ghc-9.14.1: GHC/Tc/Utils/Unify.hs
{-# LANGUAGE DataKinds #-}
{-# LANGUAGE DeriveTraversable #-}
{-# LANGUAGE DerivingStrategies #-}
{-# LANGUAGE LambdaCase #-}
{-# LANGUAGE ScopedTypeVariables #-}
{-# LANGUAGE TupleSections #-}
{-# LANGUAGE RecursiveDo #-}
{-# LANGUAGE MultiWayIf #-}
{-# LANGUAGE RecordWildCards #-}
{-# LANGUAGE TypeApplications #-}
{-
(c) The University of Glasgow 2006
(c) The GRASP/AQUA Project, Glasgow University, 1992-1998
-}
-- | Type subsumption and unification
module GHC.Tc.Utils.Unify (
-- Full-blown subsumption
tcWrapResult, tcWrapResultO, tcWrapResultMono,
tcSubType, tcSubTypeSigma, tcSubTypePat, tcSubTypeDS,
addSubTypeCtxt,
tcSubTypeAmbiguity, tcSubMult,
checkConstraints, checkTvConstraints,
buildImplicationFor, buildTvImplication, emitResidualTvConstraint,
-- Skolemisation
DeepSubsumptionFlag(..), getDeepSubsumptionFlag, isRhoTyDS,
tcSkolemise, tcSkolemiseCompleteSig, tcSkolemiseExpectedType,
dsInstantiate,
-- Various unifications
unifyType, unifyKind, unifyInvisibleType,
unifyExprType, unifyTypeAndEmit, promoteTcType,
swapOverTyVars, touchabilityTest, checkTopShape, lhsPriority,
UnifyEnv(..), updUEnvLoc, setUEnvRole,
uType,
mightEqualLater,
makeTypeConcrete,
--------------------------------
-- Holes
matchExpectedListTy,
matchExpectedTyConApp,
matchExpectedAppTy,
matchExpectedFunTys,
matchExpectedFunKind,
matchActualFunTy, matchActualFunTys,
checkTyEqRhs, recurseIntoTyConApp, recurseIntoFamTyConApp,
PuResult(..), okCheckRefl, mapCheck,
TyEqFlags(..),
notUnifying_TEFTask, unifyingLHSMetaTyVar_TEFTask, defaulting_TEFTask,
pureTyEqFlags_LHSMetaTyVar, mkTEFA_Break,
OccursCheck(..), LevelCheck(..), ConcreteCheck(..),
TyEqFamApp(..), FamAppBreaker(..),
checkPromoteFreeVars,
simpleUnifyCheck, UnifyCheckCaller(..), SimpleUnifyResult(..),
fillInferResult, fillInferResultNoInst
) where
import GHC.Prelude
import GHC.Hs
import GHC.Tc.Errors.Types ( ErrCtxtMsg(..) )
import GHC.Tc.Errors.Ppr ( pprErrCtxtMsg )
import GHC.Tc.Utils.Concrete
import GHC.Tc.Utils.Env
import GHC.Tc.Utils.Instantiate
import GHC.Tc.Utils.Monad
import GHC.Tc.Utils.TcMType
import GHC.Tc.Utils.TcType
import GHC.Tc.Types.Evidence
import GHC.Tc.Types.Constraint
import GHC.Tc.Types.CtLoc( CtLoc, mkKindEqLoc, adjustCtLoc, updateCtLocOrigin, ctLocOrigin )
import GHC.Tc.Types.Origin
import GHC.Tc.Zonk.TcType
import GHC.Tc.Utils.TcMType qualified as TcM
import GHC.Tc.Solver.InertSet
import GHC.Core.Type
import GHC.Core.TyCo.Rep hiding (Refl)
import GHC.Core.TyCo.FVs( isInjectiveInType )
import GHC.Core.TyCo.Ppr( debugPprType {- pprTyVar -} )
import GHC.Core.TyCon
import GHC.Core.Coercion
import GHC.Core.Unify
import GHC.Core.Predicate( EqRel(..), mkEqPredRole, mkNomEqPred )
import GHC.Core.Multiplicity
import GHC.Core.Reduction
import qualified GHC.LanguageExtensions as LangExt
import GHC.Builtin.Types
import GHC.Types.Name
import GHC.Types.Id( idType )
import GHC.Types.Var as Var
import GHC.Types.Var.Set
import GHC.Types.Var.Env
import GHC.Types.Basic
import GHC.Types.Unique.Set (nonDetEltsUniqSet)
import GHC.Utils.Misc
import GHC.Utils.Outputable as Outputable
import GHC.Utils.Panic
import GHC.Driver.DynFlags
import GHC.Data.Bag
import GHC.Data.FastString
import GHC.Data.Maybe (firstJusts)
import Control.Monad
import Data.Functor.Identity (Identity(..))
import qualified Data.List.NonEmpty as NE
import Data.Monoid as DM ( Any(..) )
import qualified Data.Semigroup as S ( (<>) )
import Data.Traversable (for)
{- *********************************************************************
* *
matchActualFunTys
* *
********************************************************************* -}
-- | 'matchActualFunTy' looks for just one function arrow,
-- returning an uninstantiated sigma-type.
--
-- Invariant: the returned argument type has a syntactically fixed
-- RuntimeRep in the sense of Note [Fixed RuntimeRep]
-- in GHC.Tc.Utils.Concrete.
--
-- See Note [Return arguments with a fixed RuntimeRep].
matchActualFunTy
:: ExpectedFunTyOrigin
-- ^ See Note [Herald for matchExpectedFunTys]
-> Maybe TypedThing
-- ^ The thing with type TcSigmaType
-> (Arity, TcType)
-- ^ Total number of value args in the call, and
-- the original function type
-- (Both are used only for error messages)
-> TcRhoType
-- ^ Type to analyse: a TcRhoType
-> TcM (HsWrapper, Scaled TcSigmaTypeFRR, TcSigmaType)
-- This function takes in a type to analyse (a RhoType) and returns
-- an argument type and a result type (splitting apart a function arrow).
-- The returned argument type is a SigmaType with a fixed RuntimeRep;
-- as explained in Note [Return arguments with a fixed RuntimeRep].
--
-- See Note [matchActualFunTy error handling] for the first three arguments
-- If (wrap, arg_ty, res_ty) = matchActualFunTy ... fun_ty
-- then wrap :: fun_ty ~> (arg_ty -> res_ty)
-- and NB: res_ty is an (uninstantiated) SigmaType
matchActualFunTy herald mb_thing err_info fun_ty
= assertPpr (isRhoTy fun_ty) (ppr fun_ty) $
go fun_ty
where
-- Does not allocate unnecessary meta variables: if the input already is
-- a function, we just take it apart. Not only is this efficient,
-- it's important for higher rank: the argument might be of form
-- (forall a. ty) -> other
-- If allocated (fresh-meta-var1 -> fresh-meta-var2) and unified, we'd
-- hide the forall inside a meta-variable
go :: TcRhoType -- The type we're processing, perhaps after
-- expanding type synonyms
-> TcM (HsWrapper, Scaled TcSigmaTypeFRR, TcSigmaType)
go ty | Just ty' <- coreView ty = go ty'
go (FunTy { ft_af = af, ft_mult = w, ft_arg = arg_ty, ft_res = res_ty })
= assert (isVisibleFunArg af) $
do { hasFixedRuntimeRep_syntactic (FRRExpectedFunTy herald 1) arg_ty
; return (idHsWrapper, Scaled w arg_ty, res_ty) }
go ty@(TyVarTy tv)
| isMetaTyVar tv
= do { cts <- readMetaTyVar tv
; case cts of
Indirect ty' -> go ty'
Flexi -> defer ty }
-- In all other cases we bale out into ordinary unification
-- However unlike the meta-tyvar case, we are sure that the
-- number of arguments doesn't match arity of the original
-- type, so we can add a bit more context to the error message
-- (cf #7869).
--
-- It is not always an error, because specialized type may have
-- different arity, for example:
--
-- > f1 = f2 'a'
-- > f2 :: Monad m => m Bool
-- > f2 = undefined
--
-- But in that case we add specialized type into error context
-- anyway, because it may be useful. See also #9605.
go ty = addErrCtxtM (mk_ctxt ty) (defer ty)
------------
defer fun_ty
= do { arg_ty <- new_check_arg_ty herald 1
; res_ty <- newOpenFlexiTyVarTy
; let unif_fun_ty = mkScaledFunTys [arg_ty] res_ty
; co <- unifyType mb_thing fun_ty unif_fun_ty
; return (mkWpCastN co, arg_ty, res_ty) }
------------
mk_ctxt :: TcType -> TidyEnv -> ZonkM (TidyEnv, ErrCtxtMsg)
mk_ctxt _res_ty = mkFunTysMsg herald err_info
{- Note [matchActualFunTy error handling]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
matchActualFunTy is made much more complicated by the
desire to produce good error messages. Consider the application
f @Int x y
In GHC.Tc.Gen.Head.tcInstFun we instantiate the function type, one
argument at a time. It must instantiate any type/dictionary args,
before looking for an arrow type.
But if it doesn't find an arrow type, it wants to generate a message
like "f is applied to two arguments but its type only has one".
To do that, it needs to know about the args that tcArgs has already
munched up -- hence passing in n_val_args_in_call and arg_tys_so_far;
and hence also the accumulating so_far arg to 'go'.
This allows us (in mk_ctxt) to construct f's /instantiated/ type,
with just the values-arg arrows, which is what we really want
in the error message.
Ugh!
-}
-- | Like 'matchExpectedFunTys', but used when you have an "actual" type,
-- for example in function application.
--
-- INVARIANT: the returned argument types all have a syntactically fixed RuntimeRep
-- in the sense of Note [Fixed RuntimeRep] in GHC.Tc.Utils.Concrete.
-- See Note [Return arguments with a fixed RuntimeRep].
matchActualFunTys :: ExpectedFunTyOrigin -- ^ See Note [Herald for matchExpectedFunTys]
-> CtOrigin
-> Arity
-> TcSigmaType
-> TcM (HsWrapper, [Scaled TcSigmaTypeFRR], TcRhoType)
-- If matchActualFunTys n ty = (wrap, [t1,..,tn], res_ty)
-- then wrap : ty ~> (t1 -> ... -> tn -> res_ty)
-- and res_ty is a RhoType
-- NB: the returned type is top-instantiated; it's a RhoType
matchActualFunTys herald ct_orig n_val_args_wanted top_ty
= go n_val_args_wanted [] top_ty
where
go n so_far fun_ty
| not (isRhoTy fun_ty)
= do { (wrap1, rho) <- topInstantiate ct_orig fun_ty
; (wrap2, arg_tys, res_ty) <- go n so_far rho
; return (wrap2 <.> wrap1, arg_tys, res_ty) }
go 0 _ fun_ty = return (idHsWrapper, [], fun_ty)
go n so_far fun_ty
= do { (wrap_fun1, arg_ty1, res_ty1) <- matchActualFunTy
herald Nothing
(n_val_args_wanted, top_ty)
fun_ty
; (wrap_res, arg_tys, res_ty) <- go (n-1) (arg_ty1:so_far) res_ty1
; let wrap_fun2 = mkWpFun idHsWrapper wrap_res arg_ty1 res_ty
-- NB: arg_ty1 comes from matchActualFunTy, so it has
-- a syntactically fixed RuntimeRep as needed to call mkWpFun.
; return (wrap_fun2 <.> wrap_fun1, arg_ty1:arg_tys, res_ty) }
{-
************************************************************************
* *
Skolemisation and matchExpectedFunTys
* *
************************************************************************
Note [Skolemisation overview]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose f :: (forall a. a->a) -> blah, and we have the application (f e)
Then we want to typecheck `e` pushing in the type `forall a. a->a`. But we
need to be careful:
* Roughly speaking, in (tcPolyExpr e (forall a b. rho)), we skolemise `a` and `b`,
and then call (tcExpr e rho)
* But not quite! We must be careful if `e` is a type lambda (\ @p @q -> blah).
Then we want to line up the skolemised variables `a`,`b`
with `p`,`q`, so we can't just call (tcExpr (\ @p @q -> blah) rho)
* A very similar situation arises with
(\ @p @q -> blah) :: forall a b. rho
Again, we must line up `p`, `q` with the skolemised `a` and `b`.
* Another similar situation arises with
g :: forall a b. rho
g @p @q x y = ....
Here again when skolemising `a` and `b` we must be careful to match them up
with `p` and `q`.
OK, so how exactly do we check @p binders in lambdas? First note that we only
we only attempt to deal with @p binders when /checking/. We don't do inference for
(\ @a -> blah), not yet anyway.
For checking, there are two cases to consider:
* Function LHS, where the function has a type signature
f :: forall a. a -> forall b. [b] -> blah
f @p x @q y = ...
* Lambda \ @p x @q y -> ...
\cases { @p x @q y -> ... }
(\case p behaves like \cases { p -> ... }, and p is always a term pattern.)
Both ultimately handled by matchExpectedFunTys.
* Function LHS case is handled by `GHC.Tc.Gen.Bind.tcPolyCheck`:
* It calls `tcSkolemiseCompleteSig`
* Passes the skolemised variables into `tcFunBindMatches`
* Which uses `matchExpectedFunTys` to decompose the function type to
match the arguments
* And then passes the (skolemised-variables ++ arg tys) on to `tcMatches`
* For the Lambda case there are two sub-cases:
* An expression with a type signature: (\ @a x y -> blah) :: hs_ty
This is handled by `GHC.Tc.Gen.Head.tcExprWithSig`, which kind-checks
the signature and hands off to `tcExprPolyCheck` via `tcPolyLExprSig`.
Note that the foralls at the top of hs_ty scope over the expression.
* A higher order call: h e, where h :: poly_ty -> blah
This is handlded by `GHC.Tc.Gen.Expr.tcPolyExpr`, which (in the
checking case) again hands off to `tcExprPolyCheck`. Here there is
no type-variable scoping to worry about.
So both sub-cases end up in `GHC.Tc.Gen.Expr.tcPolyExprCheck`
* This skolemises the /top-level/ invisible binders, but remembers
the binders as [ExpPatType]
* Then it looks for a lambda, and if so, calls `tcLambdaMatches` passing in
the skolemised binders so they can be matched up with the lambda binders.
* Otherwise it does deep-skolemisation if DeepSubsumption is on,
and then calls tcExpr to typecheck `e`
The outer skolemisation in tcPolyExprCheck is done using
* tcSkolemiseCompleteSig when there is a user-written signature
* tcSkolemiseGeneral when the polytype just comes from the context e.g. (f e)
The former just calls the latter, so the two cases differ only slightly:
* Both do shallow skolemisation
* Both go via checkConstraints, which uses implicationNeeded to decide whether
to build an implication constraint even if there /are/ no skolems.
See Note [When to build an implication] below.
The difference between the two cases is that `tcSkolemiseCompleteSig`
also brings the outer type variables into scope. It would do no
harm to do so in both cases, but I found that (to my surprise) doing
so caused a non-trivial (1%-ish) perf hit on the compiler.
* `tcFunBindMatches` and `tcLambdaMatches` both use `matchExpectedFunTys`, which
ensures that any trailing invisible binders are skolemised; and does so deeply
if DeepSubsumption is on.
This corresponds to the plan: "skolemise at the '=' of a function binding or
at the '->' of a lambda binding". (See #17594 and "Plan B2".)
Some wrinkles
(SK1) tcSkolemiseGeneral and tcSkolemiseCompleteSig make fresh type variables
See Note [Instantiate sig with fresh variables]
(SK2) All skolemisation (even without DeepSubsumption) builds just one implication
constraint for a nested forall like:
forall a. Eq a => forall b. Ord b => blah
The implication constraint will look like
forall a b. (Eq a, Ord b) => <constraints>
See the loop in GHC.Tc.Utils.Instantiate.topSkolemise.
and Note [Skolemisation en-bloc] in that module
Some examples:
* f :: forall a b. blah
f @p x = rhs
`tcPolyCheck` calls `tcSkolemiseCompleteSig` to skolemise the signature, and
then calls `tcFunBindMatches` passing in [a_sk, b_sk], the skolemsed
variables. The latter ultimately calls `tcMatches`, and thence `tcMatchPats`.
The latter matches up the `a_sk` with `@p`, and discards the `b_sk`.
* f :: forall (a::Type) (b::a). blah
f @(p::b) x = rhs
`tcSkolemiseCompleteSig` brings `a` and `b` into scope, bound to `a_sk` and `b_sk` resp.
When `tcMatchPats` typechecks the pattern `@(p::b)` it'll find that `b` is in
scope (as a result of tcSkolemiseCompleteSig) which is a bit strange. But
it'll then unify the kinds `Type ~ b`, which will fail as it should.
* f :: Int -> forall (a::Type) (b::a). blah
f x @p = rhs
`matchExpectedFunTys` does shallow skolemisation eagerly, so we'll skolemise the
forall a b. Then `tcMatchPats` will bind [p :-> a_sk], and discard `b_sk`.
Discarding the `b_sk` means that
f x @p = \ @q -> blah
or f x @p = let .. in \ @q -> blah
will both be rejected: this is Plan B2: skolemise at the "=".
* Suppose DeepSubsumption is on
f :: forall a. a -> forall b. b -> b -> forall z. z
f @p x @q y = rhs
The `tcSkolemiseCompleteSig` uses shallow skolemisation, so it only skolemises
and brings into scope [a :-> a_sk]. Then `matchExpectedFunTys` skolemises the
forall b, because it needs to expose two value arguments. Finally
`matchExpectedFunTys` concludes with deeply skolemising the remaining type.
So we end up with `[p :-> a_sk, q :-> b_sk]`. Notice that we must not
deeply-skolemise /first/ or we'd get the tyvars [a_sk, b_sk, c_sk] which would
not line up with the patterns [@p, x, @q, y]
-}
tcSkolemiseGeneral
:: DeepSubsumptionFlag
-> UserTypeCtxt
-> TcType -> TcType -- top_ty and expected_ty
-- Here, top_ty is the type we started to skolemise; used only in SigSkol
-- - expected_ty is the type we are actually skolemising
-- matchExpectedFunTys walks down the type, skolemising as it goes,
-- keeping the same top_ty, but successively smaller expected_tys
-> ([(Name, TcInvisTVBinder)] -> TcType -> TcM result)
-> TcM (HsWrapper, result)
tcSkolemiseGeneral ds_flag ctxt top_ty expected_ty thing_inside
| isRhoTyDS ds_flag expected_ty
-- Fast path for a very very common case: no skolemisation to do
-- But still call checkConstraints in case we need an implication regardless
= do { let sig_skol = SigSkol ctxt top_ty []
; (ev_binds, result) <- checkConstraints sig_skol [] [] $
thing_inside [] expected_ty
; return (mkWpLet ev_binds, result) }
| otherwise
= do { -- rec {..}: see Note [Keeping SkolemInfo inside a SkolemTv]
-- in GHC.Tc.Utils.TcType
; rec { (wrap, tv_prs, given, rho_ty) <- case ds_flag of
Deep -> deeplySkolemise skol_info expected_ty
Shallow -> topSkolemise skol_info expected_ty
; let sig_skol = SigSkol ctxt top_ty (map (fmap binderVar) tv_prs)
; skol_info <- mkSkolemInfo sig_skol }
; let skol_tvs = map (binderVar . snd) tv_prs
; traceTc "tcSkolemiseGeneral" (pprUserTypeCtxt ctxt <+> ppr skol_tvs <+> ppr given)
; (ev_binds, result) <- checkConstraints sig_skol skol_tvs given $
thing_inside tv_prs rho_ty
; return (wrap <.> mkWpLet ev_binds, result) }
-- The ev_binds returned by checkConstraints is very
-- often empty, in which case mkWpLet is a no-op
tcSkolemiseCompleteSig :: TcCompleteSig
-> ([ExpPatType] -> TcRhoType -> TcM result)
-> TcM (HsWrapper, result)
-- ^ The wrapper has type: spec_ty ~> expected_ty
-- See Note [Skolemisation] for the differences between
-- tcSkolemiseCompleteSig and tcTopSkolemise
tcSkolemiseCompleteSig (CSig { sig_bndr = poly_id, sig_ctxt = ctxt, sig_loc = loc })
thing_inside
= do { cur_loc <- getSrcSpanM
; let poly_ty = idType poly_id
; setSrcSpan loc $ -- Sets the location for the implication constraint
tcSkolemiseGeneral Shallow ctxt poly_ty poly_ty $ \tv_prs rho_ty ->
setSrcSpan cur_loc $ -- Revert to the original location
tcExtendNameTyVarEnv (map (fmap binderVar) tv_prs) $
thing_inside (map (mkInvisExpPatType . snd) tv_prs) rho_ty }
tcSkolemiseExpectedType :: TcSigmaType
-> ([ExpPatType] -> TcRhoType -> TcM result)
-> TcM (HsWrapper, result)
-- Just like tcSkolemiseCompleteSig, except that we don't have a user-written
-- type signature, we only have a type comimg from the context.
-- Eg. f :: (forall a. blah) -> blah
-- In the call (f e) we will call tcSkolemiseExpectedType on (forall a.blah)
-- before typececking `e`
tcSkolemiseExpectedType exp_ty thing_inside
= tcSkolemiseGeneral Shallow GenSigCtxt exp_ty exp_ty $ \tv_prs rho_ty ->
thing_inside (map (mkInvisExpPatType . snd) tv_prs) rho_ty
tcSkolemise :: DeepSubsumptionFlag -> UserTypeCtxt -> TcSigmaType
-> (TcRhoType -> TcM result)
-> TcM (HsWrapper, result)
tcSkolemise ds_flag ctxt expected_ty thing_inside
= tcSkolemiseGeneral ds_flag ctxt expected_ty expected_ty $ \_ rho_ty ->
thing_inside rho_ty
checkConstraints :: SkolemInfoAnon
-> [TcTyVar] -- Skolems
-> [EvVar] -- Given
-> TcM result
-> TcM (TcEvBinds, result)
-- checkConstraints is careful to build an implication even if
-- `skol_tvs` and `given` are both empty, under certain circumstances
-- See Note [When to build an implication]
checkConstraints skol_info skol_tvs given thing_inside
= do { implication_needed <- implicationNeeded skol_info skol_tvs given
; if implication_needed
then do { (tclvl, wanted, result) <- pushLevelAndCaptureConstraints thing_inside
; (implics, ev_binds) <- buildImplicationFor tclvl skol_info skol_tvs given wanted
; traceTc "checkConstraints" (ppr tclvl $$ ppr skol_tvs)
; emitImplications implics
; return (ev_binds, result) }
else -- Fast path. We check every function argument with tcCheckPolyExpr,
-- which uses tcTopSkolemise and hence checkConstraints.
-- So this fast path is well-exercised
do { res <- thing_inside
; return (emptyTcEvBinds, res) } }
checkTvConstraints :: SkolemInfo
-> [TcTyVar] -- Skolem tyvars
-> TcM result
-> TcM result
checkTvConstraints skol_info skol_tvs thing_inside
= do { (tclvl, wanted, result) <- pushLevelAndCaptureConstraints thing_inside
; emitResidualTvConstraint skol_info skol_tvs tclvl wanted
; return result }
emitResidualTvConstraint :: SkolemInfo -> [TcTyVar]
-> TcLevel -> WantedConstraints -> TcM ()
emitResidualTvConstraint skol_info skol_tvs tclvl wanted
| not (isEmptyWC wanted) ||
checkTelescopeSkol skol_info_anon
= -- checkTelescopeSkol: in this case, /always/ emit this implication
-- even if 'wanted' is empty. We need the implication so that we check
-- for a bad telescope. See Note [Skolem escape and forall-types] in
-- GHC.Tc.Gen.HsType
do { implic <- buildTvImplication skol_info_anon skol_tvs tclvl wanted
; emitImplication implic }
| otherwise -- Empty 'wanted', emit nothing
= return ()
where
skol_info_anon = getSkolemInfo skol_info
buildTvImplication :: SkolemInfoAnon -> [TcTyVar]
-> TcLevel -> WantedConstraints -> TcM Implication
buildTvImplication skol_info skol_tvs tclvl wanted
= assertPpr (all (isSkolemTyVar <||> isTyVarTyVar) skol_tvs) (ppr skol_tvs) $
do { ev_binds <- newNoTcEvBinds -- Used for equalities only, so all the constraints
-- are solved by filling in coercion holes, not
-- by creating a value-level evidence binding
; implic <- newImplication
; let implic' = implic { ic_tclvl = tclvl
, ic_skols = skol_tvs
, ic_given_eqs = NoGivenEqs
, ic_wanted = wanted
, ic_binds = ev_binds
, ic_info = skol_info }
; checkImplicationInvariants implic'
; return implic' }
implicationNeeded :: SkolemInfoAnon -> [TcTyVar] -> [EvVar] -> TcM Bool
-- See Note [When to build an implication]
implicationNeeded skol_info skol_tvs given
| null skol_tvs
, null given
, not (alwaysBuildImplication skol_info)
= -- Empty skolems and givens
do { tc_lvl <- getTcLevel
; if not (isTopTcLevel tc_lvl) -- No implication needed if we are
then return False -- already inside an implication
else
do { dflags <- getDynFlags -- If any deferral can happen,
-- we must build an implication
; return (gopt Opt_DeferTypeErrors dflags ||
gopt Opt_DeferTypedHoles dflags ||
gopt Opt_DeferOutOfScopeVariables dflags) } }
| otherwise -- Non-empty skolems or givens
= return True -- Definitely need an implication
alwaysBuildImplication :: SkolemInfoAnon -> Bool
-- See Note [When to build an implication]
alwaysBuildImplication _ = False
{- Commmented out for now while I figure out about error messages.
