGivenEvVar loc) pts emitNewWantedEq :: CtLoc -> Role -> TcType -> TcType -> TcS Coercion -- | Emit a new Wanted equality into the work-list emitNewWantedEq loc role ty1 ty2 = do { (ev, co) <- newWantedEq loc role ty1 ty2 ; updWorkListTcS (extendWorkListEq (mkNonCanonical ev)) ; return co } -- | Make a new equality CtEvidence newWantedEq :: CtLoc -> Role -> TcType -> TcType -> TcS (CtEvidence, Coercion) newWantedEq = newWantedEq_SI WDeriv newWantedEq_SI :: ShadowInfo -> CtLoc -> Role -> TcType -> TcType -> TcS (CtEvidence, Coercion) newWantedEq_SI si loc role ty1 ty2 = do { hole <- wrapTcS $ TcM.newCoercionHole pty ; traceTcS "Emitting new coercion hole" (ppr hole <+> dcolon <+> ppr pty) ; return ( CtWanted { ctev_pred = pty, ctev_dest = HoleDest hole , ctev_nosh = si , ctev_loc = loc} , mkHoleCo hole ) } where pty = mkPrimEqPredRole role ty1 ty2 -- no equalities here. Use newWantedEq instead newWantedEvVarNC :: CtLoc -> TcPredType -> TcS CtEvidence newWantedEvVarNC = newWantedEvVarNC_SI WDeriv newWantedEvVarNC_SI :: ShadowInfo -> CtLoc -> TcPredType -> TcS CtEvidence -- Don't look up in the solved/inerts; we know it's not there newWantedEvVarNC_SI si loc pty = do { new_ev <- newEvVar pty ; traceTcS "Emitting new wanted" (ppr new_ev <+> dcolon <+> ppr pty $$ pprCtLoc loc) ; return (CtWanted { ctev_pred = pty, ctev_dest = EvVarDest new_ev , ctev_nosh = si , ctev_loc = loc })} newWantedEvVar :: CtLoc -> TcPredType -> TcS MaybeNew newWantedEvVar = newWantedEvVar_SI WDeriv newWantedEvVar_SI :: ShadowInfo -> CtLoc -> TcPredType -> TcS MaybeNew -- For anything except ClassPred, this is the same as newWantedEvVarNC newWantedEvVar_SI si loc pty = do { mb_ct <- lookupInInerts loc pty ; case mb_ct of Just ctev | not (isDerived ctev) -> do { traceTcS "newWantedEvVar/cache hit" $ ppr ctev ; return $ Cached (ctEvExpr ctev) } _ -> do { ctev <- newWantedEvVarNC_SI si loc pty ; return (Fresh ctev) } } newWanted :: CtLoc -> PredType -> TcS MaybeNew -- Deals with both equalities and non equalities. Tries to look -- up non-equalities in the cache newWanted = newWanted_SI WDeriv newWanted_SI :: ShadowInfo -> CtLoc -> PredType -> TcS MaybeNew newWanted_SI si loc pty | Just (role, ty1, ty2) <- getEqPredTys_maybe pty = Fresh . fst <$> newWantedEq_SI si loc role ty1 ty2 | otherwise = newWantedEvVar_SI si loc pty -- deals with both equalities and non equalities. Doesn't do any cache lookups. newWantedNC :: CtLoc -> PredType -> TcS CtEvidence newWantedNC loc pty | Just (role, ty1, ty2) <- getEqPredTys_maybe pty = fst <$> newWantedEq loc role ty1 ty2 | otherwise = newWantedEvVarNC loc pty emitNewDeriveds :: CtLoc -> [TcPredType] -> TcS () emitNewDeriveds loc preds | null preds = return () | otherwise = do { evs <- mapM (newDerivedNC loc) preds ; traceTcS "Emitting new deriveds" (ppr evs) ; updWorkListTcS (extendWorkListDeriveds evs) } emitNewDerivedEq :: CtLoc -> Role -> TcType -> TcType -> TcS () -- Create new equality Derived and put it in the work list -- There's no caching, no lookupInInerts emitNewDerivedEq loc role ty1 ty2 = do { ev <- newDerivedNC loc (mkPrimEqPredRole role ty1 ty2) ; traceTcS "Emitting new derived equality" (ppr ev $$ pprCtLoc loc) ; updWorkListTcS (extendWorkListEq (mkNonCanonical ev)) } -- Very important: put in the wl_eqs -- See Note [Prioritise equalities] (Avoiding fundep iteration) newDerivedNC :: CtLoc -> TcPredType -> TcS CtEvidence newDerivedNC loc pred = return $ CtDerived { ctev_pred = pred, ctev_loc = loc } -- --------- Check done in GHC.Tc.Solver.Interact.selectNewWorkItem???? --------- -- | Checks if the depth of the given location is too much. Fails if -- it's too big, with an appropriate error message. checkReductionDepth :: CtLoc -> TcType -- ^ type being reduced -> TcS () checkReductionDepth loc ty = do { dflags <- getDynFlags ; when (subGoalDepthExceeded dflags (ctLocDepth loc)) $ wrapErrTcS $ solverDepthErrorTcS loc ty } matchFam :: TyCon -> [Type] -> TcS (Maybe (CoercionN, TcType)) -- Given (F tys) return (ty, co), where co :: ty ~N F tys matchFam tycon args = fmap (fmap (first mkTcSymCo)) $ wrapTcS $ matchFamTcM tycon args matchFamTcM :: TyCon -> [Type] -> TcM (Maybe (CoercionN, TcType)) -- Given (F tys) return (ty, co), where co :: F tys ~N ty matchFamTcM tycon args = do { fam_envs <- FamInst.tcGetFamInstEnvs ; let match_fam_result = reduceTyFamApp_maybe fam_envs Nominal tycon args ; TcM.traceTc "matchFamTcM" $ vcat [ text "Matching:" <+> ppr (mkTyConApp tycon args) , ppr_res match_fam_result ] ; return match_fam_result } where ppr_res Nothing = text "Match failed" ppr_res (Just (co,ty)) = hang (text "Match succeeded:") 2 (vcat [ text "Rewrites to:" <+> ppr ty , text "Coercion:" <+> ppr co ]) {- Note [Residual implications] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The wl_implics in the WorkList are the residual implication constraints that are generated while solving or canonicalising the current worklist. Specifically, when canonicalising (forall a. t1 ~ forall a. t2) from which we get the implication (forall a. t1 ~ t2) See GHC.Tc.Solver.Monad.deferTcSForAllEq -} {- ************************************************************************ * * Breaking type variable cycles * * ************************************************************************ -} -- | Replace all type family applications in the RHS with fresh variables, -- emitting givens that relate the type family application to the variable. -- See Note [Type variable cycles in Givens] in GHC.Tc.Solver.Canonical. breakTyVarCycle :: CtLoc -> TcType -- the RHS -> TcS TcType -- new RHS that doesn't have any type families -- This could be considerably more efficient. See Detail (5) of Note. breakTyVarCycle loc = go where go ty | Just ty' <- rewriterView ty = go ty' go (Rep.TyConApp tc tys) | isTypeFamilyTyCon tc = do { let (fun_args, extra_args) = splitAt (tyConArity tc) tys fun_app = mkTyConApp tc fun_args fun_app_kind = tcTypeKind fun_app ; new_tv <- wrapTcS (TcM.newCycleBreakerTyVar fun_app_kind) ; let new_ty = mkTyVarTy new_tv given_pred = mkHeteroPrimEqPred fun_app_kind fun_app_kind fun_app new_ty given_term = evCoercion $ mkNomReflCo new_ty -- See Detail (4) of Note ; new_given <- newGivenEvVar loc (given_pred, given_term) ; traceTcS "breakTyVarCycle replacing type family" (ppr new_given) ; emitWorkNC [new_given] ; updInertTcS $ \is -> is { inert_cycle_breakers = (new_tv, fun_app) : inert_cycle_breakers is } ; extra_args' <- mapM go extra_args ; return (mkAppTys new_ty extra_args') } -- Worried that this substitution will change kinds? -- See Detail (3) of Note | otherwise = mkTyConApp tc <$> mapM go tys go (Rep.AppTy ty1 ty2) = mkAppTy <$> go ty1 <*> go ty2 go (Rep.FunTy vis w arg res) = mkFunTy vis <$> go w <*> go arg <*> go res go (Rep.CastTy ty co) = mkCastTy <$> go ty <*> pure co go ty@(Rep.TyVarTy {}) = return ty go ty@(Rep.LitTy {}) = return ty go ty@(Rep.ForAllTy {}) = return ty -- See Detail (1) of Note go ty@(Rep.CoercionTy {}) = return ty -- See Detail (2) of Note -- | Fill in CycleBreakerTvs with the variables they stand for. -- See Note [Type variable cycles in Givens] in GHC.Tc.Solver.Canonical. restoreTyVarCycles :: InertSet -> TcM () restoreTyVarCycles is = forM_ (inert_cycle_breakers is) $ \ (cycle_breaker_tv, orig_ty) -> TcM.writeMetaTyVar cycle_breaker_tv orig_ty -- Unwrap a type synonym only when either: -- The type synonym is forgetful, or -- the type synonym mentions a type family in its expansion -- See Note [Rewriting synonyms] in GHC.Tc.Solver.Rewrite. rewriterView :: TcType -> Maybe TcType rewriterView ty@(Rep.TyConApp tc _) | isForgetfulSynTyCon tc || (isTypeSynonymTyCon tc && not (isFamFreeTyCon tc)) = tcView ty rewriterView _other = Nothing