See #14185
Caution: we get some duplication of errors if we build more implications.
Because we get one error for each function RHS, even if it's for
the same class constraint.
alwaysBuildImplication (SigSkol ctxt _ _)
= case ctxt of
FunSigCtxt {} -> True -- RHS of a binding with a signature
_ -> False
alwaysBuildImplication (RuleSkol {}) = True
alwaysBuildImplication (InstSkol {}) = True
alwaysBuildImplication (FamInstSkol {}) = True
alwaysBuildImplication _ = False
-}
buildImplicationFor :: TcLevel -> SkolemInfoAnon -> [TcTyVar]
-> [EvVar] -> WantedConstraints
-> TcM (Bag Implication, TcEvBinds)
buildImplicationFor tclvl skol_info skol_tvs given wanted
| isEmptyWC wanted && null given
-- Optimisation : if there are no wanteds, and no givens
-- don't generate an implication at all.
-- Reason for the (null given): we don't want to lose
-- the "inaccessible alternative" error check
= return (emptyBag, emptyTcEvBinds)
| otherwise
= assertPpr (all (isSkolemTyVar <||> isTyVarTyVar) skol_tvs) (ppr skol_tvs) $
-- Why allow TyVarTvs? Because implicitly declared kind variables in
-- non-CUSK type declarations are TyVarTvs, and we need to bring them
-- into scope as a skolem in an implication. This is OK, though,
-- because TyVarTvs will always remain tyvars, even after unification.
do { ev_binds_var <- newTcEvBinds
; implic <- newImplication
; let implic' = implic { ic_tclvl = tclvl
, ic_skols = skol_tvs
, ic_given = given
, ic_wanted = wanted
, ic_binds = ev_binds_var
, ic_info = skol_info }
; checkImplicationInvariants implic'
; return (unitBag implic', TcEvBinds ev_binds_var) }
{- Note [When to build an implication]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have some 'skolems' and some 'givens', and we are
considering whether to wrap the constraints in their scope into an
implication. We must /always/ do so if either 'skolems' or 'givens' are
non-empty. But what if both are empty? You might think we could
always drop the implication. Other things being equal, the fewer
implications the better. Less clutter and overhead. But we must
take care:
* If we have an unsolved [W] g :: a ~# b, and -fdefer-type-errors,
we'll make a /term-level/ evidence binding for 'g = error "blah"'.
We must have an EvBindsVar those bindings!, otherwise they end up as
top-level unlifted bindings, which are verboten. This only matters
at top level, so we check for that
See also Note [Deferred errors for coercion holes] in GHC.Tc.Errors.
cf #14149 for an example of what goes wrong.
* This is /necessary/ for top level but may be /desirable/ even for
nested bindings, because if the deferred coercion is bound too far
out it will be reported even if that thunk (say) is not evaluated.
* If you have
f :: Int; f = f_blah
g :: Bool; g = g_blah
If we don't build an implication for f or g (no tyvars, no givens),
the constraints for f_blah and g_blah are solved together. And that
can yield /very/ confusing error messages, because we can get
[W] C Int b1 -- from f_blah
[W] C Int b2 -- from g_blan
and fundeps can yield [W] b1 ~ b2, even though the two functions have
literally nothing to do with each other. #14185 is an example.
Building an implication keeps them separate.
Note [Herald for matchExpectedFunTys]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The 'herald' always looks like:
"The equation(s) for 'f' have"
"The abstraction (\x.e) takes"
"The section (+ x) expects"
"The function 'f' is applied to"
This is used to construct a message of form
The abstraction `\Just 1 -> ...' takes two arguments
but its type `Maybe a -> a' has only one
The equation(s) for `f' have two arguments
but its type `Maybe a -> a' has only one
The section `(f 3)' requires 'f' to take two arguments
but its type `Int -> Int' has only one
The function 'f' is applied to two arguments
but its type `Int -> Int' has only one
When visible type applications (e.g., `f @Int 1 2`, as in #13902) enter the
picture, we have a choice in deciding whether to count the type applications as
proper arguments:
The function 'f' is applied to one visible type argument
and two value arguments
but its type `forall a. a -> a` has only one visible type argument
and one value argument
Or whether to include the type applications as part of the herald itself:
The expression 'f @Int' is applied to two arguments
but its type `Int -> Int` has only one
The latter is easier to implement and is arguably easier to understand, so we
choose to implement that option.
Note [matchExpectedFunTys]
~~~~~~~~~~~~~~~~~~~~~~~~~~
matchExpectedFunTys checks that a sigma has the form
of an n-ary function. It passes the decomposed type to the
thing_inside, and returns a wrapper to coerce between the two types
It's used wherever a language construct must have a functional type,
namely:
A lambda expression
A function definition
An operator section
This function must be written CPS'd because it needs to fill in the
ExpTypes produced for arguments before it can fill in the ExpType
passed in.
Note [Return arguments with a fixed RuntimeRep]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The functions
- matchExpectedFunTys,
- matchActualFunTy,
- matchActualFunTys,
peel off argument types, as explained in Note [matchExpectedFunTys].
It's important that these functions return argument types that have
a fixed runtime representation, otherwise we would be in violation
of the representation-polymorphism invariants of
Note [Representation polymorphism invariants] in GHC.Core.
This is why all these functions have an additional invariant,
that the argument types they return all have a syntactically fixed RuntimeRep,
in the sense of Note [Fixed RuntimeRep] in GHC.Tc.Utils.Concrete.
Example:
Suppose we have
type F :: Type -> RuntimeRep
type family F a where { F Int = LiftedRep }
type Dual :: Type -> Type
type family Dual a where
Dual a = a -> ()
f :: forall (a :: TYPE (F Int)). Dual a
f = \ x -> ()
The body of `f` is a lambda abstraction, so we must be able to split off
one argument type from its type. This is handled by `matchExpectedFunTys`
(see 'GHC.Tc.Gen.Match.tcLambdaMatches'). We end up with desugared Core that
looks like this:
f :: forall (a :: TYPE (F Int)). Dual (a |> (TYPE F[0]))
f = \ @(a :: TYPE (F Int)) ->
(\ (x :: (a |> (TYPE F[0]))) -> ())
`cast`
(Sub (Sym (Dual[0] <(a |> (TYPE F[0]))>)))
Two important transformations took place:
1. We inserted casts around the argument type to ensure that it has
a fixed runtime representation, as required by invariant (I1) from
Note [Representation polymorphism invariants] in GHC.Core.
2. We inserted a cast around the whole lambda to make everything line up
with the type signature.
-}
-- | Use this function to split off arguments types when you have an
-- \"expected\" type.
--
-- This function skolemises at each polytype.
--
-- Invariant: this function only applies the provided function
-- to a list of argument types which all have a syntactically fixed RuntimeRep
-- in the sense of Note [Fixed RuntimeRep] in GHC.Tc.Utils.Concrete.
-- See Note [Return arguments with a fixed RuntimeRep].
matchExpectedFunTys :: forall a.
ExpectedFunTyOrigin -- See Note [Herald for matchExpectedFunTys]
-> UserTypeCtxt
-> VisArity
-> ExpSigmaType
-> ([ExpPatType] -> ExpRhoType -> TcM a)
-> TcM (HsWrapper, a)
-- If matchExpectedFunTys n ty = (wrap, _)
-- then wrap : (t1 -> ... -> tn -> ty_r) ~> ty,
-- where [t1, ..., tn], ty_r are passed to the thing_inside
--
-- Unconditionally concludes by skolemising any trailing invisible
-- binders and, if DeepSubsumption is on, it does so deeply.
--
-- Postcondition:
-- If exp_ty is Check {}, then [ExpPatType] and ExpRhoType results are all Check{}
-- If exp_ty is Infer {}, then [ExpPatType] and ExpRhoType results are all Infer{}
matchExpectedFunTys herald _ctxt arity (Infer inf_res) thing_inside
= do { arg_tys <- mapM (new_infer_arg_ty herald) [1 .. arity]
; res_ty <- newInferExpType (ir_inst inf_res)
; result <- thing_inside (map ExpFunPatTy arg_tys) res_ty
; arg_tys <- mapM (\(Scaled m t) -> Scaled m <$> readExpType t) arg_tys
; res_ty <- readExpType res_ty
-- Remarks:
-- 1. use tcMkScaledFunTys rather than mkScaledFunTys, as we might
-- have res_ty :: kappa[tau] for a meta ty-var kappa, in which case
-- mkScaledFunTys would crash. See #26277.
-- 2. tcMkScaledFunTys arg_tys res_ty does not contain any foralls
-- (even nested ones), so no need to instantiate.
; co <- fillInferResultNoInst (tcMkScaledFunTys arg_tys res_ty) inf_res
; return (mkWpCastN co, result) }
matchExpectedFunTys herald ctx arity (Check top_ty) thing_inside
= check arity [] top_ty
where
check :: VisArity -> [ExpPatType] -> TcSigmaType -> TcM (HsWrapper, a)
-- `check` is called only in the Check{} case
-- It collects rev_pat_tys in reversed order
-- n_req is the number of /visible/ arguments still needed
----------------------------
-- Skolemise quantifiers, both visible (up to n_req) and invisible
-- See Note [Visible type application and abstraction] in GHC.Tc.Gen.App
check n_req rev_pat_tys ty
| isSigmaTy ty -- An invisible quantifier at the top
|| (n_req > 0 && isForAllTy ty) -- A visible quantifier at top, and we need it
= do { rec { (n_req', wrap_gen, tv_nms, bndrs, given, inner_ty) <- skolemiseRequired skol_info n_req ty
; let sig_skol = SigSkol ctx top_ty (tv_nms `zip` skol_tvs)
skol_tvs = binderVars bndrs
; skol_info <- mkSkolemInfo sig_skol }
-- rec {..}: see Note [Keeping SkolemInfo inside a SkolemTv]
-- in GHC.Tc.Utils.TcType
; (ev_binds, (wrap_res, result))
<- checkConstraints (getSkolemInfo skol_info) skol_tvs given $
check n_req'
(reverse (map ExpForAllPatTy bndrs) ++ rev_pat_tys)
inner_ty
; assertPpr (not (null bndrs && null given)) (ppr ty) $
-- The guard ensures that we made some progress
return (wrap_gen <.> mkWpLet ev_binds <.> wrap_res, result) }
----------------------------
-- Base case: (n_req == 0): no more args
-- The earlier skolemisation ensurs that rho_ty has no top-level invisible quantifiers
-- If there is deep subsumption, do deep skolemisation now
check n_req rev_pat_tys rho_ty
| n_req == 0
= do { let pat_tys = reverse rev_pat_tys
; ds_flag <- getDeepSubsumptionFlag
; case ds_flag of
Shallow -> do { res <- thing_inside pat_tys (mkCheckExpType rho_ty)
; return (idHsWrapper, res) }
Deep -> tcSkolemiseGeneral Deep ctx top_ty rho_ty $ \_ rho_ty ->
-- "_" drop the /deeply/-skolemise binders
-- They do not line up with binders in the Match
thing_inside pat_tys (mkCheckExpType rho_ty) }
----------------------------
-- Function types
check n_req rev_pat_tys (FunTy { ft_af = af, ft_mult = mult
, ft_arg = arg_ty, ft_res = res_ty })
= assert (isVisibleFunArg af) $
do { let arg_pos = arity - n_req + 1 -- 1 for the first argument etc
; (arg_co, arg_ty) <- hasFixedRuntimeRep (FRRExpectedFunTy herald arg_pos) arg_ty
; (wrap_res, result) <- check (n_req - 1)
(mkCheckExpFunPatTy (Scaled mult arg_ty) : rev_pat_tys)
res_ty
; let wrap_arg = mkWpCastN arg_co
fun_wrap = mkWpFun wrap_arg wrap_res (Scaled mult arg_ty) res_ty
; return (fun_wrap, result) }
----------------------------
-- Type variables
check n_req rev_pat_tys ty@(TyVarTy tv)
| isMetaTyVar tv
= do { cts <- readMetaTyVar tv
; case cts of
Indirect ty' -> check n_req rev_pat_tys ty'
Flexi -> defer n_req rev_pat_tys ty }
----------------------------
-- NOW do coreView. We didn't do it before, so that we do not unnecessarily
-- unwrap a synonym in the returned rho_ty
check n_req rev_pat_tys ty
| Just ty' <- coreView ty = check n_req rev_pat_tys ty'
-- In all other cases we bale out into ordinary unification
-- However unlike the meta-tyvar case, we are sure that the
-- number of arguments doesn't match arity of the original
-- type, so we can add a bit more context to the error message
-- (cf #7869).
--
-- It is not always an error, because specialized type may have
-- different arity, for example:
--
-- > f1 = f2 'a'
-- > f2 :: Monad m => m Bool
-- > f2 = undefined
--
-- But in that case we add specialized type into error context
-- anyway, because it may be useful. See also #9605.
check n_req rev_pat_tys res_ty
= addErrCtxtM (mkFunTysMsg herald (arity, top_ty)) $
defer n_req rev_pat_tys res_ty
------------
defer :: VisArity -> [ExpPatType] -> TcRhoType -> TcM (HsWrapper, a)
defer n_req rev_pat_tys fun_ty
= do { more_arg_tys <- mapM (new_check_arg_ty herald) [arity - n_req + 1 .. arity]
; let all_pats = reverse rev_pat_tys ++ map mkCheckExpFunPatTy more_arg_tys
; res_ty <- newOpenFlexiTyVarTy
; result <- thing_inside all_pats (mkCheckExpType res_ty)
; co <- unifyType Nothing (mkScaledFunTys more_arg_tys res_ty) fun_ty
; return (mkWpCastN co, result) }
new_infer_arg_ty :: ExpectedFunTyOrigin -> Int -> TcM (Scaled ExpRhoTypeFRR)
new_infer_arg_ty herald arg_pos -- position for error messages only
= do { mult <- newFlexiTyVarTy multiplicityTy
; inf_hole <- newInferExpTypeFRR IIF_DeepRho (FRRExpectedFunTy herald arg_pos)
; return (mkScaled mult inf_hole) }
new_check_arg_ty :: ExpectedFunTyOrigin -> Int -> TcM (Scaled TcType)
new_check_arg_ty herald arg_pos -- Position for error messages only, 1 for first arg
= do { mult <- newFlexiTyVarTy multiplicityTy
; arg_ty <- newOpenFlexiFRRTyVarTy (FRRExpectedFunTy herald arg_pos)
; return (mkScaled mult arg_ty) }
mkFunTysMsg :: ExpectedFunTyOrigin
-> (VisArity, TcType)
-> TidyEnv -> ZonkM (TidyEnv, ErrCtxtMsg)
-- See Note [Reporting application arity errors]
mkFunTysMsg herald (n_vis_args_in_call, fun_ty) env
= do { (env', fun_ty) <- zonkTidyTcType env fun_ty
; let (pi_ty_bndrs, _) = splitPiTys fun_ty
n_fun_args = count isVisiblePiTyBinder pi_ty_bndrs
; return (env', FunTysCtxt herald fun_ty n_vis_args_in_call n_fun_args) }
{- Note [Reporting application arity errors]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider f :: Int -> Int -> Int
and the call foo = f 3 4 5
We'd like to get an error like:
• Couldn't match expected type ‘t0 -> t’ with actual type ‘Int’
• The function ‘f’ is applied to three visible arguments, -- What are "visible" arguments?
but its type ‘Int -> Int -> Int’ has only two -- See Note [Visibility and arity] in GHC.Types.Basic
That is what `mkFunTysMsg` tries to do. But what is the "type of the function".
Most obviously, we can report its full, polymorphic type; that is simple and
explicable. But sometimes a bit odd. Consider
f :: Bool -> t Int Int
foo = f True 5 10
We get this error:
• Couldn't match type ‘Int’ with ‘t0 -> t’
Expected: Int -> t0 -> t
Actual: Int -> Int
• The function ‘f’ is applied to three visible arguments,
but its type ‘Bool -> t Int Int’ has only one
That's not /quite/ right beause we can instantiate `t` to an arrow and get
two arrows (but not three!). With that in mind, one could consider reporting
the /instantiated/ type, and GHC used to do so. But it's more work, and in
some ways more confusing, especially when nested quantifiers are concerned, e.g.
f :: Bool -> forall t. t Int Int
So we just keep it simple and report the original function type.
************************************************************************
* *
Other matchExpected functions
* *
********************************************************************* -}
matchExpectedListTy :: TcRhoType -> TcM (TcCoercionN, TcRhoType)
-- Special case for lists
matchExpectedListTy exp_ty
= do { (co, [elt_ty]) <- matchExpectedTyConApp listTyCon exp_ty
; return (co, elt_ty) }
---------------------
matchExpectedTyConApp :: TyCon -- T :: forall kv1 ... kvm. k1 -> ... -> kn -> *
-> TcRhoType -- orig_ty
-> TcM (TcCoercionN, -- T k1 k2 k3 a b c ~N orig_ty
[TcSigmaType]) -- Element types, k1 k2 k3 a b c
-- It's used for wired-in tycons, so we call checkWiredInTyCon
-- Precondition: never called with FunTyCon
-- Precondition: input type :: *
-- Postcondition: (T k1 k2 k3 a b c) is well-kinded
matchExpectedTyConApp tc orig_ty
= assertPpr (isAlgTyCon tc) (ppr tc) $
go orig_ty
where
go ty
| Just ty' <- coreView ty
= go ty'
go ty@(TyConApp tycon args)
| tc == tycon -- Common case
= return (mkNomReflCo ty, args)
go (TyVarTy tv)
| isMetaTyVar tv
= do { cts <- readMetaTyVar tv
; case cts of
Indirect ty -> go ty
Flexi -> defer }
go _ = defer
-- If the common case does not occur, instantiate a template
-- T k1 .. kn t1 .. tm, and unify with the original type
-- Doing it this way ensures that the types we return are
-- kind-compatible with T. For example, suppose we have
-- matchExpectedTyConApp T (f Maybe)
-- where data T a = MkT a
-- Then we don't want to instantiate T's data constructors with
-- (a::*) ~ Maybe
-- because that'll make types that are utterly ill-kinded.
-- This happened in #7368
defer
= do { (_, arg_tvs) <- newMetaTyVars (tyConTyVars tc)
; traceTc "matchExpectedTyConApp" (ppr tc $$ ppr (tyConTyVars tc) $$ ppr arg_tvs)
; let args = mkTyVarTys arg_tvs
tc_template = mkTyConApp tc args
; co <- unifyType Nothing tc_template orig_ty
; return (co, args) }
----------------------
matchExpectedAppTy :: TcRhoType -- orig_ty
-> TcM (TcCoercion, -- m a ~N orig_ty
(TcSigmaType, TcSigmaType)) -- Returns m, a
-- If the incoming type is a mutable type variable of kind k, then
-- matchExpectedAppTy returns a new type variable (m: * -> k); note the *.
matchExpectedAppTy orig_ty
= go orig_ty
where
go ty
| Just ty' <- coreView ty = go ty'
| Just (fun_ty, arg_ty) <- tcSplitAppTy_maybe ty
= return (mkNomReflCo orig_ty, (fun_ty, arg_ty))
go (TyVarTy tv)
| isMetaTyVar tv
= do { cts <- readMetaTyVar tv
; case cts of
Indirect ty -> go ty
Flexi -> defer }
go _ = defer
-- Defer splitting by generating an equality constraint
defer
= do { ty1 <- newFlexiTyVarTy kind1
; ty2 <- newFlexiTyVarTy kind2
; co <- unifyType Nothing (mkAppTy ty1 ty2) orig_ty
; return (co, (ty1, ty2)) }
orig_kind = typeKind orig_ty
kind1 = mkVisFunTyMany liftedTypeKind orig_kind
kind2 = liftedTypeKind -- m :: * -> k
-- arg type :: *
{- **********************************************************************
*
fillInferResult
*
********************************************************************** -}
-- | Fill an 'InferResult' with the given type.
--
-- If @co = fillInferResult t1 infer_res@, then @co :: t1 ~# t2@,
-- where @t2@ is the type stored in the 'ir_ref' field of @infer_res@.
--
-- This function enforces the following invariants:
--
-- - Level invariant.
-- The stored type @t2@ is at the same level as given by the
-- 'ir_lvl' field.
-- - FRR invariant.
-- Whenever the 'ir_frr' field is `IFRR_Check`, @t2@ is guaranteed
-- to have a syntactically fixed RuntimeRep, in the sense of
-- Note [Fixed RuntimeRep] in GHC.Tc.Utils.Concrete.
fillInferResultNoInst :: TcType -> InferResult -> TcM TcCoercionN
fillInferResultNoInst act_res_ty (IR { ir_uniq = u
, ir_lvl = res_lvl
, ir_frr = mb_frr
, ir_ref = ref })
= do { mb_exp_res_ty <- readTcRef ref
; case mb_exp_res_ty of
Just exp_res_ty
-- We progressively refine the type stored in 'ref',
-- for example when inferring types across multiple equations.
--
-- Example:
--
-- \ x -> case y of { True -> x ; False -> 3 :: Int }
--
-- When inferring the return type of this function, we will create
-- an 'Infer' 'ExpType', which will first be filled by the type of 'x'
-- after typechecking the first equation, and then filled again with
-- the type 'Int', at which point we want to ensure that we unify
-- the type of 'x' with 'Int'. This is what is happening below when
-- we are "joining" several inferred 'ExpType's.
-> do { traceTc "Joining inferred ExpType" $
ppr u <> colon <+> ppr act_res_ty <+> char '~' <+> ppr exp_res_ty
; cur_lvl <- getTcLevel
; unless (cur_lvl `sameDepthAs` res_lvl) $
ensureMonoType act_res_ty -- See (FIR1)
; unifyType Nothing act_res_ty exp_res_ty }
Nothing
-> do { traceTc "Filling inferred ExpType" $
ppr u <+> text ":=" <+> ppr act_res_ty
-- Enforce the level invariant: ensure the TcLevel of
-- the type we are writing to 'ref' matches 'ir_lvl'.
; (prom_co, act_res_ty) <- promoteTcType res_lvl act_res_ty
-- Enforce the FRR invariant: ensure the type has a syntactically
-- fixed RuntimeRep (if necessary, i.e. 'mb_frr' is not 'Nothing').
; (frr_co, act_res_ty) <-
case mb_frr of
IFRR_Any -> return (mkNomReflCo act_res_ty, act_res_ty)
IFRR_Check frr_orig -> hasFixedRuntimeRep frr_orig act_res_ty
-- Compose the two coercions.
; let final_co = prom_co `mkTransCo` frr_co
; writeTcRef ref (Just act_res_ty)
; return final_co } }
fillInferResult :: CtOrigin -> TcType -> InferResult -> TcM HsWrapper
-- See Note [Instantiation of InferResult]
fillInferResult ct_orig res_ty ires@(IR { ir_inst = iif })
= case iif of
IIF_Sigma -> do { co <- fillInferResultNoInst res_ty ires
; return (mkWpCastN co) }
IIF_ShallowRho -> do { (wrap, res_ty') <- topInstantiate ct_orig res_ty
; co <- fillInferResultNoInst res_ty' ires
; return (mkWpCastN co <.> wrap) }
IIF_DeepRho -> do { (wrap, res_ty') <- dsInstantiate ct_orig res_ty
; co <- fillInferResultNoInst res_ty' ires
; return (mkWpCastN co <.> wrap) }
{- Note [fillInferResult]
~~~~~~~~~~~~~~~~~~~~~~~~~
When inferring, we use fillInferResult to "fill in" the hole in InferResult
data InferResult = IR { ir_uniq :: Unique
, ir_lvl :: TcLevel
, ir_ref :: IORef (Maybe TcType) }
There are two things to worry about:
1. What if it is under a GADT or existential pattern match?
- GADTs: a unification variable (and Infer's hole is similar) is untouchable
- Existentials: be careful about skolem-escape
2. What if it is filled in more than once? E.g. multiple branches of a case
case e of
T1 -> e1
T2 -> e2
Our typing rules are:
* The RHS of a existential or GADT alternative must always be a
monotype, regardless of the number of alternatives.
* Multiple non-existential/GADT branches can have (the same)
higher rank type (#18412). E.g. this is OK:
case e of
True -> hr
False -> hr
where hr:: (forall a. a->a) -> Int
c.f. Section 7.1 of "Practical type inference for arbitrary-rank types"
We use choice (2) in that Section.
(GHC 8.10 and earlier used choice (1).)
But note that
case e of
True -> hr
False -> \x -> hr x
will fail, because we still /infer/ both branches, so the \x will get
a (monotype) unification variable, which will fail to unify with
(forall a. a->a)
For (1) we can detect the GADT/existential situation by seeing that
the current TcLevel is greater than that stored in ir_lvl of the Infer
ExpType. We bump the level whenever we go past a GADT/existential match.
Then, before filling the hole use promoteTcType to promote the type
to the outer ir_lvl. promoteTcType does this
- create a fresh unification variable alpha at level ir_lvl
- emits an equality alpha[ir_lvl] ~ ty
- fills the hole with alpha
That forces the type to be a monotype (since unification variables can
only unify with monotypes); and catches skolem-escapes because the
alpha is untouchable until the equality floats out.
For (2), we simply look to see if the hole is filled already.
- if not, we promote (as above) and fill the hole
- if it is filled, we simply unify with the type that is
already there
(FIR1) There is one wrinkle. Suppose we have
case e of
T1 -> e1 :: (forall a. a->a) -> Int
G2 -> e2
where T1 is not GADT or existential, but G2 is a GADT. Then suppose the
T1 alternative fills the hole with (forall a. a->a) -> Int, which is fine.
But now the G2 alternative must not *just* unify with that else we'd risk
allowing through (e2 :: (forall a. a->a) -> Int). If we'd checked G2 first
we'd have filled the hole with a unification variable, which enforces a
monotype.
So if we check G2 second, we still want to emit a constraint that restricts
the RHS to be a monotype. This is done by ensureMonoType, and it works
by simply generating a constraint (alpha ~ ty), where alpha is a fresh
unification variable. We discard the evidence.
Note [Instantiation of InferResult]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When typechecking expressions (not types, not patterns), we always almost
always instantiate before filling in `InferResult`, so that the result is a
TcRhoType. This behaviour is controlled by the `ir_inst :: InferInstFlag`
field of `InferResult`.
If we do instantiate (ir_inst = IIF_DeepRho), and DeepSubsumption is enabled,
we instantiate deeply. See `tcInferResult`.
Usually this field is `IIF_DeepRho` meaning "return a (possibly deep) rho-type".
Why is this the common case? See #17173 for discussion. Here are some examples
of why:
1. Consider
f x = (*)
We want to instantiate the type of (*) before returning, else we
will infer the type
f :: forall {a}. a -> forall b. Num b => b -> b -> b
This is surely confusing for users.
And worse, the monomorphism restriction won't work properly. The MR is
dealt with in simplifyInfer, and simplifyInfer has no way of
instantiating. This could perhaps be worked around, but it may be
hard to know even when instantiation should happen.
2. Another reason. Consider
f :: (?x :: Int) => a -> a
g y = let ?x = 3::Int in f
Here want to instantiate f's type so that the ?x::Int constraint
gets discharged by the enclosing implicit-parameter binding.
3. Suppose one defines plus = (+). If we instantiate lazily, we will
infer plus :: forall a. Num a => a -> a -> a. However, the monomorphism
restriction compels us to infer
plus :: Integer -> Integer -> Integer
(or similar monotype). Indeed, the only way to know whether to apply
the monomorphism restriction at all is to instantiate
HOWEVER, not always! Here are places where we want `IIF_Sigma` meaning
"return a sigma-type":
* IIF_Sigma: In GHC.Tc.Module.tcRnExpr, which implements GHCi's :type
command, we want to return a completely uninstantiated type.
See Note [Implementing :type] in GHC.Tc.Module.
* IIF_Sigma: In types we can't lambda-abstract, so we must be careful not to instantiate
at all. See calls to `runInferHsType`
* IIF_Sigma: in patterns we don't want to instantiate at all. See the use of
`runInferSigmaFRR` in GHC.Tc.Gen.Pat
* IIF_ShallowRho: in the expression part of a view pattern, we must top-instantiate
but /not/ deeply instantiate (#26331). See Note [View patterns and polymorphism]
in GHC.Tc.Gen.Pat. This the only place we use IIF_ShallowRho.
Why do we want to deeply instantiate, ever? Why isn't top-instantiation enough?
Answer: to accept the following program (T26225b) with -XDeepSubsumption, we
need to deeply instantiate when inferring in checkResultTy:
f :: Int -> (forall a. a->a)
g :: Int -> Bool -> Bool
test b =
case b of
True -> f
False -> g
If we don't deeply instantiate in the branches of the case expression, we will
try to unify the type of 'f' with that of 'g', which fails. If we instead
deeply instantiate 'f', we will fill the 'InferResult' with 'Int -> alpha -> alpha'
which then successfully unifies with the type of 'g' when we come to fill the
'InferResult' hole a second time for the second case branch.
-}
{-
************************************************************************
* *
Subsumption checking
* *
************************************************************************
Note [Subsumption checking: tcSubType]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
All the tcSubType calls have the form
tcSubType actual_ty expected_ty
which checks
actual_ty <= expected_ty
That is, that a value of type actual_ty is acceptable in
a place expecting a value of type expected_ty. I.e. that
actual ty is more polymorphic than expected_ty
It returns a wrapper function
co_fn :: actual_ty ~ expected_ty
which takes an HsExpr of type actual_ty into one of type
expected_ty.
Note [Ambiguity check and deep subsumption]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider
f :: (forall b. Eq b => a -> a) -> Int
Does `f` have an ambiguous type? The ambiguity check usually checks
that this definition of f' would typecheck, where f' has the exact same
type as f:
f' :: (forall b. Eq b => a -> a) -> Intp
f' = f
This will be /rejected/ with DeepSubsumption but /accepted/ with
ShallowSubsumption. On the other hand, this eta-expanded version f''
would be rejected both ways:
f'' :: (forall b. Eq b => a -> a) -> Intp
f'' x = f x
This is squishy in the same way as other examples in GHC.Tc.Validity
Note [The squishiness of the ambiguity check]
The situation in June 2022. Since we have SimpleSubsumption at the moment,
we don't want introduce new breakage if you add -XDeepSubsumption, by
rejecting types as ambiguous that weren't ambiguous before. So, as a
holding decision, we /always/ use SimpleSubsumption for the ambiguity check
(erring on the side accepting more programs). Hence tcSubTypeAmbiguity.
-}
-----------------
-- tcWrapResult needs both un-type-checked (for origins and error messages)
-- and type-checked (for wrapping) expressions
tcWrapResult :: HsExpr GhcRn -> HsExpr GhcTc -> TcSigmaType -> ExpRhoType
-> TcM (HsExpr GhcTc)
tcWrapResult rn_expr = tcWrapResultO (exprCtOrigin rn_expr) rn_expr
tcWrapResultO :: CtOrigin -> HsExpr GhcRn -> HsExpr GhcTc -> TcSigmaType -> ExpRhoType
-> TcM (HsExpr GhcTc)
tcWrapResultO orig rn_expr expr actual_ty res_ty
= do { traceTc "tcWrapResult" (vcat [ text "Actual: " <+> ppr actual_ty
, text "Expected:" <+> ppr res_ty ])
; wrap <- tcSubType orig GenSigCtxt (Just $ HsExprRnThing rn_expr) actual_ty res_ty
; return (mkHsWrap wrap expr) }
-- | A version of 'tcWrapResult' to use when the actual type is a
-- rho-type, so nothing to instantiate; just go straight to unify.
-- It means we don't need to pass in a CtOrigin.
tcWrapResultMono :: HasDebugCallStack
=> HsExpr GhcRn -> HsExpr GhcTc
-> TcRhoType -- ^ Actual; a rho-type, not a sigma-type
-> ExpRhoType -- ^ Expected
-> TcM (HsExpr GhcTc)
tcWrapResultMono rn_expr expr act_ty res_ty
= do { co <- tcSubTypeMono rn_expr act_ty res_ty
; return (mkHsWrapCo co expr) }
-- | A version of 'tcSubType' to use when the actual type is a rho-type,
-- so that no instantiation is needed.
tcSubTypeMono :: HasDebugCallStack
=> HsExpr GhcRn
-> TcRhoType -- ^ Actual; a rho-type, not a sigma-type
-> ExpRhoType -- ^ Expected
-> TcM TcCoercionN
tcSubTypeMono rn_expr act_ty exp_ty
= assertPpr (isDeepRhoTy act_ty)
(vcat [ text "Actual type is not a (deep) rho-type."
, text "act_ty:" <+> ppr act_ty
, text "rn_expr:" <+> ppr rn_expr]) $
case exp_ty of
Infer inf_res -> fillInferResultNoInst act_ty inf_res
Check exp_ty -> unifyType (Just $ HsExprRnThing rn_expr) act_ty exp_ty
------------------------
tcSubTypePat :: CtOrigin -> UserTypeCtxt
-> ExpSigmaType -> TcSigmaType -> TcM HsWrapper
-- Used in patterns; polarity is backwards compared
-- to tcSubType
-- If wrap = tc_sub_type_et t1 t2
-- => wrap :: t1 ~> t2
tcSubTypePat inst_orig ctxt (Check ty_actual) ty_expected
= tc_sub_type unifyTypeET inst_orig ctxt ty_actual ty_expected
tcSubTypePat _ _ (Infer inf_res) ty_expected
= do { co <- fillInferResultNoInst ty_expected inf_res
-- In patterns we do not instantatiate
; return (mkWpCastN (mkSymCo co)) }
---------------
-- | A subtype check that performs deep subsumption.
-- See also 'tcSubTypeMono', for when no instantiation is required.
tcSubTypeDS :: HsExpr GhcRn
-> TcRhoType -- Actual type -- a rho-type not a sigma-type
-> TcRhoType -- Expected type
-- DeepSubsumption <=> when checking, this type
-- is deeply skolemised
-> TcM HsWrapper
-- Only one call site, in GHC.Tc.Gen.App.tcApp
tcSubTypeDS rn_expr act_rho exp_rho
= tc_sub_type_deep Top (unifyExprType rn_expr) orig GenSigCtxt act_rho exp_rho
where
orig = exprCtOrigin rn_expr
---------------
-- | Checks that the 'actual' type is more polymorphic than the 'expected' type.
tcSubType :: CtOrigin -- ^ Used when instantiating
-> UserTypeCtxt -- ^ Used when skolemising
-> Maybe TypedThing -- ^ The expression that has type 'actual' (if known)
-> TcSigmaType -- ^ Actual type
-> ExpRhoType -- ^ Expected type
-> TcM HsWrapper
tcSubType inst_orig ctxt m_thing ty_actual res_ty
= case res_ty of
Check ty_expected -> tc_sub_type (unifyType m_thing) inst_orig ctxt
ty_actual ty_expected
Infer inf_res -> fillInferResult inst_orig ty_actual inf_res
---------------
tcSubTypeSigma :: CtOrigin -- where did the actual type arise / why are we
-- doing this subtype check?
-> UserTypeCtxt -- where did the expected type arise?
-> TcSigmaType -> TcSigmaType -> TcM HsWrapper
-- External entry point, but no ExpTypes on either side
-- Checks that actual <= expected
-- Returns HsWrapper :: actual ~ expected
tcSubTypeSigma orig ctxt ty_actual ty_expected
= tc_sub_type (unifyType Nothing) orig ctxt ty_actual ty_expected
---------------
tcSubTypeAmbiguity :: UserTypeCtxt -- Where did this type arise
-> TcSigmaType -> TcSigmaType -> TcM HsWrapper
-- See Note [Ambiguity check and deep subsumption]
tcSubTypeAmbiguity ctxt ty_actual ty_expected
= tc_sub_type_ds Top Shallow (unifyType Nothing)
(AmbiguityCheckOrigin ctxt)
ctxt ty_actual ty_expected
---------------
addSubTypeCtxt :: TcType -> ExpType -> TcM a -> TcM a
addSubTypeCtxt ty_actual ty_expected thing_inside
| isRhoTy ty_actual -- If there is no polymorphism involved, the
, isRhoExpTy ty_expected -- TypeEqOrigin stuff (added by the _NC functions)
= thing_inside -- gives enough context by itself
| otherwise
= addErrCtxtM mk_msg thing_inside
where
mk_msg tidy_env
= do { (tidy_env, ty_actual) <- zonkTidyTcType tidy_env ty_actual
; ty_expected <- readExpType ty_expected
-- A worry: might not be filled if we're debugging. Ugh.
; (tidy_env, ty_expected) <- zonkTidyTcType tidy_env ty_expected
; return (tidy_env, SubTypeCtxt ty_expected ty_actual) }
---------------
tc_sub_type :: (TcType -> TcType -> TcM TcCoercionN) -- How to unify
-> CtOrigin -- Used when instantiating
-> UserTypeCtxt -- Used when skolemising
-> TcSigmaType -- Actual; a sigma-type
-> TcSigmaType -- Expected; also a sigma-type
-> TcM HsWrapper
-- Checks that actual_ty is more polymorphic than expected_ty
-- If wrap = tc_sub_type t1 t2
-- => wrap :: t1 ~> t2
--
-- The "how to unify argument" is always a call to `uType TypeLevel orig`,
-- but with different ways of constructing the CtOrigin `orig` from
-- the argument types and context.
----------------------
tc_sub_type unify inst_orig ctxt ty_actual ty_expected
= do { ds_flag <- getDeepSubsumptionFlag
; tc_sub_type_ds Top ds_flag unify inst_orig ctxt ty_actual ty_expected }
----------------------
tc_sub_type_ds :: Position p -- ^ position in the type (for error messages only)
-> DeepSubsumptionFlag
-> (TcType -> TcType -> TcM TcCoercionN)
-> CtOrigin -> UserTypeCtxt -> TcSigmaType
-> TcSigmaType -> TcM HsWrapper
-- tc_sub_type_ds is the main subsumption worker function
-- It takes an explicit DeepSubsumptionFlag
tc_sub_type_ds pos ds_flag unify inst_orig ctxt ty_actual ty_expected
| definitely_poly ty_expected -- See Note [Don't skolemise unnecessarily]
, isRhoTyDS ds_flag ty_actual
= do { traceTc "tc_sub_type (drop to equality)" $
vcat [ text "ty_actual =" <+> ppr ty_actual
, text "ty_expected =" <+> ppr ty_expected ]
; mkWpCastN <$>
unify ty_actual ty_expected }
| otherwise -- This is the general case
= do { traceTc "tc_sub_type (general case)" $
vcat [ text "ty_actual =" <+> ppr ty_actual
, text "ty_expected =" <+> ppr ty_expected ]
; (sk_wrap, inner_wrap)
<- tcSkolemise ds_flag ctxt ty_expected $ \sk_rho ->
case ds_flag of
Deep -> tc_sub_type_deep pos unify inst_orig ctxt ty_actual sk_rho
Shallow -> tc_sub_type_shallow unify inst_orig ty_actual sk_rho
; return (sk_wrap <.> inner_wrap) }
----------------------
tc_sub_type_shallow :: (TcType -> TcType -> TcM TcCoercionN)
-> CtOrigin
-> TcSigmaType
-> TcRhoType -- Skolemised (shallow-ly)
-> TcM HsWrapper
tc_sub_type_shallow unify inst_orig ty_actual sk_rho
= do { (wrap, rho_a) <- topInstantiate inst_orig ty_actual
; cow <- unify rho_a sk_rho
; return (mkWpCastN cow <.> wrap) }
----------------------
definitely_poly :: TcType -> Bool
-- A very conservative test:
-- see Note [Don't skolemise unnecessarily]
definitely_poly ty
| (tvs, theta, tau) <- tcSplitSigmaTy ty
, (tv:_) <- tvs -- At least one tyvar
, null theta -- No constraints; see (DP1)
, tv `isInjectiveInType` tau
-- The tyvar actually occurs (DP2),
-- and occurs in an injective position (DP3).
= True
| otherwise
= False
{- Note [Don't skolemise unnecessarily]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we are trying to solve
ty_actual <= ty_expected
(Char->Char) <= (forall a. a->a)
We could skolemise the 'forall a', and then complain
that (Char ~ a) is insoluble; but that's a pretty obscure
error. It's better to say that
(Char->Char) ~ (forall a. a->a)
fails.
If we prematurely go to equality we'll reject a program we should
accept (e.g. #13752). So the test (which is only to improve error
message) is very conservative:
* ty_actual is /definitely/ monomorphic: see `definitely_mono`
This definitely_mono test comes in "shallow" and "deep" variants
* ty_expected is /definitely/ polymorphic: see `definitely_poly`
This definitely_poly test is more subtle than you might think.
Here are three cases where expected_ty looks polymorphic, but
isn't, and where it would be /wrong/ to switch to equality:
(DP1) (Char->Char) <= (forall a. (a~Char) => a -> a)
(DP2) (Char->Char) <= (forall a. Char -> Char)
(DP3) (Char->Char) <= (forall a. F [a] Char -> Char)
where type instance F [x] t = t
Note [Coercion errors in tcSubMult]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
At the moment, we insist that all sub-multiplicity tests turn out
(once the typechecker has finished its work) to be equalities,
i.e. implementable by ReflCo. Why? Because our type system has
no way to express non-Refl sub-multiplicities.
How can we check that every call to `tcSubMult` returns `Refl`?
It might not be `Refl` *yet*.
[TODO: add counterexample #25130]
So we take the following approach
* In `tcEqMult`:
- Emit a perfectly ordinary Wanted equality constraint for the equality,
returning a coercion.
- Wrap that coercion with `DE_Multiplicity` to make a `DelayedError`, and put
that delayed error into `wc_errors` of the current WantedConstraints. This
is done by `ensureReflMultiplicityCo`.
* When solving constraints, discard any `DE_Multiplicity` errors that wrap a
reflective coercion, of kind `ty ~ ty`. This is done in
`GHC.Tc.Solver.simplifyDelayedErrors`
* After constraint solving is complete report an error if there are any
remaining `DE_Multiplicity` errors. See
`GHC.Tc.Errors.reportMultiplicityCoercionErrs`
Wrinkles
(DME1) If the multiplicity constraint is /solved/, but with a non-reflective
coercion, we'll have just a `DE_Multiplicity` error left over.
But if the multiplicity constraint is /unsolved/ (e.g. ManyTy ~ OneTy), we
will have /both/ an unsolved Wanted in `wc_simple`, /and/ a `DE_Multiplicity`
in `wc_errors`. We don't want to report both. Solution: suppress all
`DE_Multiplicity` constraints if there are any unsolved wanted.
This way, the delayed error is indeed only reported when the constraint is
solved with a non-reflexivity coercion.
An alternative would be to have a kind of constraint which can only produce
trivial evidence. This would allow such checks to happen in the constraint
solver (#18756). This would be similar to the existing setup for Concrete, see
Note [The Concrete mechanism] in GHC.Tc.Utils.Concrete (PHASE 1 in particular).
-}
tcSubMult :: CtOrigin -> Mult -> Mult -> TcM ()
tcSubMult origin w_actual w_expected
| Just (w1, w2) <- isMultMul w_actual =
do { tcSubMult origin w1 w_expected
; tcSubMult origin w2 w_expected }
-- Currently, we consider p*q and sup p q to be equal. Therefore, p*q <= r is
-- equivalent to p <= r and q <= r. For other cases, we approximate p <= q by p
-- ~ q. This is not complete, but it's sound. See also Note [Overapproximating
-- multiplicities] in Multiplicity.
tcSubMult origin w_actual w_expected =
case submult w_actual w_expected of
Submult -> return ()
Unknown -> tcEqMult origin w_actual w_expected
tcEqMult :: CtOrigin -> Mult -> Mult -> TcM ()
tcEqMult origin w_actual w_expected = do
{
-- Note that here we do not call to `submult`, so we check
-- for strict equality.
; coercion <- unifyTypeAndEmit TypeLevel origin w_actual w_expected
-- See Note [Coercion errors in tcSubMult].
; ensureReflMultiplicityCo coercion origin }
{- *********************************************************************
* *
Deep subsumption
* *
********************************************************************* -}
{- Note [Deep subsumption]
~~~~~~~~~~~~~~~~~~~~~~~~~~
The DeepSubsumption extension, documented here
https://github.com/ghc-proposals/ghc-proposals/pull/511.
makes a best-efforts attempt implement deep subsumption as it was
prior to the Simplify Subsumption proposal:
https://github.com/ghc-proposals/ghc-proposals/pull/287
The effects are in these main places:
1. In the subsumption check, tcSubType, we must do deep skolemisation:
see the call to tcSkolemise Deep in tc_sub_type_deep
2. In tcPolyExpr we must do deep skolemisation:
see the call to tcSkolemise in tcSkolemiseExpType
3. for expression type signatures (e :: ty), and functions with type
signatures (e.g. f :: ty; f = e), we must deeply skolemise the type;
see the call to tcDeeplySkolemise in tcSkolemiseScoped.
4. In GHC.Tc.Gen.App.tcApp we call tcSubTypeDS to match the result
type. Without deep subsumption, tcSubTypeMono would be sufficent.
In all these cases note that the deep skolemisation must be done /first/.
Consider (1)
(forall a. Int -> a -> a) <= Int -> (forall b. b -> b)
We must skolemise the `forall b` before instantiating the `forall a`.
See also Note [Deep skolemisation].
Wrinkles:
(DS1) Note that we /always/ use shallow subsumption in the ambiguity check.
See Note [Ambiguity check and deep subsumption].
(DS2) When doing deep subsumption, we must be careful not to needlessly
drop down to unification, e.g. in cases such as:
(Bool -> ∀ d. d->d) <= alpha beta gamma
See Note [FunTy vs non-FunTy case in tc_sub_type_deep].
(DS3) The interaction between deep subsumption and required foralls
(forall a -> ty) is a bit subtle. See #24696 and
Note [Deep subsumption and required foralls]
Note [Deep skolemisation]
~~~~~~~~~~~~~~~~~~~~~~~~~
deeplySkolemise decomposes and skolemises a type, returning a type
with all its arrows visible (ie not buried under foralls)
Examples:
deeplySkolemise (Int -> forall a. Ord a => blah)
= ( wp, [a], [d:Ord a], Int -> blah )
where wp = \x:Int. /\a. \(d:Ord a). <hole> x
deeplySkolemise (forall a. Ord a => Maybe a -> forall b. Eq b => blah)
= ( wp, [a,b], [d1:Ord a,d2:Eq b], Maybe a -> blah )
where wp = /\a.\(d1:Ord a).\(x:Maybe a)./\b.\(d2:Ord b). <hole> x
In general,
if deeplySkolemise ty = (wrap, tvs, evs, rho)
and e :: rho
then wrap e :: ty
and 'wrap' binds tvs, evs
ToDo: this eta-abstraction plays fast and loose with termination,
because it can introduce extra lambdas. Maybe add a `seq` to
fix this
Note [FunTy vs FunTy case in tc_sub_type_deep]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The goal of tc_sub_type_deep is to produce an HsWrapper that "proves" that the
actual type is a subtype of the expected type. The most important case is how
we deal with function arrows. Suppose we have:
ty_actual = act_arg -> act_res
ty_expected = exp_arg -> exp_res
To produce fun_wrap :: (act_arg -> act_res) ~> (exp_arg -> exp_res), we use
the fact that the function arrow is contravariant in its argument type and
covariant in its result type. Thus we recursively perform subtype checks
on the argument types (with actual/expected switched) and the result types,
to get:
arg_wrap :: exp_arg ~> act_arg -- NB: expected/actual have switched sides
res_wrap :: act_res ~> exp_res
Then fun_wrap = mkWpFun arg_wrap res_wrap.
Wrinkle [Representation-polymorphism checking during subtyping]
Inserting a WpFun HsWrapper amounts to impedance matching in deep subsumption
via eta-expansion:
f ==> \ (x :: exp_arg) -> res_wrap [ f (arg_wrap [x]) ]
As we produce a lambda, we must enforce the representation polymorphism
invariants described in Note [Representation polymorphism invariants] in GHC.Core.
That is, we must ensure that both x (the lambda binder) and (arg_wrap [x]) (the function argument)
have a fixed runtime representation.
Note however that desugaring mkWpFun does not always introduce a lambda: if
both the argument and result HsWrappers are casts, then a FunCo cast suffices,
in which case we should not perform representation-polymorphism checking.
This means that, in the FunTy/FunTy case of tc_sub_type_deep, we can skip
the representation-polymorphism checks if the produced argument and result
wrappers are identities or casts.
It is important to do so, otherwise we reject valid programs.
Here's a contrived example (there are undoubtedly more natural examples)
(see testsuite/tests/rep-poly/NoEtaRequired):
type Id :: k -> k
type family Id a where
type T :: TYPE r -> TYPE (Id r)
type family T a where
test :: forall r (a :: TYPE r). a :~~: T a -> ()
test HRefl =
let
f :: (a -> a) -> ()
f _ = ()
g :: T a -> T a
g = undefined
in f g
We don't need to eta-expand `g` to make `f g` typecheck; a cast suffices.
Hence we should not perform representation-polymorphism checks; they would
fail here.
Note [Setting the argument context]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider we are doing the ambiguity check for the (bogus)
f :: (forall a b. C b => a -> a) -> Int
We'll call
tcSubType ((forall a b. C b => a->a) -> Int )
((forall a b. C b => a->a) -> Int )
with a UserTypeCtxt of (FunSigCtxt "f"). Then we'll do the co/contra thing
on the argument type of the (->) -- and at that point we want to switch
to a UserTypeCtxt of GenSigCtxt. Why?
* Error messages. If we stick with FunSigCtxt we get errors like
* Could not deduce: C b
from the context: C b0
bound by the type signature for:
f :: forall a b. C b => a->a
But of course f does not have that type signature!
Example tests: T10508, T7220a, Simple14
* Implications. We may decide to build an implication for the whole
ambiguity check, but we don't need one for each level within it,
and TcUnify.alwaysBuildImplication checks the UserTypeCtxt.
See Note [When to build an implication]
Note [Multiplicity in deep subsumption]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider
t1 ->{mt} t2 <= s1 ->{ms} s2
At the moment we /unify/ ms~mt, via tcEqMult.
Arguably we should use `tcSubMult`. But then if mt=m0 (a unification
variable) and ms=Many, `tcSubMult` is a no-op (since anything is a
sub-multiplicty of Many). But then `m0` may never get unified with
anything. It is then skolemised by the zonker; see GHC.HsToCore.Binds
Note [Free tyvars on rule LHS]. So we in RULE foldr/app in GHC.Base
we get this
"foldr/app" [1] forall ys m1 m2. foldr (\x{m1} \xs{m2}. (:) x xs) ys
= \xs -> xs ++ ys
where we eta-expanded that (:). But now foldr expects an argument
with ->{Many} and gets an argument with ->{m1} or ->{m2}, and Lint
complains.
The easiest solution was to unify the multiplicities in tc_sub_type_deep,
insisting on equality. This is only in the DeepSubsumption code anyway.
Note [FunTy vs non-FunTy case in tc_sub_type_deep]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider this, without Quick Look, but with Deep Subsumption:
f :: ∀a b c. a b c -> Int
g :: Bool -> ∀d. d -> d
To typecheck the application (f g), we need to do the subsumption test
(Bool -> ∀ d. d->d) <= alpha beta gamma
where alpha, beta, gamma are the unification variables that instantiate a,b,c
(respectively). We must not drop down to unification, or we will reject the call.
Instead, we should only unify alpha := (->), in which case we end up with the
usual FunTy vs FunTy case of Note [FunTy vs FunTy case in tc_sub_type_deep]:
(Bool -> ∀ d. d->d) <= beta -> gamma
which is straightforwardly solved by beta := Bool, using covariance in the return
type of the function arrow, and instantiating the forall before unifying with gamma.
The conclusion is this: when doing a deep subtype check (in tc_sub_type_deep),
if the LHS is a FunTy and the RHS is a rho-type which is not a FunTy,
then unify the RHS with a FunTy and continue by performing a sub-type check on
the LHS vs the new RHS. And vice-versa (if it's the RHS that is a FunTy).
See T11305 and T26225 for examples of when this is important.
Note [Deep subsumption and required foralls]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
A required forall, (forall a -> ty) behaves like a "rho-type", one with no
top-level quantification. In particular, it is neither implicitly instantiated nor
skolemised. So
rid1 :: forall a -> a -> a
rid1 = id
rid2 :: forall a -> a -> a
rid2 a = id
Here `rid2` wll typecheck, but `rid1` will not, because we don't implicitly skolemise
the type.
This "no implicit subsumption nor skolemisation" applies during subsumption.
For example
(forall a. a->a) <= (forall a -> a -> a) -- NOT!
does /not/ hold, because that would require implicitly skoleming the (forall a->).
Note also that, in Core, `eqType` distinguishes between
(forall a. blah) and forall a -> blah)
See discussion on #22762 and these Notes in GHC.Core.TyCo.Compare
* Note [ForAllTy and type equality]
* Note [Comparing visibility]
So during deep subsumption we simply stop (and drop down to equality) when we encounter
a (forall a->). This is a little odd:
* Deep subsumption looks inside invisible foralls (forall a. ty)
* Deep subsumption looks inside arrows (t1 -> t2)
* But it does not look inside required foralls (forall a -> ty)
There is discussion on #24696. How is this implemented?
* In `tc_sub_type_deep`, the calls to `topInstantiate` and `deeplyInstantiate`
instantiate only /invisible/ binders.
* In `tc_sub_type_ds`, the call to `tcSkolemise` skolemises only /invisible/
binders.
Here is a slightly more powerful alternative
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In the story above, if we have
(forall a -> Eq a => a -> a) <= (forall a -> Ord a => a -> a)
we'll reject it, because both are rho-types but they aren't equal. But in the
"drop to equality" stage we could instead see if both rho-types are headed with
(forall a ->) and if so strip that off and go back into deep subsumption.
This is a bit more powerful, but also a bit more complicated, so GHC
doesn't do it yet, awaiting credible user demand. See #24696.
-}
data DeepSubsumptionFlag = Deep | Shallow
instance Outputable DeepSubsumptionFlag where
ppr Deep = text "Deep"
ppr Shallow = text "Shallow"
getDeepSubsumptionFlag :: TcM DeepSubsumptionFlag
getDeepSubsumptionFlag = do { ds <- xoptM LangExt.DeepSubsumption
; if ds then return Deep else return Shallow }
-- | 'tc_sub_type_deep' is where the actual work happens for deep subsumption.
--
-- Given @ty_actual@ (a sigma-type) and @ty_expected@ (deeply skolemised, i.e.
-- a deep rho type), it returns an 'HsWrapper' @wrap :: ty_actual ~> ty_expected@.
tc_sub_type_deep :: HasDebugCallStack
=> Position p -- ^ Position in the type (for error messages only)
-> (TcType -> TcType -> TcM TcCoercionN) -- ^ How to unify
-> CtOrigin -- ^ Used when instantiating
-> UserTypeCtxt -- ^ Used when skolemising
-> TcSigmaType -- ^ Actual; a sigma-type
-> TcRhoType -- ^ Expected; deeply skolemised
-> TcM HsWrapper
-- If wrap = tc_sub_type_deep t1 t2
-- => wrap :: t1 ~> t2
-- Here is where the work actually happens!
-- Precondition: ty_expected is deeply skolemised
tc_sub_type_deep pos unify inst_orig ctxt ty_actual ty_expected
= assertPpr (isDeepRhoTy ty_expected) (ppr ty_expected) $
do { traceTc "tc_sub_type_deep" $
vcat [ text "ty_actual =" <+> ppr ty_actual
, text "ty_expected =" <+> ppr ty_expected ]
; go ty_actual ty_expected }
where
-- 'unwrap' removes top-level type synonyms & looks through filled meta-tyvars
unwrap :: TcType -> TcM TcType
unwrap ty
| Just ty' <- coreView ty
= unwrap ty'
unwrap ty@(TyVarTy tv)
= do { lookup_res <- isFilledMetaTyVar_maybe tv
; case lookup_res of
Just ty' -> unwrap ty'
Nothing -> return ty }
unwrap ty = return ty
go, go1 :: TcType -> TcType -> TcM HsWrapper
go ty_a ty_e =
do { ty_a' <- unwrap ty_a
; ty_e' <- unwrap ty_e
; go1 ty_a' ty_e' }
-- If ty_actual is not a rho-type, instantiate it first; otherwise
-- unification has no chance of succeeding.
go1 ty_a ty_e
| let (tvs, theta, _) = tcSplitSigmaTy ty_a
, not (null tvs && null theta)
= do { (in_wrap, in_rho) <- topInstantiate inst_orig ty_a
; body_wrap <- go in_rho ty_e
; return (body_wrap <.> in_wrap) }
-- Main case: FunTy vs FunTy. go_fun does the work.
go1 (FunTy { ft_af = af1, ft_mult = act_mult, ft_arg = act_arg, ft_res = act_res })
(FunTy { ft_af = af2, ft_mult = exp_mult, ft_arg = exp_arg, ft_res = exp_res })
| isVisibleFunArg af1
, isVisibleFunArg af2
= go_fun af1 act_mult act_arg act_res
af2 exp_mult exp_arg exp_res
-- See Note [FunTy vs non-FunTy case in tc_sub_type_deep]
go1 (FunTy { ft_af = af1, ft_mult = act_mult, ft_arg = act_arg, ft_res = act_res }) ty_e
| isVisibleFunArg af1
= do { exp_mult <- newMultiplicityVar
; exp_arg <- newOpenFlexiTyVarTy -- NB: no FRR check needed; we might not need to eta-expand
; exp_res <- newOpenFlexiTyVarTy
; let exp_funTy = FunTy { ft_af = af1, ft_mult = exp_mult, ft_arg = exp_arg, ft_res = exp_res }
; unify_wrap <- just_unify exp_funTy ty_e
; fun_wrap <- go_fun af1 act_mult act_arg act_res af1 exp_mult exp_arg exp_res
; return $ unify_wrap <.> fun_wrap
-- unify_wrap :: exp_funTy ~> ty_e
-- fun_wrap :: ty_a ~> exp_funTy
}
go1 ty_a (FunTy { ft_af = af2, ft_mult = exp_mult, ft_arg = exp_arg, ft_res = exp_res })
| isVisibleFunArg af2
= do { act_mult <- newMultiplicityVar
; act_arg <- newOpenFlexiTyVarTy -- NB: no FRR check needed; we might not need to eta-expand
; act_res <- newOpenFlexiTyVarTy
; let act_funTy = FunTy { ft_af = af2, ft_mult = act_mult, ft_arg = act_arg, ft_res = act_res }
; unify_wrap <- just_unify ty_a act_funTy
; fun_wrap <- go_fun af2 act_mult act_arg act_res af2 exp_mult exp_arg exp_res
; return $ fun_wrap <.> unify_wrap
-- unify_wrap :: ty_a ~> act_funTy
-- fun_wrap :: act_funTy ~> ty_e
}
-- Otherwise, revert to unification.
go1 ty_a ty_e = just_unify ty_a ty_e
just_unify ty_a ty_e = do { cow <- unify ty_a ty_e
; return (mkWpCastN cow) }
-- FunTy/FunTy case: this is where we insert any necessary eta-expansions.
go_fun :: FunTyFlag -> Mult -> TcType -> TcType -- actual FunTy
-> FunTyFlag -> Mult -> TcType -> TcType -- expected FunTy
-> TcM HsWrapper
go_fun act_af act_mult act_arg act_res exp_af exp_mult exp_arg exp_res
-- See Note [FunTy vs FunTy case in tc_sub_type_deep]
= do { arg_wrap <- tc_sub_type_ds (Argument pos) Deep unify given_orig GenSigCtxt exp_arg act_arg
-- GenSigCtxt: See Note [Setting the argument context]
; res_wrap <- tc_sub_type_deep (Result pos) unify inst_orig ctxt act_res exp_res
; mkWpFun_FRR unify pos
act_af act_mult act_arg act_res
exp_af exp_mult exp_arg exp_res
arg_wrap res_wrap
}
where
given_orig = GivenOrigin (SigSkol GenSigCtxt exp_arg [])
-- | Like 'mkWpFun', except that it performs the necessary
-- representation-polymorphism checks on the argument type in the case that
-- we introduce a lambda abstraction.
mkWpFun_FRR
:: (TcType -> TcType -> TcM TcCoercionN) -- ^ how to unify
-> Position p
-> FunTyFlag -> Type -> TcType -> Type -- actual FunTy
-> FunTyFlag -> Type -> TcType -> Type -- expected FunTy
-> HsWrapper -- ^ exp_arg ~> act_arg
-> HsWrapper -- ^ act_res ~> exp_res
-> TcM HsWrapper -- ^ act_funTy ~> exp_funTy
mkWpFun_FRR unify pos act_af act_mult act_arg act_res exp_af exp_mult exp_arg exp_res arg_wrap res_wrap
= do { ((exp_arg_co, exp_arg_frr), (act_arg_co, _act_arg_frr)) <-
if needs_frr_checks
-- See Wrinkle [Representation-polymorphism checking during subtyping]
then do { exp_frr_wrap <- hasFixedRuntimeRep (frr_ctxt True ) exp_arg
; act_frr_wrap <- hasFixedRuntimeRep (frr_ctxt False) act_arg
; return (exp_frr_wrap, act_frr_wrap) }
else return ((mkNomReflCo exp_arg, exp_arg), (mkNomReflCo act_arg, act_arg))
-- Enforce equality of multiplicities (not the more natural sub-multiplicity).
-- See Note [Multiplicity in deep subsumption]
; act_arg_mult_co <- unify act_mult exp_mult
-- NB: don't use tcEqMult: that would require the evidence for
-- equality to be Refl, but it might well not be (#26332).
; let
exp_arg_fun_co =
mkFunCo Nominal exp_af
(mkReflCo Nominal exp_mult)
(mkSymCo exp_arg_co)
(mkReflCo Nominal exp_res)
act_arg_fun_co =
mkFunCo Nominal act_af
act_arg_mult_co
act_arg_co
(mkReflCo Nominal act_res)
arg_wrap_frr =
mkWpCastN (mkSymCo exp_arg_co) <.> arg_wrap <.> mkWpCastN act_arg_co
-- exp_arg_co :: exp_arg ~> exp_arg_frr
-- act_arg_co :: act_arg ~> act_arg_frr
-- arg_wrap :: exp_arg ~> act_arg
-- arg_wrap_frr :: exp_arg_frr ~> act_arg_frr
; return $
mkWpCastN exp_arg_fun_co
<.>
mkWpFun arg_wrap_frr res_wrap (Scaled exp_mult exp_arg_frr) exp_res
<.>
mkWpCastN act_arg_fun_co
}
where
needs_frr_checks :: Bool
needs_frr_checks =
not (hole_or_cast arg_wrap)
||
not (hole_or_cast res_wrap)
hole_or_cast :: HsWrapper -> Bool
hole_or_cast WpHole = True
hole_or_cast (WpCast {}) = True
hole_or_cast _ = False
frr_ctxt :: Bool -> FixedRuntimeRepContext
frr_ctxt is_exp_ty =
FRRDeepSubsumption
{ frrDSExpected = is_exp_ty
, frrDSPosition = pos
}
-----------------------
deeplySkolemise :: SkolemInfo -> TcSigmaType
-> TcM ( HsWrapper
, [(Name,TcInvisTVBinder)] -- All skolemised variables
, [EvVar] -- All "given"s
, TcRhoType )
-- See Note [Deep skolemisation]
deeplySkolemise skol_info ty
= go init_subst ty
where
init_subst = mkEmptySubst (mkInScopeSet (tyCoVarsOfType ty))
go subst ty
| Just (arg_tys, bndrs, theta, ty') <- tcDeepSplitSigmaTy_maybe ty
= do { let arg_tys' = substScaledTys subst arg_tys
; ids1 <- newSysLocalIds (fsLit "dk") arg_tys'
; (subst', bndrs1) <- tcInstSkolTyVarBndrsX skol_info subst bndrs
; ev_vars1 <- newEvVars (substTheta subst' theta)
; (wrap, tvs_prs2, ev_vars2, rho) <- go subst' ty'
; let tvs = binderVars bndrs
tvs1 = binderVars bndrs1
tv_prs1 = map tyVarName tvs `zip` bndrs1
; return ( mkWpEta ids1 (mkWpTyLams tvs1
<.> mkWpEvLams ev_vars1
<.> wrap)
, tv_prs1 ++ tvs_prs2
, ev_vars1 ++ ev_vars2
, mkScaledFunTys arg_tys' rho ) }
| otherwise
= return (idHsWrapper, [], [], substTy subst ty)
-- substTy is a quick no-op on an empty substitution
dsInstantiate :: CtOrigin -> TcType -> TcM (HsWrapper, Type)
-- Do topInstantiate or deeplyInstantiate, depending on -XDeepSubsumption
dsInstantiate orig ty
= do { ds_flag <- getDeepSubsumptionFlag
; case ds_flag of
Shallow -> topInstantiate orig ty
Deep -> deeplyInstantiate orig ty }
deeplyInstantiate :: CtOrigin -> TcType -> TcM (HsWrapper, Type)
-- Instantiate invisible foralls, even ones nested
-- (to the right) under arrows
deeplyInstantiate orig ty
= go init_subst ty
where
init_subst = mkEmptySubst (mkInScopeSet (tyCoVarsOfType ty))
go subst ty
| Just (arg_tys, bndrs, theta, rho) <- tcDeepSplitSigmaTy_maybe ty
= do { let tvs = binderVars bndrs
; (subst', tvs') <- newMetaTyVarsX subst tvs
; let arg_tys' = substScaledTys subst' arg_tys
theta' = substTheta subst' theta
; ids1 <- newSysLocalIds (fsLit "di") arg_tys'
; wrap1 <- instCall orig (mkTyVarTys tvs') theta'
; (wrap2, rho2) <- go subst' rho
; return (mkWpEta ids1 (wrap2 <.> wrap1),
mkScaledFunTys arg_tys' rho2) }
| otherwise
= do { let ty' = substTy subst ty
; return (idHsWrapper, ty') }
tcDeepSplitSigmaTy_maybe
:: TcSigmaType -> Maybe ([Scaled TcType], [TcInvisTVBinder], ThetaType, TcSigmaType)
-- Looks for a *non-trivial* quantified type, under zero or more function arrows
-- By "non-trivial" we mean either tyvars or constraints are non-empty
tcDeepSplitSigmaTy_maybe ty
= go ty
where
go ty | Just (arg_ty, res_ty) <- tcSplitFunTy_maybe ty
, Just (arg_tys, tvs, theta, rho) <- go res_ty
= Just (arg_ty:arg_tys, tvs, theta, rho)
| (tvs, theta, rho) <- tcSplitSigmaTyBndrs ty
, not (null tvs && null theta)
= Just ([], tvs, theta, rho)
| otherwise = Nothing
isDeepRhoTy :: TcType -> Bool
-- True if there are no foralls or (=>) at the top, or nested under
-- arrows to the right. e.g
-- forall a. a False
-- Int -> forall a. a False
-- (forall a. a) -> Int True
-- Returns True iff tcDeepSplitSigmaTy_maybe returns Nothing
isDeepRhoTy ty
| not (isRhoTy ty) = False -- Foralls or (=>) at top
| Just (_, res) <- tcSplitFunTy_maybe ty = isDeepRhoTy res
| otherwise = True -- No forall, (=>), or (->) at top
isRhoTyDS :: DeepSubsumptionFlag -> TcType -> Bool
isRhoTyDS ds_flag ty
= case ds_flag of
Shallow -> isRhoTy ty -- isRhoTy: no top level forall or (=>)
Deep -> isDeepRhoTy ty -- "deep" version: no nested forall or (=>)
{-
************************************************************************
* *
Boxy unification
* *
************************************************************************
The exported functions are all defined as versions of some
non-exported generic functions.
-}
unifyExprType :: HsExpr GhcRn -> TcType -> TcType -> TcM TcCoercionN
unifyExprType rn_expr ty1 ty2
= unifyType (Just (HsExprRnThing rn_expr)) ty1 ty2
unifyType :: Maybe TypedThing -- ^ If present, the thing that has type ty1
-> TcTauType -> TcTauType -- ty1 (actual), ty2 (expected)
-> TcM TcCoercionN -- :: ty1 ~# ty2
-- Actual and expected types
-- Returns a coercion : ty1 ~ ty2
unifyType thing ty1 ty2
= unifyTypeAndEmit TypeLevel origin ty1 ty2
where
origin = TypeEqOrigin { uo_actual = ty1
, uo_expected = ty2
, uo_thing = thing
, uo_visible = True }
unifyInvisibleType :: TcTauType -> TcTauType -- ty1 (actual), ty2 (expected)
-> TcM TcCoercionN -- :: ty1 ~# ty2
-- Actual and expected types
-- Returns a coercion : ty1 ~ ty2
unifyInvisibleType ty1 ty2
= unifyTypeAndEmit TypeLevel origin ty1 ty2
where
origin = TypeEqOrigin { uo_actual = ty1
, uo_expected = ty2
, uo_thing = Nothing
, uo_visible = False } -- This is the "invisible" bit
unifyTypeET :: TcTauType -> TcTauType -> TcM CoercionN
-- Like unifyType, but swap expected and actual in error messages
-- This is used when typechecking patterns
unifyTypeET ty1 ty2
= unifyTypeAndEmit TypeLevel origin ty1 ty2
where
origin = TypeEqOrigin { uo_actual = ty2 -- NB swapped
, uo_expected = ty1 -- NB swapped
, uo_thing = Nothing
, uo_visible = True }
unifyKind :: Maybe TypedThing -> TcKind -> TcKind -> TcM CoercionN
unifyKind mb_thing ty1 ty2
= unifyTypeAndEmit KindLevel origin ty1 ty2
where
origin = TypeEqOrigin { uo_actual = ty1
, uo_expected = ty2
, uo_thing = mb_thing
, uo_visible = True }
unifyTypeAndEmit :: TypeOrKind -> CtOrigin -> TcType -> TcType -> TcM CoercionN
-- Make a ref-cell, unify, emit the collected constraints
unifyTypeAndEmit t_or_k orig ty1 ty2
= do { ref <- newTcRef emptyBag
; loc <- getCtLocM orig (Just t_or_k)
; let env = UE { u_loc = loc, u_role = Nominal
, u_rewriters = emptyRewriterSet -- ToDo: check this
, u_defer = ref, u_unified = Nothing }
-- The hard work happens here
; co <- uType env ty1 ty2
; cts <- readTcRef ref
; unless (null cts) (emitSimples cts)
; return co }
{-
%************************************************************************
%* *
uType and friends
%* *
%************************************************************************
Note [The eager unifier]
~~~~~~~~~~~~~~~~~~~~~~~~
The eager unifier, `uType`, is called by
* The constraint generator (e.g. in GHC.Tc.Gen.Expr),
via the wrappers `unifyType`, `unifyKind` etc
* The constraint solver (e.g. in GHC.Tc.Solver.Equality),
via `GHC.Tc.Solver.Monad.wrapUnifierTcS`.
`uType` runs in the TcM monad, but it carries a UnifyEnv that tells it
what to do when unifying a variable or deferring a constraint. Specifically,
* it collects deferred constraints in `u_defer`, and
* it records which unification variables it has unified in `u_unified`
Then it is up to the wrappers (one for the constraint generator, one for
the constraint solver) to deal with these collected sets.
Although `uType` runs in the TcM monad for convenience, really it could
operate just with the ability to
* write to the accumulators of deferred constraints
and unification variables in UnifyEnv.
* read and update existing unification variables
* zonk types befire unifying (`zonkTcType` in `uUnfilledVar`, and
`zonkTyCoVarKind` in `uUnfilledVar1`
* create fresh coercion holes (`newCoercionHole`)
* emit tracing info for debugging
* look at the ambient TcLevel: `getTcLevel`
A job for the future.
-}
data UnifyEnv
= UE { u_role :: Role
, u_loc :: CtLoc
, u_rewriters :: RewriterSet
-- Deferred constraints
, u_defer :: TcRef (Bag Ct)
-- Which variables are unified;
-- if Nothing, we don't care
, u_unified :: Maybe (TcRef [TcTyVar])
}
setUEnvRole :: UnifyEnv -> Role -> UnifyEnv
setUEnvRole uenv role = uenv { u_role = role }
updUEnvLoc :: UnifyEnv -> (CtLoc -> CtLoc) -> UnifyEnv
updUEnvLoc uenv@(UE { u_loc = loc }) upd = uenv { u_loc = upd loc }
mkKindEnv :: UnifyEnv -> TcType -> TcType -> UnifyEnv
-- Modify the UnifyEnv to be right for unifing
-- the kinds of these two types
mkKindEnv env@(UE { u_loc = ctloc }) ty1 ty2
= env { u_role = Nominal, u_loc = mkKindEqLoc ty1 ty2 ctloc }
uType, uType_defer
:: UnifyEnv
-> TcType -- ty1 is the *actual* type
-> TcType -- ty2 is the *expected* type
-> TcM CoercionN
-- It is always safe to defer unification to the main constraint solver
-- See Note [Deferred unification]
uType_defer (UE { u_loc = loc, u_defer = ref
, u_role = role, u_rewriters = rewriters })
ty1 ty2 -- ty1 is "actual", ty2 is "expected"
= do { let pred_ty = mkEqPredRole role ty1 ty2
; hole <- newCoercionHole pred_ty
; let ct = mkNonCanonical $ CtWanted $
WantedCt { ctev_pred = pred_ty
, ctev_dest = HoleDest hole
, ctev_loc = loc
, ctev_rewriters = rewriters }
co = HoleCo hole
; updTcRef ref (`snocBag` ct)
-- snocBag: see Note [Work-list ordering] in GHC.Tc.Solver.Equality
-- Error trace only
-- NB. do *not* call mkErrCtxt unless tracing is on,
-- because it is hugely expensive (#5631)
; whenDOptM Opt_D_dump_tc_trace $
do { ctxt <- getErrCtxt
; err_ctxt <- mkErrCtxt emptyTidyEnv ctxt
; traceTc "utype_defer" $
vcat ( ppr role
: debugPprType ty1
: debugPprType ty2
: map pprErrCtxtMsg err_ctxt )
; traceTc "utype_defer2" (ppr co) }
; return co }
--------------
uType env@(UE { u_role = role }) orig_ty1 orig_ty2
| Phantom <- role
= do { kind_co <- uType (mkKindEnv env orig_ty1 orig_ty2)
(typeKind orig_ty1) (typeKind orig_ty2)
; return (mkPhantomCo kind_co orig_ty1 orig_ty2) }
| otherwise
= do { tclvl <- getTcLevel
; traceTc "u_tys" $ vcat
[ text "tclvl" <+> ppr tclvl
, sep [ ppr orig_ty1, text "~" <> ppr role, ppr orig_ty2] ]
; co <- go orig_ty1 orig_ty2
; if isReflCo co
then traceTc "u_tys yields no coercion" Outputable.empty
else traceTc "u_tys yields coercion:" (ppr co)
; return co }
where
go :: TcType -> TcType -> TcM CoercionN
-- The arguments to 'go' are always semantically identical
-- to orig_ty{1,2} except for looking through type synonyms
-- Unwrap casts before looking for variables. This way, we can easily
-- recognize (t |> co) ~ (t |> co), which is nice. Previously, we
-- didn't do it this way, and then the unification above was deferred.
go (CastTy t1 co1) t2
= do { co_tys <- uType env t1 t2
; return (mkCoherenceLeftCo role t1 co1 co_tys) }
go t1 (CastTy t2 co2)
= do { co_tys <- uType env t1 t2
; return (mkCoherenceRightCo role t2 co2 co_tys) }
-- Variables; go for uUnfilledVar
-- Note that we pass in *original* (before synonym expansion),
-- so that type variables tend to get filled in with
-- the most informative version of the type
go (TyVarTy tv1) ty2
= do { lookup_res <- isFilledMetaTyVar_maybe tv1
; case lookup_res of
Just ty1 -> do { traceTc "found filled tyvar" (ppr tv1 <+> text ":->" <+> ppr ty1)
; uType env ty1 orig_ty2 }
Nothing -> uUnfilledVar env NotSwapped tv1 ty2 }
go ty1 (TyVarTy tv2)
= do { lookup_res <- isFilledMetaTyVar_maybe tv2
; case lookup_res of
Just ty2 -> do { traceTc "found filled tyvar" (ppr tv2 <+> text ":->" <+> ppr ty2)
; uType env orig_ty1 ty2 }
Nothing -> uUnfilledVar env IsSwapped tv2 ty1 }
-- See Note [Unifying type synonyms] in GHC.Core.Unify
go ty1@(TyConApp tc1 []) (TyConApp tc2 [])
| tc1 == tc2
= return $ mkReflCo role ty1
-- Now expand synonyms
-- See Note [Expanding synonyms during unification]
--
-- Also NB that we recurse to 'go' so that we don't push a
-- new item on the origin stack. As a result if we have
-- type Foo = Int
-- and we try to unify Foo ~ Bool
-- we'll end up saying "can't match Foo with Bool"
-- rather than "can't match "Int with Bool". See #4535.
go ty1 ty2
| Just ty1' <- coreView ty1 = go ty1' ty2
| Just ty2' <- coreView ty2 = go ty1 ty2'
-- Functions (t1 -> t2) just check the two parts
go (FunTy { ft_af = af1, ft_mult = w1, ft_arg = arg1, ft_res = res1 })
(FunTy { ft_af = af2, ft_mult = w2, ft_arg = arg2, ft_res = res2 })
| isVisibleFunArg af1 -- Do not attempt (c => t); just defer
, af1 == af2 -- Important! See #21530
= do { co_w <- uType (env { u_role = funRole role SelMult }) w1 w2
; co_l <- uType (env { u_role = funRole role SelArg }) arg1 arg2
; co_r <- uType (env { u_role = funRole role SelRes }) res1 res2
; return $ mkNakedFunCo role af1 co_w co_l co_r }
-- Always defer if a type synonym family (type function)
-- is involved. (Data families behave rigidly.)
go ty1@(TyConApp tc1 _) ty2
| isTypeFamilyTyCon tc1 = defer ty1 ty2
go ty1 ty2@(TyConApp tc2 _)
| isTypeFamilyTyCon tc2 = defer ty1 ty2
go (TyConApp tc1 tys1) (TyConApp tc2 tys2)
-- See Note [Mismatched type lists and application decomposition]
| tc1 == tc2, equalLength tys1 tys2
, isInjectiveTyCon tc1 role -- don't look under newtypes at Rep equality
= assertPpr (isGenerativeTyCon tc1 role) (ppr tc1) $
do { traceTc "go-tycon" (ppr tc1 $$ ppr tys1 $$ ppr tys2 $$ ppr (take 10 (tyConRoleListX role tc1)))
; cos <- zipWith4M u_tc_arg (tyConVisibilities tc1) -- Infinite
(tyConRoleListX role tc1) -- Infinite
tys1 tys2
; return $ mkTyConAppCo role tc1 cos }
go (LitTy m) ty@(LitTy n)
| m == n
= return $ mkReflCo role ty
-- See Note [Care with type applications]
-- Do not decompose FunTy against App;
-- it's often a type error, so leave it for the constraint solver
go ty1@(AppTy s1 t1) ty2@(AppTy s2 t2)
= go_app (isNextArgVisible s1) ty1 s1 t1 ty2 s2 t2
go ty1@(AppTy s1 t1) ty2@(TyConApp tc2 ts2)
| Just (ts2', t2') <- snocView ts2
= assert (not (tyConMustBeSaturated tc2)) $
go_app (isNextTyConArgVisible tc2 ts2')
ty1 s1 t1 ty2 (TyConApp tc2 ts2') t2'
go ty1@(TyConApp tc1 ts1) ty2@(AppTy s2 t2)
| Just (ts1', t1') <- snocView ts1
= assert (not (tyConMustBeSaturated tc1)) $
go_app (isNextTyConArgVisible tc1 ts1')
ty1 (TyConApp tc1 ts1') t1' ty2 s2 t2
go ty1@(CoercionTy co1) ty2@(CoercionTy co2)
= do { kco <- uType (mkKindEnv env ty1 ty2)
(coercionType co1) (coercionType co2)
; return $ mkProofIrrelCo role kco co1 co2 }
-- Anything else fails
-- E.g. unifying for-all types, which is relative unusual
go ty1 ty2 = defer ty1 ty2
------------------
defer ty1 ty2 -- See Note [Check for equality before deferring]
| ty1 `tcEqType` ty2 = return (mkReflCo role ty1)
| otherwise = uType_defer env orig_ty1 orig_ty2
------------------
u_tc_arg is_vis role ty1 ty2
= do { traceTc "u_tc_arg" (ppr role $$ ppr ty1 $$ ppr ty2)
; uType env_arg ty1 ty2 }
where
env_arg = env { u_loc = adjustCtLoc is_vis False (u_loc env)
, u_role = role }
------------------
-- For AppTy, decompose only nominal equalities
-- See Note [Decomposing AppTy equalities] in GHC.Tc.Solver.Equality
go_app vis ty1 s1 t1 ty2 s2 t2
| Nominal <- role
= -- Unify arguments t1/t2 before function s1/s2, because
-- the former have smaller kinds, and hence simpler error messages
-- c.f. GHC.Tc.Solver.Equality.can_eq_app
-- Example: test T8603
do { let env_arg = env { u_loc = adjustCtLoc vis False (u_loc env) }
; co_t <- uType env_arg t1 t2
; co_s <- uType env s1 s2
; return $ mkAppCo co_s co_t }
| otherwise
= defer ty1 ty2
{- Note [Check for equality before deferring]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Particularly in ambiguity checks we can get equalities like (ty ~ ty).
If ty involves a type function we may defer, which isn't very sensible.
An egregious example of this was in test T9872a, which has a type signature
Proxy :: Proxy (Solutions Cubes)
Doing the ambiguity check on this signature generates the equality
Solutions Cubes ~ Solutions Cubes
and currently the constraint solver normalises both sides at vast cost.
This little short-cut in 'defer' helps quite a bit.
Note [Care with type applications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Note: type applications need a bit of care!
They can match FunTy and TyConApp, so use splitAppTy_maybe
NB: we've already dealt with type variables and Notes,
so if one type is an App the other one jolly well better be too
Note [Mismatched type lists and application decomposition]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When we find two TyConApps, you might think that the argument lists
are guaranteed equal length. But they aren't. Consider matching
w (T x) ~ Foo (T x y)
We do match (w ~ Foo) first, but in some circumstances we simply create
a deferred constraint; and then go ahead and match (T x ~ T x y).
This came up in #3950.
So either
(a) either we must check for identical argument kinds
when decomposing applications,
(b) or we must be prepared for ill-kinded unification sub-problems
Currently we adopt (b) since it seems more robust -- no need to maintain
a global invariant.
Note [Expanding synonyms during unification]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We expand synonyms during unification, but:
* We expand *after* the variable case so that we tend to unify
variables with un-expanded type synonym. This just makes it
more likely that the inferred types will mention type synonyms
understandable to the user
* Similarly, we expand *after* the CastTy case, just in case the
CastTy wraps a variable.
* We expand *before* the TyConApp case. For example, if we have
type Phantom a = Int
and are unifying
Phantom Int ~ Phantom Char
it is *wrong* to unify Int and Char.
* The problem case immediately above can happen only with arguments
to the tycon. So we check for nullary tycons *before* expanding.
This is particularly helpful when checking (* ~ *), because * is
now a type synonym. See Note [Unifying type synonyms] in GHC.Core.Unify.
Note [Deferred unification]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
We may encounter a unification ty1 ~ ty2 that cannot be performed syntactically,
and yet its consistency is undetermined. Previously, there was no way to still
make it consistent. So a mismatch error was issued.
Now these unifications are deferred until constraint simplification, where type
family instances and given equations may (or may not) establish the consistency.
Deferred unifications are of the form
F ... ~ ...
or x ~ ...
where F is a type function and x is a type variable.
E.g.
id :: x ~ y => x -> y
id e = e
involves the unification x = y. It is deferred until we bring into account the
context x ~ y to establish that it holds.
If available, we defer original types (rather than those where closed type
synonyms have already been expanded via tcCoreView). This is, as usual, to
improve error messages.
************************************************************************
* *
uUnfilledVar and friends
* *
************************************************************************
@uunfilledVar@ is called when at least one of the types being unified is a
variable. It does {\em not} assume that the variable is a fixed point
of the substitution; rather, notice that @uVar@ (defined below) nips
back into @uTys@ if it turns out that the variable is already bound.
-}
----------
uUnfilledVar, uUnfilledVar1
:: UnifyEnv
-> SwapFlag
-> TcTyVar -- Tyvar 1: not necessarily a meta-tyvar
-- definitely not a /filled/ meta-tyvar
-> TcTauType -- Type 2
-> TcM CoercionN
-- "Unfilled" means that the variable is definitely not a filled-in meta tyvar
-- It might be a skolem, or untouchable, or meta
uUnfilledVar env swapped tv1 ty2
| Nominal <- u_role env
= do { ty2 <- liftZonkM $ zonkTcType ty2
-- Zonk to expose things to the occurs check, and so
-- that if ty2 looks like a type variable then it
-- /is/ a type variable
; uUnfilledVar1 env swapped tv1 ty2 }
| otherwise -- See Note [Do not unify representational equalities]
-- in GHC.Tc.Solver.Equality
= unSwap swapped (uType_defer env) (mkTyVarTy tv1) ty2
uUnfilledVar1 env -- Precondition: u_role==Nominal
swapped
tv1
ty2 -- ty2 is zonked
| Just tv2 <- getTyVar_maybe ty2
= go tv2
| otherwise
= uUnfilledVar2 env swapped tv1 ty2
where
-- 'go' handles the case where both are
-- tyvars so we might want to swap
-- E.g. maybe tv2 is a meta-tyvar and tv1 is not
go tv2 | tv1 == tv2 -- Same type variable => no-op
= return (mkNomReflCo (mkTyVarTy tv1))
| swapOverTyVars False tv1 tv2 -- Distinct type variables
-- Swap meta tyvar to the left if poss
= do { tv1 <- liftZonkM $ zonkTyCoVarKind tv1
-- We must zonk tv1's kind because that might
-- not have happened yet, and it's an invariant of
-- uUnfilledTyVar2 that ty2 is fully zonked
-- Omitting this caused #16902
; uUnfilledVar2 env (flipSwap swapped) tv2 (mkTyVarTy tv1) }
| otherwise
= uUnfilledVar2 env swapped tv1 ty2
----------
uUnfilledVar2 :: UnifyEnv -- Precondition: u_role==Nominal
-> SwapFlag
-> TcTyVar -- Tyvar 1: not necessarily a meta-tyvar
-- definitely not a /filled/ meta-tyvar
-> TcTauType -- Type 2, zonked
-> TcM CoercionN
uUnfilledVar2 env@(UE { u_defer = def_eq_ref }) swapped tv1 ty2
= do { cur_lvl <- getTcLevel
-- See Note [Unification preconditions], (UNTOUCHABLE) wrinkles
-- Here we don't know about given equalities; so we treat
-- /any/ level outside this one as untouchable. Hence cur_lvl.
; if simpleUnifyCheck UC_OnTheFly cur_lvl tv1 ty2 /= SUC_CanUnify
then not_ok_so_defer cur_lvl
else
do { def_eqs <- readTcRef def_eq_ref -- Capture current state of def_eqs
-- Attempt to unify kinds
-- When doing so, be careful to preserve orientation;
-- see Note [Kind Equality Orientation] in GHC.Tc.Solver.Equality
-- and wrinkle (W2) in Note [Fundeps with instances, and equality orientation]
-- in GHC.Tc.Solver.Dict
-- Failing to preserve orientation led to #25597.
; let kind_env = unSwap swapped (mkKindEnv env) ty1 ty2
; co_k <- unSwap swapped (uType kind_env) (tyVarKind tv1) (typeKind ty2)
; traceTc "uUnfilledVar2 ok" $
vcat [ ppr tv1 <+> dcolon <+> ppr (tyVarKind tv1)
, ppr ty2 <+> dcolon <+> ppr (typeKind ty2)
, ppr (isReflCo co_k), ppr co_k ]
; if isReflCo co_k
-- Only proceed if the kinds match
-- NB: tv1 should still be unfilled, despite the kind unification
-- because tv1 is not free in ty2' (or, hence, in its kind)
then do { liftZonkM $ writeMetaTyVar tv1 ty2
; case u_unified env of
Nothing -> return ()
Just uref -> updTcRef uref (tv1 :)
; return (mkNomReflCo ty2) } -- Unification is always Nominal
else -- The kinds don't match yet, so defer instead.
do { writeTcRef def_eq_ref def_eqs
-- Since we are discarding co_k, also discard any constraints
-- emitted by kind unification; they are just useless clutter.
-- Do this dicarding by simply restoring the previous state
-- of def_eqs; a bit imperative/yukky but works fine.
; defer }
}}
where
ty1 = mkTyVarTy tv1
defer = unSwap swapped (uType_defer env) ty1 ty2
not_ok_so_defer cur_lvl =
do { traceTc "uUnfilledVar2 not ok" $
vcat [ text "tv1:" <+> ppr tv1
, text "ty2:" <+> ppr ty2
, text "simple-unify-chk:" <+> ppr (simpleUnifyCheck UC_OnTheFly cur_lvl tv1 ty2)
]
-- Occurs check or an untouchable: just defer
-- NB: occurs check isn't necessarily fatal:
-- eg tv1 occurred in type family parameter
; defer }
swapOverTyVars :: Bool -> TcTyVar -> TcTyVar -> Bool
swapOverTyVars is_given tv1 tv2
-- See Note [Unification variables on the left]
| not is_given, pri1 == 0, pri2 > 0 = True
| not is_given, pri2 == 0, pri1 > 0 = False
-- Level comparison: see Note [TyVar/TyVar orientation]
| lvl1 `strictlyDeeperThan` lvl2 = False
| lvl2 `strictlyDeeperThan` lvl1 = True
-- Priority: see Note [TyVar/TyVar orientation]
| pri1 > pri2 = False
| pri2 > pri1 = True
-- Names: see Note [TyVar/TyVar orientation]
| isSystemName tv2_name, not (isSystemName tv1_name) = True
| otherwise = False
where
lvl1 = tcTyVarLevel tv1
lvl2 = tcTyVarLevel tv2
pri1 = lhsPriority tv1
pri2 = lhsPriority tv2
tv1_name = Var.varName tv1
tv2_name = Var.varName tv2
lhsPriority :: TcTyVar -> Int
-- Higher => more important to be on the LHS
-- => more likely to be eliminated
-- Only used when the levels are identical
-- See Note [TyVar/TyVar orientation]
lhsPriority tv
= assertPpr (isTyVar tv) (ppr tv) $
case tcTyVarDetails tv of
RuntimeUnk -> 0
SkolemTv {} -> 0
MetaTv { mtv_info = info, mtv_tclvl = lvl }
| QLInstVar <- lvl
-> 5 -- Eliminate instantiation variables first
| otherwise
-> case info of
CycleBreakerTv -> 0
TyVarTv -> 1
ConcreteTv {} -> 2
TauTv -> 3
RuntimeUnkTv -> 4
{- Note [Unification preconditions]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Question: given a homogeneous equality (alpha ~# ty), when is it OK to
unify alpha := ty?
(This note only applies to /homogeneous/ equalities, in which both
sides have the same kind.)
There are five reasons not to unify:
1. (SKOL-ESC) Skolem-escape
Consider the constraint
forall[2] a[2]. alpha[1] ~ Maybe a[2]
If we unify alpha := Maybe a, the skolem 'a' may escape its scope.
The level alpha[1] says that alpha may be used outside this constraint,
where 'a' is not in scope at all. So we must not unify.
Bottom line: when looking at a constraint alpha[n] := ty, do not unify
if any free variable of 'ty' has level deeper (greater) than n
2. (UNTOUCHABLE) Untouchable unification variables
Consider the constraint
forall[2] a[2]. b[1] ~ Int => alpha[1] ~ Int
There is no (SKOL-ESC) problem with unifying alpha := Int, but it might
not be the principal solution. Perhaps the "right" solution is alpha := b.
We simply can't tell. See "OutsideIn(X): modular type inference with local
assumptions", section 2.2. We say that alpha[1] is "untouchable" inside
this implication.
Bottom line: at ambient level 'l', when looking at a constraint
alpha[n] ~ ty, do not unify alpha := ty if there are any given equalities
between levels 'n' and 'l'.
Exactly what is a "given equality" for the purpose of (UNTOUCHABLE)?
Answer: see Note [Tracking Given equalities] in GHC.Tc.Solver.InertSet
3. (TYVAR-TV) Unifying TyVarTvs and CycleBreakerTvs
This precondition looks at the MetaInfo of the unification variable:
* TyVarTv: When considering alpha{tyv} ~ ty, if alpha{tyv} is a
TyVarTv it can only unify with a type variable, not with a
structured type. So if 'ty' is a structured type, such as (Maybe x),
don't unify.
* CycleBreakerTv: never unified, except by restoreTyVarCycles.
4. (CONCRETE) A ConcreteTv can only unify with a concrete type,
by definition.
That is, if we have `rr[conc] ~ F Int`, we can't unify
`rr` with `F Int`, so we hold off on unifying.
Note however that the equality might get rewritten; for instance
if we can rewrite `F Int` to a concrete type, say `FloatRep`,
then we will have `rr[conc] ~ FloatRep` and we can unify `rr ~ FloatRep`.
Note that we can still make progress on unification even if
we can't fully solve an equality, e.g.
alpha[conc] ~# TupleRep '[ beta[tau], F gamma[tau] ]
we can fill beta[tau] := beta[conc]. This is why we call
'makeTypeConcrete' in startSolvingByUnification.
5. (REWRITERS) the equality does not have any unsolved equalities in its rewriter
set. If those other equalities have not been solved, unifying this equality
will propagate strange-looking errors elswhere. That is the whole point of
rewriter sets. Suppose our equality is
[W] co1 {rew = {cok}} (alpha :: k) ~ (Int |> {cok})
where co :: Type ~ k is an unsolved wanted. Note that this equality
is homogeneous; both sides have kind k. We refrain from unifying
here, because of `cok` in its rewriter set. See
Note [Unify only if the rewriter set is empty] in GHC.Solver.Equality.
Needless to say, all there are wrinkles:
* (SKOL-ESC) Promotion. Given alpha[n] ~ ty, what if beta[k] is free
in 'ty', where beta is a unification variable, and k>n? 'beta'
stands for a monotype, and since it is part of a level-n type
(equal to alpha[n]), we must /promote/ beta to level n. Just make
up a fresh gamma[n], and unify beta[k] := gamma[n].
* (TYVAR-TV) Unification variables. Suppose alpha[tyv,n] is a level-n
TyVarTv (see Note [TyVarTv] in GHC.Tc.Types.TcMType)? Now
consider alpha[tyv,n] ~ Bool. We don't want to unify because that
would break the TyVarTv invariant.
What about alpha[tyv,n] ~ beta[tau,n], where beta is an ordinary
TauTv? Again, don't unify, because beta might later be unified
with, say Bool. (If levels permit, we reverse the orientation here;
see Note [TyVar/TyVar orientation].)
* (UNTOUCHABLE) Untouchability. When considering (alpha[n] ~ ty), how
do we know whether there are any given equalities between level n
and the ambient level? We answer in two ways:
* In the eager unifier, we only unify if l=n. If not, alpha may be
untouchable, and defer to the constraint solver. This check is
made in GHC.Tc.Utils.uUnifilledVar2, in the guard
isTouchableMetaTyVar.
* In the constraint solver, we track where Given equalities occur
and use that to guard unification in
GHC.Tc.Utils.Unify.touchabilityTest. More details in
Note [Tracking Given equalities] in GHC.Tc.Solver.InertSet
Historical note: in the olden days (pre 2021) the constraint solver
also used to unify only if l=n. Equalities were "floated" out of the
implication in a separate step, so that they would become touchable.
But the float/don't-float question turned out to be very delicate,
as you can see if you look at the long series of Notes associated with
GHC.Tc.Solver.floatEqualities, around Nov 2020. It's much easier
to unify in-place, with no floating.
Note [TyVar/TyVar orientation]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
See also Note [Fundeps with instances, and equality orientation]
where the kind equality orientation is important
Given (a ~ b), should we orient the equality as (a~b) or (b~a)?
This is a surprisingly tricky question!
The question is answered by swapOverTyVars, which is used
- in the eager unifier, in GHC.Tc.Utils.Unify.uUnfilledVar1
- in the constraint solver, in GHC.Tc.Solver.Equality.canEqCanLHS2
First note: only swap if you have to!
See Note [Avoid unnecessary swaps]
So we look for a positive reason to swap, using a three-step test:
* Level comparison. If 'a' has deeper level than 'b',
put 'a' on the left. See Note [Deeper level on the left]
* Priority. If the levels are the same, look at what kind of
type variable it is, using 'lhsPriority'.
Generally speaking we always try to put a MetaTv on the left in
preference to SkolemTv or RuntimeUnkTv, because the MetaTv may be
touchable and can be unified.
Tie-breaking rules for MetaTvs:
- CycleBreakerTv: This is essentially a stand-in for another type;
it's untouchable and should have the same priority as a skolem: 0.
- TyVarTv: These can unify only with another tyvar, but we can't unify
a TyVarTv with a TauTv, because then the TyVarTv could (transitively)
get a non-tyvar type. So give these a low priority: 1.
- ConcreteTv: These are like TauTv, except they can only unify with
a concrete type. So we want to be able to write to them, but not quite
as much as TauTvs: 2.
- TauTv: This is the common case; we want these on the left so that they
can be written to: 3.
- RuntimeUnkTv: These aren't really meta-variables used in type inference,
but just a convenience in the implementation of the GHCi debugger.
Eagerly write to these: 4. See Note [RuntimeUnkTv] in
GHC.Runtime.Heap.Inspect.
* Names. If the level and priority comparisons are all
equal, try to eliminate a TyVar with a System Name in
favour of ones with a Name derived from a user type signature
* Age. At one point in the past we tried to break any remaining
ties by eliminating the younger type variable, based on their
Uniques. See Note [Eliminate younger unification variables]
(which also explains why we don't do this any more)
Note [Unification variables on the left]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
For wanteds, but not givens, swap (skolem ~ meta-tv) regardless of
level, so that the unification variable is on the left.
* We /don't/ want this for Givens because if we ave
[G] a[2] ~ alpha[1]
[W] Bool ~ a[2]
we want to rewrite the wanted to Bool ~ alpha[1],
so we can float the constraint and solve it.
* But for Wanteds putting the unification variable on
the left means an easier job when floating, and when
reporting errors -- just fewer cases to consider.
In particular, we get better skolem-escape messages:
see #18114
Note [Deeper level on the left]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The most important thing is that we want to put tyvars with
the deepest level on the left. The reason to do so differs for
Wanteds and Givens, but either way, deepest wins! Simple.
* Wanteds. Putting the deepest variable on the left maximise the
chances that it's a touchable meta-tyvar which can be solved.
* Givens. Suppose we have something like
forall a[2]. b[1] ~ a[2] => beta[1] ~ a[2]
If we orient the Given a[2] on the left, we'll rewrite the Wanted to
(beta[1] ~ b[1]), and that can float out of the implication.
Otherwise it can't. By putting the deepest variable on the left
we maximise our changes of eliminating skolem capture.
See also GHC.Tc.Solver.InertSet Note [Let-bound skolems] for another reason
to orient with the deepest skolem on the left.
IMPORTANT NOTE: this test does a level-number comparison on
skolems, so it's important that skolems have (accurate) level
numbers.
See #15009 for an further analysis of why "deepest on the left"
is a good plan.
Note [Avoid unnecessary swaps]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
If we swap without actually improving matters, we can get an infinite loop.
Consider
work item: a ~ b
inert item: b ~ c
We canonicalise the work-item to (a ~ c). If we then swap it before
adding to the inert set, we'll add (c ~ a), and therefore kick out the
inert guy, so we get
new work item: b ~ c
inert item: c ~ a
And now the cycle just repeats
Historical Note [Eliminate younger unification variables]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Given a choice of unifying
alpha := beta or beta := alpha
we try, if possible, to eliminate the "younger" one, as determined
by `ltUnique`. Reason: the younger one is less likely to appear free in
an existing inert constraint, and hence we are less likely to be forced
into kicking out and rewriting inert constraints.
This is a performance optimisation only. It turns out to fix
#14723 all by itself, but clearly not reliably so!
It's simple to implement (see nicer_to_update_tv2 in swapOverTyVars).
But, to my surprise, it didn't seem to make any significant difference
to the compiler's performance, so I didn't take it any further. Still
it seemed too nice to discard altogether, so I'm leaving these
notes. SLPJ Jan 18.
Note [Prevent unification with type families]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We prevent unification with type families because of an uneasy compromise.
It's perfectly sound to unify with type families, and it even improves the
error messages in the testsuite. It also modestly improves performance, at
least in some cases. But it's disastrous for test case perf/compiler/T3064.
Here is the problem: Suppose we have (F ty) where we also have [G] F ty ~ a.
What do we do? Do we reduce F? Or do we use the given? Hard to know what's
best. GHC reduces. This is a disaster for T3064, where the type's size
spirals out of control during reduction. If we prevent
unification with type families, then the solver happens to use the equality
before expanding the type family.
It would be lovely in the future to revisit this problem and remove this
extra, unnecessary check. But we retain it for now as it seems to work
better in practice.
Revisited in Nov '20, along with removing flattening variables. Problem
is still present, and the solution is still the same.
Note [Non-TcTyVars in GHC.Tc.Utils.Unify]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Because the same code is now shared between unifying types and unifying
kinds, we sometimes will see proper TyVars floating around the unifier.
Example (from test case polykinds/PolyKinds12):
type family Apply (f :: k1 -> k2) (x :: k1) :: k2
type instance Apply g y = g y
When checking the instance declaration, we first *kind-check* the LHS
and RHS, discovering that the instance really should be
type instance Apply k3 k4 (g :: k3 -> k4) (y :: k3) = g y
During this kind-checking, all the tyvars will be TcTyVars. Then, however,
as a second pass, we desugar the RHS (which is done in functions prefixed
with "tc" in GHC.Tc.TyCl"). By this time, all the kind-vars are proper
TyVars, not TcTyVars, get some kind unification must happen.
Thus, we always check if a TyVar is a TcTyVar before asking if it's a
meta-tyvar.
This used to not be necessary for type-checking (that is, before * :: *)
because expressions get desugared via an algorithm separate from
type-checking (with wrappers, etc.). Types get desugared very differently,
causing this wibble in behavior seen here.
-}
-- | Breaks apart a function kind into its pieces.
matchExpectedFunKind
:: TypedThing -- ^ type, only for errors
-> Arity -- ^ n: number of desired arrows
-> TcKind -- ^ fun_kind
-> TcM Coercion -- ^ co :: fun_kind ~ (arg1 -> ... -> argn -> res)
matchExpectedFunKind hs_ty n k = go n k
where
go 0 k = return (mkNomReflCo k)
go n k | Just k' <- coreView k = go n k'
go n k@(TyVarTy kvar)
| isMetaTyVar kvar
= do { maybe_kind <- readMetaTyVar kvar
; case maybe_kind of
Indirect fun_kind -> go n fun_kind
Flexi -> defer n k }
go n (FunTy { ft_af = af, ft_mult = w, ft_arg = arg, ft_res = res })
| isVisibleFunArg af
= do { co <- go (n-1) res
; return (mkNakedFunCo Nominal af (mkNomReflCo w) (mkNomReflCo arg) co) }
go n other
= defer n other
defer n k
= do { arg_kinds <- newMetaKindVars n
; res_kind <- newMetaKindVar
; let new_fun = mkVisFunTysMany arg_kinds res_kind
origin = TypeEqOrigin { uo_actual = k
, uo_expected = new_fun
, uo_thing = Just hs_ty
, uo_visible = True
}
; unifyTypeAndEmit KindLevel origin k new_fun }
{- *********************************************************************
* *
Checking alpha ~ ty
for the on-the-fly unifier
* *
********************************************************************* -}
data UnifyCheckCaller
= UC_OnTheFly -- Called from the on-the-fly unifier
| UC_QuickLook -- Called from Quick Look
| UC_Solver -- Called from constraint solver
-- | The result type of 'simpleUnifyCheck'.
data SimpleUnifyResult
-- | Definitely cannot unify (untouchable variable or incompatible top-shape)
= SUC_CannotUnify
-- | The variable is touchable and the top-shape test passed, but
-- it may or may not be OK to unify
| SUC_NotSure
-- | Definitely OK to unify
| SUC_CanUnify
deriving stock (Eq, Ord, Show)
instance Semigroup SimpleUnifyResult where
no@SUC_CannotUnify <> _ = no
SUC_CanUnify <> r = r
_ <> no@SUC_CannotUnify = no
r <> SUC_CanUnify = r
ns@SUC_NotSure <> SUC_NotSure = ns
instance Outputable SimpleUnifyResult where
ppr = \case
SUC_CannotUnify -> text "SUC_CannotUnify"
SUC_NotSure -> text "SUC_NotSure"
SUC_CanUnify -> text "SUC_CanUnify"
simpleUnifyCheck :: UnifyCheckCaller -> TcLevel -> TcTyVar -> TcType -> SimpleUnifyResult
-- ^ A fast check for unification. May return "not sure", in which case
-- unification might still be OK, but it'll take more work to do
-- (use the full 'checkTypeEq').
--
-- * Rejects if lhs_tv occurs in rhs_ty (occurs check)
-- * Rejects foralls unless
-- lhs_tv is RuntimeUnk (used by GHCi debugger)
-- or is a QL instantiation variable
-- * Rejects a non-concrete type if lhs_tv is concrete
-- * Rejects type families unless fam_ok=True
-- * Does a level-check for type variables, to avoid skolem escape
--
-- This function is pretty heavily used, so it's optimised not to allocate
simpleUnifyCheck caller given_eq_lvl lhs_tv rhs
| not $ touchabilityTest given_eq_lvl lhs_tv
= SUC_CannotUnify
| not $ checkTopShape lhs_info rhs
= SUC_CannotUnify
| rhs_is_ok rhs
= SUC_CanUnify
| otherwise
= SUC_NotSure
where
lhs_info = metaTyVarInfo lhs_tv
!(occ_in_ty, occ_in_co) = mkOccFolders (tyVarName lhs_tv)
lhs_tv_lvl = tcTyVarLevel lhs_tv
lhs_tv_is_concrete = isConcreteTyVar lhs_tv
forall_ok = case caller of
UC_QuickLook -> isQLInstTyVar lhs_tv
_ -> isRuntimeUnkTyVar lhs_tv
-- This fam_ok thing relates to a very specific perf problem
-- See Note [Prevent unification with type families]
-- A couple of QuickLook regression tests rely on unifying with type
-- families, so we let it through there (not very principled, but let's
-- see if it bites us)
fam_ok = case caller of
UC_Solver -> True
UC_QuickLook -> True
UC_OnTheFly -> False
rhs_is_ok (TyVarTy tv)
| lhs_tv == tv = False
| tcTyVarLevel tv `strictlyDeeperThan` lhs_tv_lvl = False
| lhs_tv_is_concrete, not (isConcreteTyVar tv) = False
| occ_in_ty $! (tyVarKind tv) = False
| otherwise = True
rhs_is_ok (FunTy {ft_af = af, ft_mult = w, ft_arg = a, ft_res = r})
| not forall_ok, isInvisibleFunArg af = False
| otherwise = rhs_is_ok w && rhs_is_ok a && rhs_is_ok r
rhs_is_ok (TyConApp tc tys)
| lhs_tv_is_concrete, not (isConcreteTyCon tc) = False
| not forall_ok, not (isTauTyCon tc) = False
| not fam_ok, not (isFamFreeTyCon tc) = False
| otherwise = all rhs_is_ok tys
rhs_is_ok (ForAllTy (Bndr tv _) ty)
| forall_ok = rhs_is_ok (tyVarKind tv) && (tv == lhs_tv || rhs_is_ok ty)
| otherwise = False
rhs_is_ok (AppTy t1 t2) = rhs_is_ok t1 && rhs_is_ok t2
rhs_is_ok (CastTy ty co) = not (occ_in_co co) && rhs_is_ok ty
rhs_is_ok (CoercionTy co) = not (occ_in_co co)
rhs_is_ok (LitTy {}) = True
mkOccFolders :: Name -> (TcType -> Bool, TcCoercion -> Bool)
-- These functions return True
-- * if lhs_tv occurs (incl deeply, in the kind of variable)
-- * if there is a coercion hole
-- No expansion of type synonyms
mkOccFolders lhs_tv = (getAny . check_ty, getAny . check_co)
where
!(check_ty, _, check_co, _) = foldTyCo occ_folder emptyVarSet
occ_folder = TyCoFolder { tcf_view = noView -- Don't expand synonyms
, tcf_tyvar = do_tcv, tcf_covar = do_tcv
, tcf_hole = do_hole
, tcf_tycobinder = do_bndr }
do_tcv is v = Any (not (v `elemVarSet` is) && tyVarName v == lhs_tv)
`mappend` check_ty (varType v)
do_bndr is tcv _faf = extendVarSet is tcv
do_hole _is _hole = DM.Any True -- Reject coercion holes
{- *********************************************************************
* *
Equality invariant checking
* *
********************************************************************* -}
{- Note [Checking for foralls]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We never want to unify
alpha ~ (forall a. a->a) -> Int
So we look for foralls hidden inside the type, and it's convenient
to do that at the same time as the occurs check (which looks for
occurrences of alpha).
However, it's not just a question of looking for foralls /anywhere/!
Consider
(alpha :: forall k. k->*) ~ (beta :: forall k. k->*)
This is legal; e.g. dependent/should_compile/T11635.
We don't want to reject it because of the forall in beta's kind, but
(see Note [Occurrence checking: look inside kinds] in GHC.Core.Type)
we do need to look in beta's kind. So we carry a flag saying if a
'forall' is OK, and switch the flag on when stepping inside a kind.
Why is it OK? Why does it not count as impredicative polymorphism?
The reason foralls are bad is because we reply on "seeing" foralls
when doing implicit instantiation. But the forall inside the kind is
fine. We'll generate a kind equality constraint
(forall k. k->*) ~ (forall k. k->*)
to check that the kinds of lhs and rhs are compatible. If alpha's
kind had instead been
(alpha :: kappa)
then this kind equality would rightly complain about unifying kappa
with (forall k. k->*)
Note [Forgetful synonyms in checkTyConApp]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider
type S a b = b -- Forgets 'a'
[W] alpha[2] ~ Maybe (S beta[4] gamma[2])
We don't want to promote beta to level 2; rather, we should
expand the synonym. (Currently, in checkTypeEqRhs promotion
is irrevocable, by side effect.)
To avoid this risk we eagerly expand forgetful synonyms.
This also means we won't get an occurs check in
a ~ Maybe (S a b)
The annoyance is that we might expand the synonym unnecessarily,
something we generally try to avoid. But for now, this seems
simple.
In a forgetful case like a ~ Maybe (S a b), `checkTyEqRhs` returns
a Reduction that looks
Reduction { reductionCoercion = Refl
, reductionReducedType = Maybe b }
We must jolly well use that reductionReduced type, even though the
reductionCoercion is Refl. See `canEqCanLHSFinish_no_unification`.
-}
data PuResult a b
-- | Pure unifier failure.
--
-- Invariant: the CheckTyEqResult is not 'cteOK'; that it, it specifies a problem.
= PuFail CheckTyEqResult
-- | Pure unifier success.
| PuOK (Bag a) b
deriving stock (Functor, Foldable, Traversable)
instance Applicative (PuResult a) where
pure x = PuOK emptyBag x
PuFail p1 <*> PuFail p2 = PuFail (p1 S.<> p2)
PuFail p1 <*> PuOK {} = PuFail p1
PuOK {} <*> PuFail p2 = PuFail p2
PuOK c1 f <*> PuOK c2 x = PuOK (c1 `unionBags` c2) (f x)
instance (Outputable a, Outputable b) => Outputable (PuResult a b) where
ppr (PuFail prob) = text "PuFail" <+> (ppr prob)
ppr (PuOK cts x) = text "PuOK" <> braces
(vcat [ text "redn:" <+> ppr x
, text "cts:" <+> ppr cts ])
okCheckRefl :: TcType -> PuResult a Reduction
okCheckRefl ty = PuOK emptyBag (mkReflRedn Nominal ty)
mapCheck :: Monad m
=> (x -> m (PuResult a Reduction))
-> [x]
-> m (PuResult a Reductions)
mapCheck f xs
= do { (ress :: [PuResult a Reduction]) <- mapM f xs
; return (unzipRedns <$> sequenceA ress) }
-- sequenceA :: [PuResult a Reduction] -> PuResult a [Reduction]
-- unzipRedns :: [Reduction] -> Reductions
{-# INLINEABLE mapCheck #-}
-----------------------------
-- | Options describing how to deal with a type equality
-- in the eager unifier. See 'checkTyEqRhs'
data TyEqFlags m a
= -- | TFTyFam: LHS is a type family application
-- Invariant: we are not unifying; see `notUnifying_TEFTask`
TEFTyFam
{ tefTyFam_occursCheck :: CheckTyEqProblem
-- ^ The 'CheckTyEqProblem' to report for occurs-check failures
-- (soluble or insoluble)
, tefTyFam_tyCon :: TyCon
, tefTyFam_args :: [Type]
, tef_fam_app :: TyEqFamApp m a
-- ^ How to deal with type family applications
}
-- | TEFTyVar: LHS is a 'TyVar'.
| TEFTyVar
-- NB: this constructor does not actually store a 'TyVar', in order to
-- support being called from 'makeTypeConcrete' (which works as if we
-- created a fresh 'ConcreteTv' metavariable; see (3) in Note [TyEqFlags])
{ tefTyVar_occursCheck :: OccursCheck
-- ^ Occurs check
, tefTyVar_levelCheck :: LevelCheck m
-- ^ Level check
, tefTyVar_concreteCheck :: ConcreteCheck m
-- ^ Concreteness check
, tef_fam_app :: TyEqFamApp m a
-- ^ How to deal with type family applications
}
-- | What to do when encountering a type-family application while processing
-- a type equality in the pure unifier.
--
-- See Note [Family applications in canonical constraints]
data TyEqFamApp m a where
-- | Just recurse
TEFA_Recurse :: TyEqFamApp m a
-- | Recurse, but replace with cycle breaker if that fails,
-- using the specified 'FamAppBreaker'
TEFA_Break :: FamAppBreaker a -> TyEqFamApp TcM a
-- | A defunctionalisation of family application breaker functions.
-- Interpreter: `famAppBreaker`
data FamAppBreaker a where
BreakGiven :: FamAppBreaker (TcTyVar, TcType)
BreakWanted :: CtEvidence -> TcTyVar -> FamAppBreaker Ct
data OccursCheck where
-- | No occurs check.
OC_None :: OccursCheck
-- | Do an occurs check between the LHS tyvar and the RHS.
OC_Check ::
{ occurs_tv_name :: Name
-- ^ The 'Name' to perform an occurs-check on
, occurs_problem :: CheckTyEqProblem
-- ^ Which 'CheckTyEqProblem' to report for occurs-check failures
-- (soluble or insoluble)
} -> OccursCheck
-- | What level check to perform, in a call to the pure unifier?
-- A defunctionalisation of the possible level-check functions
-- Interpreter: `tyVarLevelCheck`
data LevelCheck m where
-- | No level check.
LC_None :: LevelCheck m
-- | Do a level check between the LHS tyvar and the occurrence tyvar.
--
-- Fail if the level check fails.
--
-- True <=> lenient check, e.g. the levels have a chance of working out
-- after promotion.
LC_Check ::
{ lc_lvlc :: TcLevel
, lc_lenient :: Bool
} -> LevelCheck m
-- | Do a level check between the LHS tyvar and the occurrence tyvar.
--
-- If the level check fails, and the occurrence is a unification
-- variable, promote it.
--
-- - False <=> don't promote under type families (the common case)
-- - True <=> promote even under type families
-- (see Note [Defaulting equalities] in GHC.Tc.Solver)
LC_Promote
:: { lc_lvlp :: TcLevel
, lc_deep :: Bool
} -> LevelCheck TcM
data ConcreteCheck m where
-- | No concreteness check.
CC_None :: ConcreteCheck m
-- | Simple check for concreteness, leniently returning 'OK'
-- if concreteness could be achieved after promotion.
CC_Check :: ConcreteCheck m
-- | Proper concreteness check: promote non-concrete metavariables,
-- failing if there are any problems.
CC_Promote :: ConcreteTvOrigin -> ConcreteCheck TcM
{- Note [TyEqFlags]
~~~~~~~~~~~~~~~~~~~
When we call the eager unifier, e.g. through 'checkTyEqRhs', we specify what
kind of checks the unifier performs via the 'TyEqFlags' argument. In particular,
when the LHS type in a unification is a type variable, we might want to perform
different checks; this is achieved using the 'TEFTyVar' constructor to 'TyEqFlags':
1. `notUnifying_TEFTask`
LHS is a skolem tyvar, or an untouchable meta-tyvar.
We are not unifying; we only want to perform occurs-checks.
TEFTyVar
{ tefTyVar_occursCheck = OC_Check ...
, tefTyVar_levelCheck = LC_None
, tefTyVar_concreteCheck = CC_None
, tef_fam_app = TEFA_Recurse
}
2a. `unifyingLHSMetaTyVar_TEFTask`
We are unifying; we want to perform an occurs check, a level check,
and a concreteness check (when the meta-tyvar is a ConcreteTv).
TEFTyVar
{ tefTyVar_occursCheck = OC_Check ...
, tefTyVar_levelCheck = LC_Promote ...
, tefTyVar_concreteCheck = CC_Promote or CC_None
-- depending on whether or not the lhs tv is concrete
, tef_fam_app = TEFA_Recurse
}
2b. `defaulting_TEFTask`
We are in the top-level defaulting code, considering unifying a
touchable meta-tyvar with a type.
TEFTyVar
{ tefTyVar_occursCheck = OC_None
, tefTyVar_levelCheck = LC_None
, tefTyVar_concreteCheck = CC_Promote or CC_None
-- depending on whether or not the lhs tv is concrete
, tef_fam_app = TEFA_Recurse
}
3. `makeTypeConcrete`
LHS is a fresh ConcreteTv meta-tyvar (see call to 'checkTyEqRhs' in
`makeTypeConcrete`). We are unifying; we only want to perform
a concreteness check.
TEFTyVar
{ tefTyVar_occursCheck = OC_None
, tefTyVar_levelCheck = LC_None
, tefTyVar_concreteCheck = CC_Promote conc_orig
, tef_fam_app = TEFA_Recurse
}
4. `pureTyEqFlags_LHSMetaTyVar`
We want to perform a non-monadic check, i.e. we want to know whether we are able
to unify 'lhs_tv' with 'rhs_ty', but don't want to actually perform
a side-effecting unification. This is used in 'mightEqualLater'.
TEFTyVar
{ tefTyVar_occursCheck = OC_Check ...
, tefTyVar_levelCheck = LC_Check ...
, tefTyVar_concreteCheck = CC_Check or CC_None
-- depending on whether or not the lhs tv is concrete
, tef_fam_app = TEFA_Recurse
}
Here 'LC_Check True' and 'ConcreteSimpleCheck' make 'checkTyEqRhs' perform
conservative pure checks, without actually carrying out any promotion.
-}
-- | Create a "not unifying" 'TyEqFlags' from a 'CanEqLHS'.
--
-- See use-case (1) in Note [TyEqFlags].
notUnifying_TEFTask :: CheckTyEqProblem -> CanEqLHS -> TyEqFlags m a
notUnifying_TEFTask occ_prob = \case
TyFamLHS tc tys ->
TEFTyFam
{ tefTyFam_tyCon = tc
, tefTyFam_args = tys
, tefTyFam_occursCheck = occ_prob
, tef_fam_app = TEFA_Recurse
}
TyVarLHS tv ->
TEFTyVar
{ tefTyVar_occursCheck = OC_Check (tyVarName tv) occ_prob
, tefTyVar_levelCheck = LC_None
, tefTyVar_concreteCheck = CC_None
, tef_fam_app = TEFA_Recurse
}
-- We need an occurs-check here, but no level check.
-- See Note [Promotion and level-checking] wrinkle (W1)
-- TEFA_Recurse: see Note [Don't cycle-break Wanteds when not unifying]
-- | Create "unifying" 'TyEqFlags' from a 'TyVarLHS'.
--
-- Invariant: the argument 'TcTyVar' is a 'MetaTv'.
unifyingLHSMetaTyVar_TEFTask :: CtEvidence -> TcTyVar -> TyEqFlags TcM Ct
unifyingLHSMetaTyVar_TEFTask ev lhs_tv =
TEFTyVar
{ tefTyVar_occursCheck = OC_Check (tyVarName lhs_tv) cteInsolubleOccurs
, tefTyVar_levelCheck = LC_Promote { lc_lvlp = tcTyVarLevel lhs_tv
, lc_deep = False }
, tefTyVar_concreteCheck = mkConcreteCheck lhs_tv
, tef_fam_app = mkTEFA_Break ev NomEq (BreakWanted ev lhs_tv)
}
defaulting_TEFTask :: TcTyVar -> TyEqFlags TcM a
-- Used during top-level defauting in GHC.Tc.Solver.Default.defaultEquality
-- Invariant: the argument 'TcTyVar' is a 'MetaTv'.
defaulting_TEFTask lhs_tv =
TEFTyVar
{ tefTyVar_occursCheck = OC_Check (tyVarName lhs_tv) cteInsolubleOccurs
, tefTyVar_levelCheck = LC_Promote { lc_lvlp = tcTyVarLevel lhs_tv
, lc_deep = True }
-- LC_Promote: promote deeper unification variables (DE4)
-- lc_deep = True: ...including under type families (DE5)
, tefTyVar_concreteCheck = mkConcreteCheck lhs_tv
, tef_fam_app = TEFA_Recurse
}
mkConcreteCheck :: TcTyVar -> ConcreteCheck TcM
mkConcreteCheck lhs_tv = case isConcreteTyVar_maybe lhs_tv of
Nothing -> CC_None
Just conc_orig -> CC_Promote conc_orig
-- | 'TyEqFlags' for a pure call of 'checkTyEqRhs'.
--
-- A call to 'checkTyEqRhs' with these 'TyEqFlags' will perform an optimistic
-- check: it will return 'PuFail' if there is definitely a problem, but it
-- might return 'PuOK' if it isn't entirely sure.
pureTyEqFlags_LHSMetaTyVar :: TyVar -> TyEqFlags Identity ()
pureTyEqFlags_LHSMetaTyVar lhs_tv =
TEFTyVar
{ tefTyVar_occursCheck = OC_Check (tyVarName lhs_tv) cteInsolubleOccurs
, tefTyVar_levelCheck = LC_Check { lc_lvlc = tcTyVarLevel lhs_tv
, lc_lenient = True }
, tefTyVar_concreteCheck = case isConcreteTyVar_maybe lhs_tv of
Nothing -> CC_None
Just {} -> CC_Check -- No promotion
, tef_fam_app = TEFA_Recurse }
mkTEFA_Break :: CtEvidence -> EqRel -> FamAppBreaker a -> TyEqFamApp TcM a
mkTEFA_Break ev eq_rel breaker
| NomEq <- eq_rel
, not cycle_breaker_origin
= TEFA_Break breaker
| otherwise
= TEFA_Recurse
where
-- cycle_breaker_origin: see Detail (7) of Note [Type equality cycles]
-- in GHC.Tc.Solver.Equality
cycle_breaker_origin = case ctLocOrigin (ctEvLoc ev) of
CycleBreakerOrigin {} -> True
_ -> False
-- | Do we want to perform a concreteness check in 'checkTyEqRhs'?
tefConcrete :: TyEqFlags m a -> Bool
tefConcrete (TEFTyFam {}) = False
tefConcrete (TEFTyVar { tefTyVar_concreteCheck = conc }) =
wantConcreteCheck conc
wantConcreteCheck :: ConcreteCheck m -> Bool
wantConcreteCheck = \case
CC_None -> False
CC_Check -> True
CC_Promote {} -> True
-- | Given a family-application @ty@, return a @'Reduction' :: ty ~ cbv@
-- where @cbv@ is a fresh loop-breaker tyvar (for Given), or
-- just a fresh 'TauTv' (for Wanted)
famAppBreaker :: FamAppBreaker a -> TcType -> TcM (PuResult a Reduction)
famAppBreaker BreakGiven fam_app
= do { new_tv <- TcM.newCycleBreakerTyVar (typeKind fam_app)
; return (PuOK (unitBag (new_tv, fam_app))
(mkReflRedn Nominal (mkTyVarTy new_tv))) }
-- Why reflexive? See Detail (4) of Note [Type equality cycles]
-- in GHC.Tc.Solver.Equality
famAppBreaker (BreakWanted ev lhs_tv) fam_app
-- Occurs check or skolem escape; so flatten
= do { reason <- checkPromoteFreeVars cteInsolubleOccurs
(tyVarName lhs_tv) lhs_tv_lvl
(tyCoVarsOfType fam_app_kind)
; if not (cterHasNoProblem reason) -- Failed to promote free vars
then return $ PuFail reason
else
do { new_tv_ty <-
case lhs_tv_info of
ConcreteTv conc_info ->
-- Make a concrete tyvar if lhs_tv is concrete
-- e.g. alpha[2,conc] ~ Maybe (F beta[4])
-- We want to flatten to
-- alpha[2,conc] ~ Maybe gamma[2,conc]
-- gamma[2,conc] ~ F beta[4]
TcM.newConcreteTyVarTyAtLevel conc_info lhs_tv_lvl fam_app_kind
_ -> TcM.newMetaTyVarTyAtLevel lhs_tv_lvl fam_app_kind
; let pty = mkNomEqPred fam_app new_tv_ty
; hole <- TcM.newCoercionHole pty
; let new_ev = WantedCt { ctev_pred = pty
, ctev_dest = HoleDest hole
, ctev_loc = cb_loc
, ctev_rewriters = ctEvRewriters ev }
; return (PuOK (singleCt (mkNonCanonical $ CtWanted new_ev))
(mkReduction (HoleCo hole) new_tv_ty)) } }
where
fam_app_kind = typeKind fam_app
(lhs_tv_info, lhs_tv_lvl) =
case tcTyVarDetails lhs_tv of
MetaTv { mtv_info = info, mtv_tclvl = lvl } -> (info,lvl)
-- lhs_tv should be a meta-tyvar
_ -> pprPanic "famAppBreaker BreakWanted: lhs_tv is not a meta-tyvar"
(ppr lhs_tv)
cb_loc = updateCtLocOrigin (ctEvLoc ev) CycleBreakerOrigin
-- CycleBreakerOrigin: see Detail (7) of Note [Type equality cycles]
instance Outputable (TyEqFlags m a) where
ppr = \case
TEFTyFam occ tc tys fam_app ->
text "TEFTyFam" <> braces (
hcat [ ppr occ
, ppr (mkTyConApp tc tys)
, ppr fam_app ] )
TEFTyVar occ lc conc fam_app ->
text "TEFTyVar" <> braces (hcat (punctuate comma fields))
where
fields = [ text "OccursCheck:" <+> ppr occ ]
++
[ text "LevelCheck:" <+> ppr lc ]
++
[ text "ConcreteCheck:" <+> ppr conc ]
++
[ text "FamApp:" <+> ppr fam_app ]
instance Outputable (TyEqFamApp m a) where
ppr TEFA_Recurse = text "TEFA_Recurse"
ppr (TEFA_Break breaker) = text "TEFA_BreakGiven" <+> ppr breaker
instance Outputable (FamAppBreaker a) where
ppr BreakGiven = text "BreakGiven"
ppr (BreakWanted ev tv) = parens $ text "BreakWanted" <+> ppr ev <+> ppr tv
instance Outputable OccursCheck where
ppr OC_None = text "OC_None"
ppr (OC_Check { occurs_tv_name = nm }) = text "OC_Check" <+> ppr nm
instance Outputable (ConcreteCheck m) where
ppr CC_None = text "CC_None"
ppr CC_Check = text "CC_Check"
ppr (CC_Promote {}) = text "CC_Promote"
instance Outputable (LevelCheck m) where
ppr (LC_None) = text "LC_None"
ppr (LC_Check lvl lenient) = text "LC_Check" <+> ppr lvl <+> ppWhen lenient (text "(lenient)")
ppr (LC_Promote lvl deep) = text "LC_Promote" <+> ppr lvl <+> ppWhen deep (text "(deep)")
-- | Adjust the 'TyEqFlags' when going under a type family:
--
-- 1. Only the outer family application gets the loop-breaker treatment
-- 2. Weaken level checks for tyvar promotion. For example, in @[W] alpha[2] ~ Maybe (F beta[3])@,
-- do not promote @beta[3]@, instead promote @(F beta[3])@.
-- 3. Occurs checks become potentially soluble (after additional type family
-- reductions).
famAppArgFlags :: TyEqFlags m a -> TyEqFlags m a
famAppArgFlags flags = case flags of
TEFTyFam {} ->
flags
{ tef_fam_app = TEFA_Recurse -- (1)
, tefTyFam_occursCheck = cteSolubleOccurs -- (3)
}
TEFTyVar { tefTyVar_occursCheck = occ, tefTyVar_levelCheck = mb_lc } ->
flags
{ tef_fam_app = TEFA_Recurse -- (1)
, tefTyVar_levelCheck = zap_lc mb_lc -- (2)
, tefTyVar_occursCheck = soluble_occ occ -- (3)
}
where
soluble_occ = \case
OC_Check tv _ -> OC_Check tv cteSolubleOccurs
OC_None -> OC_None
zap_lc = \case
LC_Promote { lc_lvlp = lvl, lc_deep = deeply }
| not deeply
-> LC_Check { lc_lvlc = lvl, lc_lenient = False }
lc -> lc
{- Note [checkTyEqRhs]
~~~~~~~~~~~~~~~~~~~~~~
The key function `checkTyEqRhs ty_eq_flags rhs` is called on the
RHS of a type equality
lhs ~ rhs
and checks to see if `rhs` satisfies, or can be made to satisfy,
invariants described by `ty_eq_flags`. It can succeded or fail; in
the latter case it returns a `CheckTyEqResult` that describes why it
failed.
When `lhs` is a touchable type variable, so unification might happen, then
`checkTyEqRhs` enforces the unification preconditions of Note [Unification preconditions].
Notably, it can check for things like:
* Insoluble occurs check
e.g. alpha[tau] ~ [alpha]
or F Int ~ [F Int]
* Potentially-soluble occurs check
e.g. alpha[tau] ~ [F alpha beta]
* Impredicativity error:
e.g. alpha[tau] ~ (forall a. a->a)
* Skolem escape
e.g alpha[1] ~ (b[sk:2], Int)
* Concreteness error
e.g. alpha[conc] ~ r[sk]
Its specific behaviour is governed by the `TyEqFlags` that are passed
to it; see Note [TyEqFlags].
Note, however, that `checkTyEqRhs` specifically does /not/ check for:
* Touchability of the LHS (in the case of a unification variable)
* Shape of the LHS (e.g. we can't unify Int with a TyVarTv)
These things are checked by `simpleUnifyCheck`.
-}
-- | Perform the checks specified by the 'TyEqFlags' on the RHS, in order to
-- enforce the unification preconditions of Note [Unification preconditions].
--
-- See Note [checkTyEqRhs].
checkTyEqRhs :: forall m a
. Monad m
=> TyEqFlags m a
-> TcType -- Already zonked
-> m (PuResult a Reduction)
checkTyEqRhs flags rhs
-- Crucial special case for a top-level equality of the form 'alpha ~ F tys'.
-- We don't want to flatten that (F tys), as this gets us right back to where
-- we started!
--
-- See also Note [Special case for top-level of Given equality]
| Just (TyFamLHS tc tys) <- canTyFamEqLHS_maybe rhs
, not $ tefConcrete flags
= recurseIntoFamTyConApp flags tc tys
| otherwise
= check_ty_eq_rhs flags rhs
{- Note [Special case for top-level of Given equality]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We take care when examining
[G] F ty ~ G (...(F ty)...)
where both sides are TyFamLHSs. We don't want to flatten that RHS to
[G] F ty ~ cbv
[G] G (...(F ty)...) ~ cbv
Instead we'd like to say "occurs-check" and swap LHS and RHS, which yields a
canonical constraint
[G] G (...(F ty)...) ~ F ty
That tends to rewrite a big type to smaller one. This happens in T15703,
where we had:
[G] Pure g ~ From1 (To1 (Pure g))
Making a loop breaker and rewriting left to right just makes much bigger
types than swapping it over.
(We might hope to have swapped it over before getting to checkTypeEq,
but better safe than sorry.)
NB: We never see a TyVarLHS here, such as
[G] a ~ F tys here
because we'd have swapped it to
[G] F tys ~ a
in canEqCanLHS2, before getting to checkTypeEq.
-}
check_ty_eq_rhs :: forall m a
. Monad m
=> TyEqFlags m a
-> TcType -- Already zonked
-> m (PuResult a Reduction)
check_ty_eq_rhs flags ty
= case ty of
LitTy {} -> return $ okCheckRefl ty
TyConApp tc tys -> checkTyConApp flags ty tc tys
TyVarTy tv -> checkTyVar flags tv
-- Don't worry about foralls inside the kind; see Note [Checking for foralls]
-- Nor can we expand synonyms; see Note [Occurrence checking: look inside kinds]
-- in GHC.Core.FVs
FunTy {ft_af = af, ft_mult = w, ft_arg = a, ft_res = r}
| isInvisibleFunArg af -- e.g. Num a => blah
-> return $ PuFail impredicativeProblem -- Not allowed (TyEq:F)
| otherwise
-> do { w_res <- check_ty_eq_rhs flags w
; a_res <- check_ty_eq_rhs flags a
; r_res <- check_ty_eq_rhs flags r
; return (mkFunRedn Nominal af <$> w_res <*> a_res <*> r_res) }
AppTy fun arg -> do { fun_res <- check_ty_eq_rhs flags fun
; arg_res <- check_ty_eq_rhs flags arg
; return (mkAppRedn <$> fun_res <*> arg_res) }
CastTy ty co -> do { ty_res <- check_ty_eq_rhs flags ty
; co_res <- checkCo flags co
; return (mkCastRedn1 Nominal ty <$> co_res <*> ty_res) }
CoercionTy co -> do { co_res <- checkCo flags co
; return (mkReflCoRedn Nominal <$> co_res) }
ForAllTy {} -> return $ PuFail impredicativeProblem -- Not allowed (TyEq:F)
{-# INLINEABLE check_ty_eq_rhs #-}
-------------------
checkCo :: Monad m => TyEqFlags m a -> Coercion -> m (PuResult a Coercion)
-- See Note [checkCo]
checkCo flags co =
case flags of
TEFTyFam {} ->
-- NB: 'TEFTyFam' case means we are not unifying.
return (pure co)
TEFTyVar
{ tefTyVar_concreteCheck = conc
, tefTyVar_levelCheck = lc
, tefTyVar_occursCheck = occ
}
-- Coercions cannot appear in concrete types.
--
-- See Note [Concrete types] in GHC.Tc.Utils.Concrete.
| case conc of { CC_None -> False; _ -> True }
-> return $ PuFail (cteProblem cteConcrete)
-- Occurs check (can promote)
| OC_Check lhs_tv occ_prob <- occ
, LC_Promote { lc_lvlp = lhs_tv_lvl } <- lc
-> do { reason <- checkPromoteFreeVars occ_prob lhs_tv lhs_tv_lvl (tyCoVarsOfCo co)
; return $
if cterHasNoProblem reason
then pure co
else PuFail reason }
-- Occurs check (no promotion)
| OC_Check lhs_tv occ_prob <- occ
, nameUnique lhs_tv `elemVarSetByKey` tyCoVarsOfCo co
-> return $ PuFail (cteProblem occ_prob)
| otherwise
-> return (pure co)
{- Note [checkCo]
~~~~~~~~~~~~~~~~~
We don't often care about the contents of coercions, so checking
coercions before making an equality constraint may be surprising.
But there are several cases we need to be wary of:
(1) When we're unifying a variable, we must make sure that the variable
appears nowhere on the RHS -- even in a coercion. Otherwise, we'll
create a loop.
(2) We must still make sure that no variable in a coercion is at too
high a level. But, when unifying, we can promote any variables we encounter.
Note [Promotion and level-checking]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
"Promotion" happens when we have this:
[W] w1: alpha[2] ~ Maybe beta[4]
Here we must NOT unify alpha := Maybe beta, because beta may turn out
to stand for a type involving some inner skolem. Yikes!
Skolem-escape. So instead we /promote/ beta, like this:
beta[4] := beta'[2]
[W] w1: alpha[2] ~ Maybe beta'[2]
Now we can unify alpha := Maybe beta', which might unlock other
constraints. But if some other constraint wants to unify beta with a
nested skolem, it'll get stuck with a skolem-escape error.
Now consider `w2` where a type family is involved (#22194):
[W] w2: alpha[2] ~ Maybe (F gamma beta[4])
In `w2`, it may or may not be the case that `beta` is level 2; suppose
we later discover gamma := Int, and type instance F Int _ = Int.
So, instead, we promote the entire funcion call:
[W] w2': alpha[2] ~ Maybe gamma[2]
[W] w3: gamma[2] ~ F gamma beta[4]
Now we can unify alpha := Maybe gamma, which is a Good Thng.
Wrinkle (W1)
There is an important wrinkle: /all this only applies when unifying/.
For example, suppose we have
[G] a[2] ~ Maybe b[4]
where 'a' is a skolem. This Given might arise from a GADT match, and
we can absolutely use it to rewrite locally. In fact we must do so:
that is how we exploit local knowledge about the outer skolem a[2].
This applies equally for a Wanted [W] a[2] ~ Maybe b[4]. Using it for
local rewriting is fine. (It's not clear to me that it is /useful/,
but it's fine anyway.)
So we only do the level-check in checkTyVar when /unifying/ not for
skolems (or untouchable unification variables).
Note [Family applications in canonical constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
A constraint with a type family application in the RHS needs special care.
This is dealt with by `checkFamApp`.
(CFA1) First, occurs checks. If we have
[G] a ~ Maybe (F (Maybe a))
[W] alpha ~ Maybe (F (Maybe alpha))
it looks as if we have an occurs check. But go read
Note [Type equality cycles] in GHC.Tc.Solver.Equality
The same considerations apply when the LHS is a type family:
[G] G a ~ Maybe (F (Maybe (G a)))
[W] G alpha ~ Maybe (F (Maybe (G alpha)))
(CFA2) Second, promotion. If we have (#22194)
[W] alpha[2] ~ Maybe (F beta[4])
it is wrong to promote beta. Instead we want to split to
[W] alpha[2] ~ Maybe gamma[2]
[W] gamma[2] ~ F beta[4]
See Note [Promotion and level-checking] above.
(CFA3) Third, concrete type variables. If we have
[W] alpha[conc] ~ Maybe (F tys)
we want to add an extra variable thus:
[W] alpha[conc] ~ Maybe gamma[conc]
[W] gamma[conc] ~ F tys
Now we can unify alpha, and that might unlock something else.
In all these cases we want to create a fresh type variable, and
emit a new equality connecting it to the type family application.
Once these three cases are dealt with, the `tef_fam_app` field of `TypeEqFlags`
says what to do:
(CFA4) `TEFA_Recurse` is straightforward: just recurse into the arguments,
BUT use `recurseIntoFamTyConApp` to record that we are now "under" a
type-family application; see `famAppArgFlags`.
(CFA5) `TEFA_Break` is the clever one. It does a two-step process:
(1) Recurse into the arguments with `recurseIntoFamTyConApp`.
(2) If any of the arguments fail (level-check error, occurs check,
concreteness failure), use the `FamAppBreaker` to create a cycle breaker.
Remarks:
* It would be possible to use Step (2) above always, skipping Step (1).
But this would create many unnecessary cycle-breaker variables.
This was the cause of #25933.
* This always cycle-breaks the /outermost/ family application.
If we have [W] alpha ~ Maybe (F (G alpha)):
- We'll use checkFamApp on `(F (G alpha))`
- In Step (1), `recurseIntoFamTyConApp` sets `tef_fam_app := TEFA_Recurse`,
before looking at the argument `(G alpha)`. So we will not cycle-break
the latter
- The occurs check will fire when we hit `alpha`
- `checkFamApp` on `(F (G alpha))` will see the failure and invoke
the `FamAppBreaker`.
* Step (1) may fail because of a level-check problem, which activates step (2).
This is what implements (CFA2).
-}
-------------------
checkTyConApp :: Monad m
=> TyEqFlags m a
-> TcType -> TyCon -> [TcType]
-> m (PuResult a Reduction)
checkTyConApp flags tc_app tc tys
| isTypeFamilyTyCon tc
, let arity = tyConArity tc
= if tys `lengthIs` arity
then checkFamApp flags tc_app tc tys -- Common case
else do { let (fun_args, extra_args) = splitAt (tyConArity tc) tys
fun_app = mkTyConApp tc fun_args
; fun_res <- checkFamApp flags fun_app tc fun_args
; extra_res <- mapCheck (check_ty_eq_rhs flags) extra_args
; return (mkAppRedns <$> fun_res <*> extra_res) }
| Just ty' <- rewriterView tc_app
-- e.g. S a where type S a = F [a]
-- or type S a = Int
-- See Note [Forgetful synonyms in checkTyConApp]
= check_ty_eq_rhs flags ty'
| not (isTauTyCon tc)
= return $ PuFail impredicativeProblem
| tefConcrete flags
, not (isConcreteTyCon tc)
= return $ PuFail (cteProblem cteConcrete)
| otherwise -- Recurse on arguments
= recurseIntoTyConApp flags tc tys
-------------------
recurseIntoTyConApp :: Monad m
=> TyEqFlags m a
-> TyCon -> [TcType]
-> m (PuResult a Reduction)
recurseIntoTyConApp flags tc tys
= do { tys_res <- mapCheck (check_ty_eq_rhs flags) tys
; return (mkTyConAppRedn Nominal tc <$> tys_res) }
recurseIntoFamTyConApp :: Monad m
=> TyEqFlags m a
-> TyCon -> [TcType]
-> m (PuResult a Reduction)
recurseIntoFamTyConApp flags tc tys
= recurseIntoTyConApp (famAppArgFlags flags) tc tys
-- famAppArgFlags: adjust flags when going under fam app
-------------------
-- | See Note [Family applications in canonical constraints]
checkFamApp :: forall m a.
Monad m
=> TyEqFlags m a
-> TcType -> TyCon -> [TcType] -- Saturated family application
-> m (PuResult a Reduction)
checkFamApp flags fam_app tc tys =
case flags of
TEFTyFam { tefTyFam_tyCon = lhs_tc, tefTyFam_args = lhs_tys
, tefTyFam_occursCheck = occ_prob }
| tcEqTyConApps lhs_tc lhs_tys tc tys
-- There is an occurs check F ty ~ ...(F ty)...
-- Do not recurse; see (CFA1) in Note [Family applications in canonical constraints]
-> do_not_recurse occ_prob
TEFTyVar { tefTyVar_concreteCheck = conc }
| wantConcreteCheck conc
-- We are checking for concreteness, e.g. kappa[conc] ~ ...(F ty)...
-- Do not recurse: type family applications are not concrete.
-- See (CFA3) in Note [Family applications in canonical constraints]
-> do_not_recurse cteConcrete
_ ->
-- We can recurse into the arguments of the type family application.
case fam_flags of
TEFA_Recurse ->
-- Simply recurse into the arguments.
-- See (CFA4) in Note [Family applications in canonical constraints]
recurseIntoFamTyConApp flags tc tys
TEFA_Break brk ->
-- Recurse into the arguments, and cycle-break if that fails
--
-- e.g. alpha[2] ~ Maybe (F beta[2]) No problem: just unify
-- alpha[2] ~ Maybe (F beta[4]) Level-check problem: break
--
-- NB: in the latter case, don't promote beta[4]; use 'recurseIntoFamTyConApp'
-- which modifies the flags with 'famAppArgFlags'.
do { rec_res <- recurseIntoFamTyConApp flags tc tys
; case rec_res of
PuOK {} -> return rec_res
PuFail {} -> famAppBreaker brk fam_app
}
where
fam_flags = tef_fam_app flags
do_not_recurse prob =
case fam_flags of
TEFA_Recurse -> return $ PuFail (cteProblem prob)
TEFA_Break brk -> famAppBreaker brk fam_app
-------------------
-- | The result of a single check in 'checkTyVar', such as a concreteness check
-- or a level check.
data TyVarCheckResult m where
-- | Check succeded; nothing else to do.
TyVarCheck_Success :: TyVarCheckResult m
-- | Check succeeded, but requires an additional promotion.
--
-- Invariant: at least one of the fields is not 'Nothing'.
TyVarCheck_Promote
:: Maybe TcLevel
-- ^ @Just lvl@ <=> 'TyVar' needs to be promoted to @lvl@
-> Maybe ConcreteTvOrigin
-- ^ @Just conc_orig@ <=> 'TyVar' needs to be make concrete
-> TyVarCheckResult TcM
-- | Check failed with some 'CheckTyEqProblem's.
--
-- Invariant: the 'CheckTyEqResult' is not 'cteOK'.
TyVarCheck_Error
:: CheckTyEqResult -> TyVarCheckResult m
instance Semigroup (TyVarCheckResult m) where
TyVarCheck_Success <> r = r
r <> TyVarCheck_Success = r
TyVarCheck_Error e1 <> TyVarCheck_Error e2 =
TyVarCheck_Error (e1 S.<> e2)
e@(TyVarCheck_Error {}) <> _ = e
_ <> e@(TyVarCheck_Error {}) = e
TyVarCheck_Promote l1 c1 <> TyVarCheck_Promote l2 c2 =
TyVarCheck_Promote
(combineMaybe minTcLevel l1 l2)
(combineMaybe const c1 c2) -- pick one 'ConcreteTvOrigin' arbitrarily
combineMaybe :: (a -> a -> a) -> Maybe a -> Maybe a -> Maybe a
combineMaybe _ Nothing r = r
combineMaybe _ r Nothing = r
combineMaybe f (Just a) (Just b) = Just (f a b)
instance Monoid (TyVarCheckResult m) where
mempty = TyVarCheck_Success
checkTyVar :: forall m a. Monad m => TyEqFlags m a -> TcTyVar -> m (PuResult a Reduction)
checkTyVar flags occ_tv
= case flags of
TEFTyFam {}
-> success -- Nothing to do if the LHS is a type-family
TEFTyVar { tefTyVar_occursCheck = occ, tefTyVar_levelCheck = lc
, tefTyVar_concreteCheck = conc }
-> do { already_promoted <- promotionDone occ_tv lc conc
; if already_promoted
then success
-- Already promoted; job done
-- Example alpha[2] ~ Maybe (beta[4], beta[4])
-- We promote the first occurrence, and then encounter it
-- a second time; we don't want to re-promote it!
-- Remember, the entire process started with a fully zonked type
-- so the only reason for a filled meta-tyvar is promotion
else do_rhs_checks
[ simpleOccursCheck occ occ_tv
, tyVarLevelCheck lc occ_tv
, tyVarConcretenessCheck conc occ_tv ]
}
where
success = return $ okCheckRefl (mkTyVarTy occ_tv)
-- Combine the results of individual checks. See 'TyVarCheckResult'.
do_rhs_checks :: [TyVarCheckResult m] -> m (PuResult a Reduction)
do_rhs_checks checks =
case mconcat checks of
TyVarCheck_Success -> success
TyVarCheck_Promote mb_lvl mb_conc -> promote flags mb_lvl mb_conc
TyVarCheck_Error cte_prob -> return $ PuFail cte_prob
---------------------
-- occ_tv is definitely a MetaTyVar; we need to promote it/make it concrete
promote :: TyEqFlags TcM a -> Maybe TcLevel -> Maybe ConcreteTvOrigin
-> TcM (PuResult a Reduction)
promote flags mb_lhs_tv_lvl mb_conc
| MetaTv { mtv_info = info_occ, mtv_tclvl = lvl_occ } <- tcTyVarDetails occ_tv
= do { let new_info | Just conc <- mb_conc = ConcreteTv conc
| otherwise = info_occ
new_lvl =
case mb_lhs_tv_lvl of
Nothing -> lvl_occ
Just lhs_tv_lvl -> lhs_tv_lvl `minTcLevel` lvl_occ
-- c[conc,3] ~ p[tau,2]: want to clone p:=p'[conc,2]
-- c[tau,2] ~ p[tau,3]: want to clone p:=p'[tau,2]
-- Check the kind of occ_tv
--
-- This is important for several reasons:
--
-- 1. To ensure there is no occurs check or skolem-escape
-- in the kind of occ_tv.
-- 2. If the LHS is a concrete type variable and the RHS is an
-- unfilled meta-tyvar, we need to ensure that the kind of
-- 'occ_tv' is concrete. Test cases: T23051, T23176.
; let occ_kind = tyVarKind occ_tv
; kind_result <- check_ty_eq_rhs flags occ_kind
; for kind_result $ \ kind_redn ->
do { let kind_co = reductionCoercion kind_redn
new_kind = reductionReducedType kind_redn
occ_tv' = setTyVarKind occ_tv new_kind
; new_tv_ty <- promote_meta_tyvar new_info new_lvl occ_tv'
; return $ mkGReflLeftRedn Nominal new_tv_ty (mkSymCo kind_co)
} }
| otherwise = pprPanic "promote" (ppr occ_tv)
---------------------
promotionDone :: Monad m => TcTyVar -> LevelCheck m -> ConcreteCheck m -> m Bool
promotionDone occ_tv (LC_Promote {}) _ = isFilledMetaTyVar occ_tv
promotionDone occ_tv _ (CC_Promote {}) = isFilledMetaTyVar occ_tv
promotionDone _ _ _ = return False
---------------------
simpleOccursCheck :: OccursCheck -> TcTyVar -> TyVarCheckResult m
-- ^ The "interpreter" for 'OccursCheck'
simpleOccursCheck OC_None _
= TyVarCheck_Success
simpleOccursCheck (OC_Check lhs_tv occ_prob) occ_tv
| lhs_tv == tyVarName occ_tv || check_kind (tyVarKind occ_tv)
= TyVarCheck_Error (cteProblem occ_prob)
| otherwise
= TyVarCheck_Success
where
(check_kind, _) = mkOccFolders lhs_tv
-------------------------
tyVarLevelCheck :: LevelCheck m -> TcTyVar -> TyVarCheckResult m
-- ^ The "interpreter" for 'LevelCheck'
tyVarLevelCheck LC_None _ = TyVarCheck_Success
tyVarLevelCheck lc occ_tv
| not occ_is_deeper
= TyVarCheck_Success
| isSkolemTyVar occ_tv
= TyVarCheck_Error (cteProblem cteSkolemEscape)
| otherwise
= case lc of
LC_Check { lc_lenient = lenient } ->
if lenient
then TyVarCheck_Success
else TyVarCheck_Error (cteProblem cteSkolemEscape)
LC_Promote { lc_lvlp = lhs_tv_lvl } ->
TyVarCheck_Promote (Just lhs_tv_lvl) Nothing
where
occ_is_deeper =
case lc of
LC_Check { lc_lvlc = lhs_tv_lvl } ->
tcTyVarLevel occ_tv `strictlyDeeperThan` lhs_tv_lvl
LC_Promote { lc_lvlp = lhs_tv_lvl } ->
tcTyVarLevel occ_tv `strictlyDeeperThan` lhs_tv_lvl
-------------------------
tyVarConcretenessCheck :: ConcreteCheck m -> TcTyVar -> TyVarCheckResult m
-- ^ The "interpreter" for 'ConcreteCheck'
tyVarConcretenessCheck cc occ_tv
| CC_None <- cc -- No concreteness checks => ok
= TyVarCheck_Success
| isConcreteTyVar occ_tv -- It's concrete already => ok
= TyVarCheck_Success
| not (isMetaTyVar occ_tv) -- Not a meta-tyvar => error
= TyVarCheck_Error (cteProblem cteConcrete)
| otherwise -- It's a non-concrete meta-tyvar
= case cc of
CC_Check -> TyVarCheck_Success
CC_Promote conc_orig -> TyVarCheck_Promote Nothing (Just conc_orig)
-- CC_None is dealt with at the top
-------------------------
checkPromoteFreeVars :: CheckTyEqProblem -- What occurs check problem to report
-> Name -> TcLevel
-> TyCoVarSet -> TcM CheckTyEqResult
-- Check this set of TyCoVars for
-- (a) occurs check
-- (b) promote if necessary, or report skolem escape
checkPromoteFreeVars occ_prob lhs_tv lhs_tv_lvl vs
= do { oks <- mapM do_one (nonDetEltsUniqSet vs)
; return (mconcat oks) }
where
do_one :: TyCoVar -> TcM CheckTyEqResult
do_one v | isCoVar v = return cteOK
| tyVarName v == lhs_tv = return (cteProblem occ_prob)
| no_promotion = return cteOK
| not (isMetaTyVar v) = return (cteProblem cteSkolemEscape)
| otherwise = promote_one v
where
no_promotion = not (tcTyVarLevel v `strictlyDeeperThan` lhs_tv_lvl)
-- isCoVar case: coercion variables are not an escape risk
-- If an implication binds a coercion variable, it'll have equalities,
-- so the "intervening given equalities" test above will catch it
-- Coercion holes get filled with coercions, so again no problem.
promote_one :: TyVar -> TcM CheckTyEqResult
promote_one tv =
do { _ <- promote_meta_tyvar TauTv lhs_tv_lvl tv
; return cteOK }
promote_meta_tyvar :: MetaInfo -> TcLevel -> TcTyVar -> TcM TcType
promote_meta_tyvar info dest_lvl occ_tv
= do { -- Check whether occ_tv is already unified. The rhs-type
-- started zonked, but we may have promoted one of its type
-- variables, and we then encounter it for the second time.
-- But if so, it'll definitely be another already-checked TyVar
mb_filled <- isFilledMetaTyVar_maybe occ_tv
; case mb_filled of {
Just ty -> return ty ;
Nothing ->
-- OK, not done already, so clone/promote it
do { new_tv <- cloneMetaTyVarWithInfo info dest_lvl occ_tv
; liftZonkM $ writeMetaTyVar occ_tv (mkTyVarTy new_tv)
; traceTc "promoteTyVar" (ppr occ_tv <+> text "-->" <+> ppr new_tv)
; return (mkTyVarTy new_tv) } } }
-------------------------
touchabilityTest :: TcLevel -> TcTyVar -> Bool
-- ^ This is the key test for untouchability:
-- See Note [Unification preconditions] in GHC.Tc.Utils.Unify
-- and Note [Solve by unification] in GHC.Tc.Solver.Equality
--
-- @True@ <=> the variable is touchable
touchabilityTest given_eq_lvl tv
| MetaTv { mtv_tclvl = tv_lvl } <- tcTyVarDetails tv
= tv_lvl `deeperThanOrSame` given_eq_lvl
| otherwise
= False
-------------------------
-- | checkTopShape checks (TYVAR-TV)
-- Note [Unification preconditions]; returns True if these conditions
-- are satisfied. But see the Note for other preconditions, too.
checkTopShape :: MetaInfo -> TcType -> Bool
checkTopShape info xi
= case info of
TyVarTv ->
case getTyVar_maybe xi of -- Looks through type synonyms
Nothing -> False
Just tv -> case tcTyVarDetails tv of -- (TYVAR-TV) wrinkle
SkolemTv {} -> True
RuntimeUnk -> True
MetaTv { mtv_info = TyVarTv } -> True
_ -> False
CycleBreakerTv -> False -- We never unify these
_ -> True
--------------------------------------------------------------------------------
-- Making a type concrete.
-- | Try to turn the provided type into a concrete type, by ensuring
-- unfilled metavariables are appropriately marked as concrete.
--
-- Returns a coercion whose RHS is a zonked type which is "as concrete as possible",
-- and a collection of Wanted equality constraints that are necessary to make
-- the type concrete.
--
-- For example, for an input @TYPE a[sk]@ we will return a coercion with RHS
-- @TYPE gamma[conc]@ together with the Wanted equality constraint @a ~# gamma@.
--
-- INVARIANT: the RHS type of the returned coercion is equal to the input type,
-- up to zonking and the returned Wanted equality constraints.
--
-- INVARIANT: if this function returns an empty list of constraints
-- then the RHS type of the returned coercion is concrete,
-- in the sense of Note [Concrete types].
makeTypeConcrete :: FastString -> ConcreteTvOrigin
-> TcType -> TcM (TcCoercion, Cts)
makeTypeConcrete occ_fs conc_orig ty =
do { traceTc "makeTypeConcrete {" $
vcat [ text "ty:" <+> ppr ty ]
-- To make a type 'ty' concrete, we query what would happen were we
-- to try unifying
--
-- alpha[conc] ~# ty
--
-- for a fresh concrete metavariable 'alpha'.
--
-- We do this by calling 'checkTyEqRhs' with suitable 'TyEqFlags'.
-- NB: we don't actually need to create a fresh concrete metavariable
-- in order to call 'checkTyEqRhs'.
; let ty_eq_flags =
TEFTyVar
{ tefTyVar_occursCheck = OC_None
-- LHS is a fresh meta-tyvar: no occurs check needed
, tefTyVar_levelCheck = LC_None
, tefTyVar_concreteCheck = CC_Promote conc_orig
, tef_fam_app = TEFA_Recurse
}
-- NB: 'checkTyEqRhs' expects a fully zonked type as input.
; ty' <- liftZonkM $ zonkTcType ty
; pu_res <- checkTyEqRhs @TcM @() ty_eq_flags ty'
-- NB: 'checkTyEqRhs' will also check the kind, thus upholding the
-- invariant that the kind of a concrete type must also be a concrete type.
; (cts, final_co) <-
case pu_res of
PuOK _ redn ->
do { traceTc "makeTypeConcrete: unifier success" $
vcat [ text "ty:" <+> ppr ty <+> dcolon <+> ppr (typeKind ty)
, text "redn:" <+> ppr redn
]
; return (emptyBag, mkSymCo $ reductionCoercion redn)
-- NB: the unifier returns a 'Reduction' with the concrete
-- type on the left, but we want a coercion with it on the
-- right; so we use 'mkSymCo'.
}
PuFail _prob ->
do { traceTc "makeTypeConcrete: unifier failure" $
vcat [ text "ty:" <+> ppr ty <+> dcolon <+> ppr (typeKind ty)
, text "problem:" <+> ppr _prob
]
-- We failed to make 'ty' concrete. In order to continue
-- typechecking, we proceed as follows:
--
-- - create a new concrete metavariable alpha[conc]
-- - emit the equality @ty ~# alpha[conc]@.
--
-- This equality will eventually get reported as insoluble
-- to the user.
-- The kind of a concrete metavariable must itself be concrete,
-- so we need to do a concreteness check on the kind first.
; let ki = typeKind ty
; (kind_co, kind_cts) <-
if isConcreteType ki
then return (mkNomReflCo ki, emptyBag)
else makeTypeConcrete occ_fs conc_orig ki
-- Now create the new concrete metavariable.
; conc_tv <- newConcreteTyVar conc_orig occ_fs (coercionRKind kind_co)
; let conc_ty = mkTyVarTy conc_tv
pty = mkEqPredRole Nominal ty' conc_ty
; hole <- newCoercionHole pty
; loc <- getCtLocM orig (Just KindLevel)
; let ct = mkNonCanonical $ CtWanted
$ WantedCt { ctev_pred = pty
, ctev_dest = HoleDest hole
, ctev_loc = loc
, ctev_rewriters = emptyRewriterSet }
; return (kind_cts S.<> unitBag ct, HoleCo hole)
}
; traceTc "makeTypeConcrete }" $
vcat [ text "ty :" <+> _ppr_ty ty
, text "ty':" <+> _ppr_ty ty'
, text "final_co:" <+> _ppr_co final_co ]
; return (final_co, cts)
}
where
_ppr_ty ty = ppr ty <+> dcolon <+> ppr (typeKind ty)
_ppr_co co = ppr co <+> dcolon <+> parens (_ppr_ty (coercionLKind co)) <+> text "~#" <+> parens (_ppr_ty (coercionRKind co))
orig :: CtOrigin
orig = case conc_orig of
ConcreteFRR frr_orig -> FRROrigin frr_orig
{- *********************************************************************
* *
mightEqualLater
* *
********************************************************************* -}
{- Note [What might equal later?]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We must determine whether a Given might later equal a Wanted:
see Note [Instance and Given overlap] in GHC.Tc.Solver.Dict
We definitely need to account for the possibility that any metavariable might be
arbitrarily instantiated. Yet we do *not* want to allow skolems 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[tau] ~? C Int
That Given certainly might match later.
2. C a[sk] ~? C Int
No. No new givens are going to arise that will get the `a` to rewrite
to Int. Example:
f :: forall a. C a => blah
f = rhs -- Gives rise to [W] C Int
It would be silly to fail to solve ([W] C Int), just because we have
([G] C a) in the Givens!
3. C alpha[tv] ~? C Int
In this variant of (1) that alpha[tv] is a TyVarTv, unifiable only with
other type /variables/. It cannot equal Int later.
4. C (F alpha[tau]) ~? 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[sk]) ~? C Int. For example
f :: forall a. C (F a) => blah
f = rhs -- Gives rise to [W] C Int
No, this won't match later. If we could rewrite (F a), we would
have by now. But see also Red Herring below.
This arises in instance decls too. For example in GHC.Core.Ppr we see
instance Outputable (XTickishId pass)
=> Outputable (GenTickish pass) where
If we have [W] Outputable Int in the body, we don't want to fail to solve
it because (XTickishId pass) might simplify to Int.
7. C (Maybe alpha[tau]) ~? C alpha[tau]
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 led to #19107, tested in
typecheck/should_compile/T19107.
8. C alpha[cbv] ~? C Int
where alpha[cbv] = F a
The alpha[cbv] is a cycle-breaker var which stands for F a. See
Note [Type equality cycles] in GHC.Tc.Solver.Equality
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 alpha[cbv] ~? C Int
where alpha[cbv] = F beta[tau]
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[tau], Int) ~? C (Bool, F alpha[tau])
This cannot equal later, because F alpha would have to equal both Bool and
Int.
To deal with type family applications, we use the "fine-grained" Core unifier.
See Note [Apartness and type families] in GHC.Core.Unify, controlled
by the `bind_fam :: BindFamFun` function defined in `mightEqualLater`.
One tricky point: 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. However: 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. See the
use of `mentions_meta_ty_var` in `mightEqualLater`.
This treatment fixes #18910 and is tested in
typecheck/should_compile/InstanceGivenOverlap{,2}
Red Herring
~~~~~~~~~~~
In #21208, we have this scenario:
instance forall b. C b
[G] C a[sk]
[W] C (F a[sk])
What should we do with that wanted? According to the logic above, the Given
cannot match later (this is example 6), and so we use the global instance.
But wait, you say: What if we learn later (say by a future type instance F a = a)
that F a unifies with a? That looks like the Given might really match later!
This mechanism described in this Note is *not* about this kind of situation, however.
It is all asking whether a Given might match the Wanted *in this run of the solver*.
It is *not* about whether a variable might be instantiated so that the Given matches,
or whether a type instance introduced in a downstream module might make the Given match.
The reason we care about what might match later is only about avoiding order-dependence.
That is, we don't want to commit to a course of action that depends on seeing constraints
in a certain order. But an instantiation of a variable and a later type instance
don't introduce order dependency in this way, and so mightMatchLater is right to ignore
these possibilities.
Here is an example, with no type families, that is perhaps clearer:
instance forall b. C (Maybe b)
[G] C (Maybe Int)
[W] C (Maybe a)
What to do? We *might* say that the Given could match later and should thus block
us from using the global instance. But we don't do this. Instead, we rely on class
coherence to say that choosing the global instance is just fine, even if later we
call a function with (a := Int). After all, in this run of the solver, [G] C (Maybe Int)
will definitely never match [W] C (Maybe a). (Recall that we process Givens before
Wanteds, so there is no [G] a ~ Int hanging about unseen.)
Interestingly, in the first case (from #21208), the behavior changed between
GHC 8.10.7 and GHC 9.2, with the latter behaving correctly and the former
reporting overlapping instances.
Test case: typecheck/should_compile/T21208.
-}
--------------------------------------------------------------------------------
-- mightEqualLater
mightEqualLater :: InertSet -> TcPredType -> CtLoc -> TcPredType -> CtLoc -> Maybe Subst
-- See Note [What might equal later?]
-- Used to implement logic in Note [Instance and Given overlap] in GHC.Tc.Solver.Dict
mightEqualLater inert_set given_pred given_loc wanted_pred wanted_loc
| prohibitedSuperClassSolve given_loc wanted_loc
= Nothing
| otherwise
= case tcUnifyTysFG bind_fam bind_tv [given_pred] [wanted_pred] of
Unifiable subst
-> Just subst
MaybeApart reason subst
| MARInfinite <- reason -- see Example 7 in the Note.
-> Nothing
| otherwise
-> Just subst
SurelyApart -> Nothing
where
bind_tv :: BindTvFun
bind_tv tv rhs_ty
| MetaTv { mtv_info = info } <- tcTyVarDetails tv
, ok_shape tv info rhs_ty
, can_unify tv rhs_ty
= BindMe
| otherwise
= DontBindMe
bind_fam :: BindFamFun
-- See Examples (4), (5), and (6) from the Note, especially (6)
bind_fam _fam_tc fam_args _rhs
| anyFreeVarsOfTypes mentions_meta_ty_var fam_args = BindMe
| otherwise = DontBindMe
can_unify :: TcTyVar -> TcType -> Bool
can_unify tv rhs_ty
-- See Note [Use checkTyEqRhs in mightEqualLater]
| PuOK {} <- runIdentity $ checkTyEqRhs (pureTyEqFlags_LHSMetaTyVar tv) rhs_ty
= True
| otherwise
= False
-- 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 -> anyFreeVarsOfType mentions_meta_ty_var
(lookupCycleBreakerVar tv inert_set)
_ -> True
| otherwise
= False
-- Like checkTopShape, but allows cbv variables to unify
ok_shape :: TcTyVar -> MetaInfo -> Type -> Bool
ok_shape _lhs_tv TyVarTv rhs_ty -- see Example 3 from the Note
| Just rhs_tv <- getTyVar_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
ok_shape lhs_tv _other _rhs_ty = mentions_meta_ty_var lhs_tv
{- Note [Use checkTyEqRhs in mightEqualLater]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In 'mightEqualLater', we are checking whether two types ty1 and ty2 might become
equal after further unifications. To do this, we call the pure unifier, via
the function 'tcUnifyTysFG'. However, it is important that we don't return
false positives:
1. Level check: ty1 = alpha[tau:1] /~ ty2 = k[sk:3]
2. Occurs check: ty1 = beta[tau] /~ ty2 = Maybe beta[tau]
3. Concreteness: ty1 = kappa[conc] /~ ty2 = k[sk].
In these examples, ty1 and ty2 cannot unify; to inform the pure unifier of this
fact, we use 'checkTyEqRhs' to provide the 'BindTvFun'.
Failing to account for this caused #25744:
work item = [W] w1: Eq (g[sk:1] x[sk:3])
inerts = { [G] g1: Eq x[sk:3]
, [G] g2: forall y. Eq y => Eq (g[sk:1] y)
, [G] g3: Eq (f[tau:1] v[tau:1]) }
We want to solve w1 using g1 and g2. The presence of g3 is irrelevant, because
we cannot unify 'f[tau:1] v[tau:1]' and 'g[sk:1] x[sk:3]' due to skolem escape.
-}
-- | Return the type family application a CycleBreakerTv maps to.
lookupCycleBreakerVar :: TcTyVar -- ^ cbv, must be a CycleBreakerTv
-> InertSet
-> TcType -- ^ type family application the cbv maps to
lookupCycleBreakerVar cbv (IS { inert_cycle_breakers = cbvs_stack })
-- This function looks at every environment in the stack. This is necessary
-- to avoid #20231. This function (and its one usage site) is the only reason
-- that we store a stack instead of just the top environment.
| Just tyfam_app <- assert (isCycleBreakerTyVar cbv) $
firstJusts (NE.map (lookupBag cbv) cbvs_stack)
= tyfam_app
| otherwise
= pprPanic "lookupCycleBreakerVar found an unbound cycle breaker" (ppr cbv $$ ppr cbvs_stack)
--------------------------------------------------------------------------------